The AdaCore Blog https://blog.adacore.com/ An insight into the AdaCore ecosystem en-us Mon, 25 Oct 2021 02:21:32 -0400 Mon, 25 Oct 2021 02:21:32 -0400 Starting micro-controller Ada drivers in the Alire ecosystem https://blog.adacore.com/starting-micro-controller-ada-drivers-in-the-alire-ecosystem Mon, 18 Oct 2021 04:12:00 -0400 Fabien Chouteau https://blog.adacore.com/starting-micro-controller-ada-drivers-in-the-alire-ecosystem $alr init --lib samd21_hal$ cd samd21_hal $alr with gnat_arm_elf # Add a dependency on the arm-elf compiler for Target use "arm-elf"; for Runtime ("Ada") use "zfp-cortex-m0p"; $ mv src/samd21_hal.ads src/sam.ads
$sed -i 's/Samd21_Hal/SAM/g' src/sam.ads $ alr build
$wget http://packs.download.atmel.com/Atmel.SAMD21_DFP.1.3.395.atpack$ unzip Atmel.SAMD21_DFP.1.3.395.atpack -d samd21_atpack
$svd2ada samd21_atpack/samd21a/svd/ATSAMD21G18A.svd --boolean -o src -p SAM_SVD --base-types-package HAL --gen-uint-always $ alr with hal
$alr build $ cd ..
$alr init --bin metro_m0_example$ cd metro_m0_example
$alr with samd21_hal --use=../samd21_hal for Target use "arm-elf"; for Runtime ("Ada") use "zfp-cortex-m0p"; $ alr build
[...]
warning: cannot find entry symbol _start; defaulting to 0000000000008000
package Device_Configuration is
for CPU_Name use "ARM Cortex-M0P";
for Float_Handling use "soft";

for Number_Of_Interrupts use "42";

for Memories use ("RAM", "FLASH");

--  Specify from which memory bank the program will load
for Boot_Memory use "FLASH";

--  Specification of the RAM
for Mem_Kind ("RAM") use "ram";
for Size ("RAM") use "0x8000";

--  Specification of the FLASH
for Mem_Kind ("FLASH") use "rom";
for Size ("FLASH") use "0x40000";
end Device_Configuration;
$cd ..$ alr get --build startup_gen
$cd metro_m0_example$ eval alr printenv
$../startup_gen_21.0.0_75bdb097/startup-gen -P metro_m0_example.gpr -l src/link.ld -s src/crt0.S CPU: ARM Cortex-M0P Float_Handling: SOFT Name : RAM Address : 0x20000000 Size : 0x8000 Kind : RAM Name : FLASH Address : 0x08000000 Size : 0x40000 Kind : ROM for Languages use ("Ada", "ASM_CPP"); package Linker is for Switches ("Ada") use ("-T", Project'Project_dir & "/src/link.ld", "-Wl,--print-memory-usage", "-Wl,--gc-sections"); end Linker; $ alr build
[...]
Memory region         Used Size  Region Size  %age Used
FLASH:         756 B       256 KB      0.29%
RAM:        4120 B        32 KB     12.57%
$cd ..$ alr init --lib metro_m0_bsp
$cd metro_m0_bsp$ alr with samd21_hal --use=../samd21_hal
$cd ../metro_m0_example/$ alr with metro_m0_bsp --use=../metro_m0_bsp
alr build ]]> Enhancing the Security of a TCP Stack with SPARK https://blog.adacore.com/enhancing-the-security-of-a-tcp-stack-with-spark Tue, 12 Oct 2021 04:21:00 -0400 Yannick Moy https://blog.adacore.com/enhancing-the-security-of-a-tcp-stack-with-spark You've probably never heard of CycloneTCP, an open source dual IPv4/IPv6 stack dedicated to embedded applications. That may be because people don't find and publish vulnerabilities for this stack. The quality of CycloneTCP is even acknowledged by the AMNESIA:33 report, which classifies it as one of the most resilient TCP/IP stacks. To go beyond the usual best development practices and use of industrial testsuites, the developers of CycloneTCP at Oryx Embedded partnered with AdaCore. We worked together to replace the TCP part of the C codebase with SPARK code, and used the SPARK tools to prove both that the code is not vulnerable to the usual runtime errors (like buffer overflow) and that it correctly implements the TCP automaton specified in RFC 793. As part of this work, we found two subtle bugs related to memory management and concurrency. For more details, see our article or watch our online presentation on October 20th at IEEE SecDev 2021. ]]> Task Suspension with a Timeout in Ravenscar/Jorvik https://blog.adacore.com/blocking-with-a-timeout-in-ravenscar-jorvik Tue, 05 Oct 2021 06:25:00 -0400 Pat Rogers https://blog.adacore.com/blocking-with-a-timeout-in-ravenscar-jorvik This blog entry shows how to define an abstract data type that allows tasks to block on objects of the type, waiting for resumption signals from other components, for at most a specified amount of time per object. This "timeout" capability has been available in Ada from the beginning, via select statements containing timed entry calls. But what about developers working within the Ravenscar and Jorvik tasking subsets? Select statements and timed calls are not included within either profile. This new abstraction will provide some of the functionality of timed entry calls, with an implementation consistent with the Ravenscar and Jorvik subsets. In a previous blog entry we showed how to have a set of "conditions" that tasks can await, suspended, eventually to be awakened when "signaled" by some other task. Callers could await one of several conditions at the same time. However, these waiting callers blocked indefinitely, without a timeout option. That's often appropriate, but not in all cases. We will now take a different approach, defining a simpler version of "conditions" more like a condition variable or semaphore. Tasks still call Wait, now for a specific condition object passed as a parameter, and also specify how long they are willing to be suspended, waiting for some other task or interrupt handler to call Signal for that same object. The reason these blog entries focus on "conditions" is that "condition synchronization" is one of the two forms of synchronization required for concurrent programming. (Mutual exclusion is the other form.) For example, a consumer task must wait until a shared buffer is not empty before it can remove a value from the buffer. Likewise, a producer task inserting items into the shared buffer must wait until the buffer is not full. Protected entry barriers exist for the sake of expressing these sorts of Boolean conditions. However, as mentioned, for Ravenscar and Jorvik use we need an alternative mechanism. You should understand that this mechanism does not provide the full capabilities of timed entry calls. Condition objects are not entries, they are just flags, or events, and as such do not include an entry body that can provide application-specific functionality. Unlike a timed entry call, a call to Wait is not a request for a service to be provided (strictly, started to be provided) within a given time. Instead, a call to Wait requests notification that a condition has been satisfied, or if you like, an event has occurred, within the specified time. The analogue in full Ada would be a select statement containing a timed call to a protected entry with a null body. Any application-specific functionality corresponding with the Wait call's return -- that which a protected entry body would otherwise provide -- must be programmed separately from the call itself. Ada defines some standard lower-level facilities that can be used to define synchronization mechanisms, as well as used directly by applications. The most important of these are within the subsets defined by Ravenscar and Jorvik. We will use some of them to define the new capability. Having shown how to implement the facility within the Ravenscar and Jorvik subsets, we then provide a demonstration on bare-metal hardware. The API As usual, the new mechanism is designed as an abstract data type (ADT), hence a private type in Ada. As a synchronization mechanism, clients of the type have no business doing assignment between objects of this type, and language-defined equality on such objects makes no sense. Therefore, the type is limited as well as private. (As you will see, there is another good reason for the type to be limited.) The enclosing package, type declaration, and primitive subprogram declarations are as follows: package Timed_Conditions is type Timed_Condition is limited private; procedure Wait (This : in out Timed_Condition; Deadline : Time; Timed_Out : out Boolean); procedure Wait (This : in out Timed_Condition; Interval : Time_Span; Timed_Out : out Boolean); procedure Signal (This : in out Timed_Condition); private ... end Timed_Conditions; With this API clients can declare objects of type Timed_Condition and can pass them to calls to Wait and Signal. Procedure Wait is overloaded to allow expression of the timeout value either in terms of an absolute time, i.e., a point on the timeline, or a time interval. With the latter, the actual timeout is the sum of the time when the call takes place and the interval specified. Tasks calling Wait for a given Timed_Condition object suspend until either the time is reached or a call to Signal takes place for the same object. In both cases Wait returns a Boolean value indicating whether or not the call has returned due to the expiration of the time specified. For example, we could declare an object of this type like so: with Timed_Conditions; use Timed_Conditions; package User_Button is Pressed : Timed_Condition; ... end User_Button; Let's say, arbitrarily, that we want to wait at most 2 seconds for Pressed to be signaled. The task in the code below does so: with User_Button; ... task body Waiter is Time_Expired : Boolean; Timeout : constant Time_Span := Milliseconds (2_000); -- arbitrary begin loop Wait (User_Button.Pressed, Timeout, Time_Expired); if Time_Expired then ... ... end Waiter; The Implementation As hinted earlier, Ada defines standard lower-level mechanisms useful for building new kinds of concurrency constructs. We will use two: "timing events" and "suspension objects," both appearing in the full definition of the ADT in the private part of the package. The type Timing_Event is language-defined in the Ada.Real_Time.Timing_Events package. Objects of this type allow clients to specify a time when an "event" should occur. When that time is reached a user-defined protected procedure "handler" is invoked by the runtime library, performing whatever functional steps are required to implement the event. Clients may also cancel the future event, such that the handler will not be triggered. As you can imagine, this type will provide much of our timeout implementation. The pertinent parts of the API are as follows: package Ada.Real_Time.Timing_Events is type Timing_Event is tagged limited private; type Timing_Event_Handler is access protected procedure (Event : in out Timing_Event); procedure Set_Handler (Event : in out Timing_Event; At_Time : Time; Handler : Timing_Event_Handler); procedure Cancel_Handler (Event : in out Timing_Event; Cancelled : out Boolean); ... private ... end Ada.Real_Time.Timing_Events; Procedure Set_Handler allows clients to set a time when the given Timing_Event object is to be triggered, and, as well, to specify a pointer to the protected procedure to be invoked when the time is reached. Procedure Set_Handler is overloaded for convenience, the difference being a parameter of type Time_Span instead of type Time. Note the formal parameter defined for the protected procedure handler, designated by the Timing_Event_Handler access type. Any handler must be a protected procedure with a conforming formal parameter profile. Procedure Cancel_Handler cancels the timeout trigger for the given Timing_Event object. On return from the call the parameter Cancelled is True if the object was set prior to it being cancelled; otherwise, on return it is False. An object being "set" means that a timeout was pending and a pointer to a handler was currently assigned. The other required lower-level mechanism, "suspension objects," is provided by the type Suspension_Object declared in the Ada.Synchronous_Task_Control package. The pertinent parts of that package are as follows: package Ada.Synchronous_Task_Control is type Suspension_Object is limited private; procedure Set_True (S : in out Suspension_Object); procedure Suspend_Until_True (S : in out Suspension_Object); ... private ... end Ada.Synchronous_Task_Control; A Suspension_Object variable amounts to a thread-safe Boolean flag. Clients can call Set_True and Set_False to assign the values. Most significantly, via procedure Suspend_Until_True a client task can suspend itself until the specified flag becomes True. However, at most one task can be suspended on a given Suspension_Object variable at any given moment. Violations of that constraint raise Program_Error. Suspension_Object variables are initially False, automatically, and are set back to False automatically on return from a call to Suspend_Until_True. As a result, in typical code Set_False is not used. The operations Set_True and Set_False are atomic with respect to each other and with respect to Suspend_Until_True. The full declaration of our ADT using these two facilities is as follows: with Ada.Real_Time; use Ada.Real_Time; with Ada.Real_Time.Timing_Events; use Ada.Real_Time.Timing_Events; with Ada.Synchronous_Task_Control; use Ada.Synchronous_Task_Control; package Timed_Conditions is type Timed_Condition is limited private; procedure Wait (This : in out Timed_Condition; Deadline : Time; Timed_Out : out Boolean); procedure Wait (This : in out Timed_Condition; Interval : Time_Span; Timed_Out : out Boolean); procedure Signal (This : in out Timed_Condition); private type Timed_Condition is new Timing_Event with record Timed_Out : Boolean := False; Caller_Unblocked : Suspension_Object; end record; protected Timeout_Handler is pragma Interrupt_Priority; procedure Signal_Timeout (Event : in out Timing_Event); end Timeout_Handler; -- A shared, global PO defining the timing event handler procedure. All -- objects of type Timed_Condition use this one handler. Each execution of -- the procedure will necessarily execute at Interrupt_Priority'Last, so -- there's no reason to have a handler per-object. end Timed_Conditions; Our type Timed_Condition is visible to clients as a limited private type, so they must use it accordingly. The full view of the type in the package private part, however, indicates that much more is possible. In particular, the full type declaration in the private part extends type Timing_Event to define the Timed_Condition type. As a result, the new type inherits all the Timing_Event capabilities, and is a tagged type because Timing_Event is tagged. However, by design, neither the inherited operations nor the tagged nature are made part of the client API. We only want Timed_Condition clients to have the Wait and Signal operations. Completing the type declaration via inheritance in the private part of the package, rather than the public part, achieves that effect. Clients only have compile-time visibility to the partial view defined before the package private part. In contrast, the private part and package body have the full view, so the inherited operations are available there and will provide most of our timeout semantics. Furthermore, our extended type includes a Boolean component indicating whether a timeout occurred, and a Suspension_Object component used to block and unblock caller tasks. We made Timed_Condition a limited type in the visible part of the package (the client's partial view) for the reasons stated initially. In fact, the language requires us to do so, because the full type declaration in the private part (the full view) is itself limited. That required correspondence between the partial and full view makes sense because the client's view must be realistic with regard to the operations possible. If the type really is limited, as defined by the full view, then assignment really isn't possible. It wouldn't make sense for the client's view to indicate that assignment is possible if it really isn't. (By the same token, if the full view is not limited, the partial view is not required to be limited, but can be. If the partial view is limited but the full view is not, clients simply cannot do something that the full view allows within the package, i.e., assignment.) So, why is the full view of type Timed_Condition limited, even though the reserved word doesn't appear in our full view? It's because we are extending a limited type. Our new package is a client of Ada.Real_Time.Timing_Events so we have the partial view of type Timing_Event. That partial view is tagged and limited. Therefore any extension is also tagged and limited. In addition to the completion for Timed_Condition, the private part of the package also declares a single protected object, the Timeout_Handler. This protected object declares the protected procedure that will be invoked whenever any Timing_Condition object has timed out. (Note the required formal parameter's type. More on that in a moment.) When the Ceiling_Locking protocol is applied, as it is in both Ravenscar and Jorvik, the language requires Timing_Event handlers to execute at priority System.Interrupt_Priority'Last. The pragma Interrupt_Priority achieves that effect. (The expectation is that timeout handlers are executed directly by the clock interrupt handler.) It may seem surprising to have a single handler routine shared amongst all Timed_Condition objects. This approach works for a few reasons. First, the formal parameter to the handler gives us the specific object that has been triggered. Second, and most important, under these two profiles all handlers for Timing_Events must execute at a priority of System.Interrupt_Priority'Last, so all handlers will execute atomically, not concurrently. Therefore there is no benefit to having a dedicated protected object per Timing_Condition object. Given that full definition, here is the corresponding package body: package body Timed_Conditions is ---------- -- Wait -- ---------- procedure Wait (This : in out Timed_Condition; Deadline : Time; Timed_Out : out Boolean) is begin This.Set_Handler (Deadline, Timeout_Handler.Signal_Timeout'Access); Suspend_Until_True (This.Caller_Unblocked); Wait.Timed_Out := This.Timed_Out; end Wait; ---------- -- Wait -- ---------- procedure Wait (This : in out Timed_Condition; Interval : Time_Span; Timed_Out : out Boolean) is begin Wait (This, Clock + Interval, Timed_Out); end Wait; ------------ -- Signal -- ------------ procedure Signal (This : in out Timed_Condition) is Handler_Was_Set : Boolean; begin This.Cancel_Handler (Handler_Was_Set); if Handler_Was_Set then -- a caller to Wait is suspended This.Timed_Out := False; Set_True (This.Caller_Unblocked); end if; end Signal; --------------------- -- Timeout_Handler -- --------------------- protected body Timeout_Handler is -------------------- -- Signal_Timeout -- -------------------- procedure Signal_Timeout (Event : in out Timing_Event) is This : Timed_Condition renames Timed_Condition (Timing_Event'Class (Event)); begin This.Timed_Out := True; Set_True (This.Caller_Unblocked); -- note: Event's pointer to a handler becomes null automatically end Signal_Timeout; end Timeout_Handler; end Timed_Conditions; When called, procedure Wait sets a timeout deadline for the specified Timed_Condition, along with a pointer to the shared Signal_Timeout handler, and then suspends the caller. If the time expires, Signal_Timeout sets the Boolean Timed_Out flag to True and then unblocks the suspended caller in Wait. If, on the other hand, procedure Signal is called prior to the timeout, the timeout is canceled, Timed_Out is set to False, and again the caller in Wait is unblocked. In either case the Wait caller is unblocked and the Caller_Unblocked variable goes back to False automatically. (It is False initially, automatically.) At that point the internal Timed_Out Boolean flag can be assigned to the Timed_Out formal parameter. Wait then exits. Note that Signal could be called before a call to Wait has occurred for the same Timed_Condition object. And of course, it might be called after a timeout has expired. Therefore, the body of procedure Signal checks to see if Cancel_Handler actually cancelled an event timeout. It does this check via the Boolean parameter passed to Cancel_Handler, named Handler_Was_Set. If True, the timeout was pending, which means there was a caller suspended in Wait for this Timed_Condition object. In that case we set the Timed_Out Flag to False and unblock the suspended caller. If Handler_Was_Set is False, there was no pending timeout, hence no caller suspended in Wait, so nothing further is done. An important aspect of the Timing_Event operations is that they are free of race conditions, per language rules, when operating on any given Timing_Event object. In addition, execution of Set_Handler is atomic with respect to the execution of the handler for that same object. Therefore, execution of these operations' internal statements will not be interleaved. However, calls to them might be interleaved. For example, let's assume a task will call Wait and another task will call Signal, for the same Timed_Condition object. Wait could be about to make the call to Set_Handler, and then be preempted by the other task calling Cancel_Handler (via Signal). We know that Set_Handler and Cancel_Handler will be executed atomically, so either Set_Handler or Cancel_Handler will execute first, followed by the other. The if-statement in Signal ensures that either order works. If Set_Handler executes first, followed immediately by Cancel_Handler, the Boolean parameter Handler_Was_Set will return True and hence Caller_Unblocked will be set to True. When Wait resumes execution it will call Suspend_Until True but will find Caller_Unblocked True, so it will return immediately, and then finish Wait's execution. Alternatively, if Cancel_Handler executes first, Handler_Was_Set will be False and nothing further will be done in Signal. The call to Wait will then continue as usual, waiting for Signal to be called. The application must be structured such that another call to Signal does eventually occur, if that should happen prior to the timeout. These are not persistent signals. Finally, recall we said there was something to mention about the formal parameter profile for Signal_Timeout. Specifically, the type for the formal parameter must always be type Timing_Event, otherwise such a protected procedure would not be compatible with the access type. The runtime system will automatically call Signal_Timeout for us if/when the timeout expires, and will pass the specific Timed_Condition object to the handler. But although it is a Timed_Condition object, the view is as a Timing_Event object because that is the type of the formal parameter. Therefore, we have to convert the view inside the procedure from type Timing_Event to type Timed_Condition. Without the conversion, the view as a Timing_Event parameter would not allow references in the handler body to the extension components Timed_Out and Caller_Unblocked. The syntax to do the view conversion is used in the renaming declaration. It's a bit ugly, but is always the same approach: convert "up" to the "base" type, i.e., the root class-wide type, and then "down" to the specific derived type. The compiler may issue code to check that the right target type is actually involved, or it might recognize the fact that, in this case, the view conversion is always correct. The Example Now that we have the facility in place, let's have an example. We'll use one of the STM32 Discovery Kit boards that has a user button and some LEDs on it. A task will call Wait on a Timed_Condition variable, and an interrupt handler for the user button will Signal that same Timed_Condition variable. If the user doesn't press the button prior to the timeout deadline, the waiting task will turn on the orange LED and turn off the green LED. If the user does press the button in time, the waiting task will turn on the green LED and turn off the orange LED. This processing continues until power is pulled. First, here's the declaration for the library package containing the waiter task. Ravenscar and Jorvik require all tasks to be declared at the library level: package LED_Controller is task Control; end LED_Controller; We'll take the default task priority and stack size for Control. Next, the package body, which I promise is more interesting: with Ada.Real_Time; use Ada.Real_Time; with STM32.Board; use STM32.Board; with User_Button; with Timed_Conditions; use Timed_Conditions; package body LED_Controller is ------------- -- Control-- ------------- task body Control is Time_Expired : Boolean; Timeout : constant Time_Span := Milliseconds (2_000); -- arbitrary begin loop Wait (User_Button.Pressed, Timeout, Time_Expired); if Time_Expired then Orange_LED.Set; Green_LED.Clear; else Orange_LED.Clear; Green_LED.Set; end if; end loop; end Control; end LED_Controller; As the comment indicates, the timeout of two seconds is entirely arbitrary. Package STM32.Board defines the devices on the STM32F407 Discovery board. In this code we use the two LEDs and the blue user button. Package User_Button is defined here to declare the Timed_Condition variable Pressed, the button hardware initialization routine, and the button interrupt handler. Here's the package declaration: with Timed_Conditions; use Timed_Conditions; package User_Button is procedure Initialize (Use_Rising_Edge : Boolean := True); Pressed : Timed_Condition; end User_Button; There we see the variable and the hardware initialization procedure. The interrupt handler is declared within the package body: with STM32.Board; use STM32.Board; with STM32.Device; use STM32.Device; with STM32.GPIO; use STM32.GPIO; with STM32.EXTI; use STM32.EXTI; with System; package body User_Button is Button_High : Boolean := True; EXTI_Line : constant External_Line_Number := User_Button_Point.Interrupt_Line_Number; ------------ -- Button -- ------------ protected Button with Interrupt_Priority => System.Interrupt_Priority'Last is procedure Interrupt with Attach_Handler => User_Button_Interrupt; end Button; ------------ -- Button -- ------------ protected body Button is --------------- -- Interrupt -- --------------- procedure Interrupt is begin Clear_External_Interrupt (EXTI_Line); if (Button_High and then User_Button_Point.Set) or else (not Button_High and then not User_Button_Point.Set) then -- we would de-bounce the button, but no need for this demo Timed_Conditions.Signal (Pressed); end if; end Interrupt; end Button; ---------------- -- Initialize -- ---------------- procedure Initialize (Use_Rising_Edge : Boolean := True) is begin Enable_Clock (User_Button_Point); User_Button_Point.Configure_IO ((Mode => Mode_In, Resistors => (if Use_Rising_Edge then Pull_Down else Pull_Up))); -- Connect the button's pin to the External Interrupt Handler User_Button_Point.Configure_Trigger (if Use_Rising_Edge then Interrupt_Rising_Edge else Interrupt_Falling_Edge); Button_High := Use_Rising_Edge; end Initialize; end User_Button; The protected procedure Button.Interrupt is the handler for the interrupt, not surprisingly. When the hardware interrupt occurs, if the physical button has been pressed Signal is called. The details of setting up the interrupt are not particularly pertinent. It is worth mentioning, for the sake of clarity, that User_Button_Point is a GPIO port/pin pair that is defined by package STM32.Board. Finally, the main procedure. The task and interrupt handler do all the work, but the main procedure first initializes the hardware, including the two LEDs. with Ada.Real_Time; use Ada.Real_Time; with Last_Chance_Handler; pragma Unreferenced (Last_Chance_Handler); -- The "last chance handler" is the user-defined routine that is called when -- an exception is propagated. We want it in the executable, therefore it -- must be somewhere in the closure of the context clauses. with STM32.Board; with System; with LED_Controller; pragma Unreferenced (LED_Controller); with User_Button; procedure Test is pragma Priority (System.Priority'Last); begin STM32.Board.Initialize_LEDs; STM32.Board.All_LEDs_Off; User_Button.Initialize; loop delay until Time_Last; end loop; end Test; The main procedure specifies the highest non-interrupt priority for the environment task so all the hardware initialization occurs first. The task in package LED_Controller is activated automatically, and eventually calls Wait. At some point someone will press the physical button on the board, generating the interrupt. The LEDs will be lit accordingly. The use of pragma Unreferenced prevents the compiler from issuing warnings about the fact that the main procedure doesn't do anything with certain packages. Ordinarily we'd want such warnings. But if they are not referenced, why mention them? For the packages to appear in the executable, they must appear somewhere in the transitive closure of the with-clauses. There's no reference to them elsewhere in the example code, so we pull them in here, in the main, and tell the compiler that it's OK. After setting up the hardware, the main procedure goes into an infinite loop. That's required because tasks in Ravenscar and Jorvik should never complete and terminate -- including the implicit environment task that calls the main procedure. My personal preference is to have an extremely long delay inside the loop so that the main doesn't consume CPU cycles. But a null loop would work too, as long as the environment task is given a priority lower than any tasks that will be doing actual application processing. In this example we wanted the main procedure to do something before the Control task, without actually synchronizing with the task, so we gave it the highest priority. With that priority something that actually suspended the environment task was required, rather than a null loop. An elegant alternative to the long delay would be to suspend on another Suspension_Object variable that will never be set to True. (Thanks to Bob Duff for that suggestion.) The code we used for the STM32 board and drivers is part of the Ada Drivers Library (ADL) provided by AdaCore and the Ada community. The ADL is available on GitHub for both non-proprietary and commercial use here: https://github.com/AdaCore/Ada_Drivers_Library. As always, questions and comments are welcome! ]]> A design pattern for OOP in Ada https://blog.adacore.com/a-design-pattern-for-oop-in-ada Mon, 13 Sep 2021 11:04:00 -0400 Fabien Chouteau https://blog.adacore.com/a-design-pattern-for-oop-in-ada When I do Object Oriented Programming with Ada, I tend to follow a design pattern that makes it easier for me and hopefully also for people reading my code. For this post I will use the all time classic example of OOP, a Graphic User Interface framework. Here’s what a widget specification looks like: package Widget.Button is subtype Parent is Widget.Instance; type Instance is new Parent with private; subtype Class is Instance'Class; type Acc is access all Instance; type Any_Acc is access all Class; overriding procedure Event (This : in out Instance; Evt : Event_Kind); overriding procedure Draw (This : in out Instance); private subtype Dispatch is Instance'Class; type Instance is new Parent with record C : Boolean := False; end record; end Widget.Button; Now let’s have a look at each element. Type Instance is [...] package Widget.Button is type Instance is [...] with private; I define one tagged type (object) per package and the name of this type is always “Instance”. As we all know, naming is the hardest thing in programming, so having to find meaningful names for both the package and type is annoying at best. Another solution is to use plural for packages names and singular for the types: package Widgets.Buttons is type Button is [...] with private; But there is one other benefit to using the same type name in every package: easier refactoring. We know that Ada is an amazing language when it comes to safe refactoring, thanks to its strong typing and powerful specifications. One might say that Ada provides “fearless refactoring”. The drawback is that changing the signature of a method, for instance, means a lot of code to edit. With this design pattern, an inherited methods look exactly the same for all types: overriding procedure Event (This : in out Instance; Evt : Event_Kind); So we can just copy/past it everywhere it is needed when the signature changes. Using the other plural/singular naming convention, the type of “This” changes every time. overriding procedure Event (This : in out Button; Evt : Event_Kind); overriding procedure Event (This : in out Checkbox; Evt : Event_Kind); This may look like a detail but it makes sense to me, and I like the consistency of this convention. The declaration of widgets looks like this: with Widget.Button; with Widget.Checkbox; use Widget; [...] B : Button.Instance; C : Checkbox.Instance; Class, Acc and Any_Acc subtype Class is Instance'Class; type Acc is access all Instance; type Any_Acc is access all Class; The next types and subtype declaration follow the same idea, they are always the same for every object. The subtype “Class” is useful when writing class-wide subprograms, e.g.: procedure Something (B : Button.Class); The type “Acc” is a general access type for the instance: B : Button.Acc := new Button.Instance; The type “Any_Acc” is a access type for any object in the hierarchy: procedure Something (B : not null Button.Any_Acc); You can even define more access types like: type Const_Acc is access constant Instance; type Any_Const_Acc is access constant Class; And if you don’t like “Acc” you can use “Reference” or Ref, just stay consistent across your hierarchy: type Reference is access all Instance; type Any_Reference is access all Class; type Ref is access all Instance; type Any_Ref is access all Class; subtype Parent is Widget.Instance; subtype Parent is Widget.Instance; type Instance is new Parent with private; Emmanuel Briot already wrote a blog post here on this pattern. Always defining a subtype that names the parent type of the object makes inherited subprogram call easy, readable and safe: overriding procedure Event (This : in out Instance; Evt : Event_Kind) is begin Parent (This).Event (Evt); end Event; I will let you read Emmanuel’s post to see the pitfalls of other approaches. subtype Dispatch is Instance'Class; private subtype Dispatch is Instance'Class; In Ada, dynamic dispatching on subprograms only occurs on class-wide types. This is unsettling for many, and I was myself caught by this when I started OOP in Ada. By defining a “Dispatch” subtype, we can make dispatching call explicit, easier to spot and use: overriding procedure Draw (This : in out Instance) is begin Dispatch (This).Event (Draw_Event); end Draw; Conclusion I hope this pattern will be useful to some of you. Let me know in the comments what is your opinion on this, and maybe what other patterns you are using in Ada. ]]> When the RISC-V ISA is the Weakest Link https://blog.adacore.com/when-the-isa-is-the-weakest-link Thu, 02 Sep 2021 04:26:00 -0400 Yannick Moy https://blog.adacore.com/when-the-isa-is-the-weakest-link NVIDIA has been using SPARK for some time now to develop safety- and security-critical firmware applications. At the recent DEF CON 29, hackers Zabrocki and Matrosov presented how they went about attacking NVIDIA firmware written in SPARK but ended up attacking the RISC-V ISA instead! Zabrocki starts by explaining the context for their red teaming exercise at NVIDIA, followed by a description of SPARK and their evaluation of the language from a security attack perspective. He shows how they used an extension of Ghidra to decompile the binary code generated by GNAT and describes the vulnerability they identified in the RISC-V ISA thanks to that decompilation. Matrosov goes on to explain how they glitched the NVIDIA chip to exploit this vulnerability. Finally, Zabrocki talks about projects used to harden RISC-V platforms. What I found amazing about this presentation is that because of the protection provided by the NVIDIA team’s developing the software in SPARK and proving it free of runtime errors, the hackers had to turn to something other than the software to find vulnerabilities - which led them to find one in the RISC-V ISA itself! Zabrocki correctly pointed out that memory exhaustion issues are not detected by the SPARK tool, GNATprove. Instead, you should use, for example, GNATstack to detect (some of) them. This is a perfect example of non-functional requirements that are not checked by the SPARK tool. Other similar requirements include timing constraints for real-time software and robustness against cosmic rays for satellite software. Finally, note that SPARK supports safe pointers (an enhancement added in SPARK Pro 2020), and that the classes of problems detected by the tool are clearly defined in the tool documentation. ]]> Security-Hardening Software Libraries with Ada and SPARK https://blog.adacore.com/security-hardening-with-spark Thu, 08 Jul 2021 03:34:00 -0400 Kyriakos Georgiou https://blog.adacore.com/security-hardening-with-spark Part of AdaCore's ongoing efforts under the HICLASS project is to demonstrate how the SPARK technology can play an integral part in the security-hardening of existing software libraries written in other non-security-oriented programming languages such as C. This blog post presents the first white paper under this work-stream, “Security-Hardening Software Libraries with Ada and SPARK”. In this paper, we have taken a quantitative approach, where SPARK is put under test to demonstrate that it can be relatively easily adopted and that the massive benefits of its adoption do not come with a significant negative impact on the performance of a program. To showcase this, a set of benchmarks from a modern C embedded benchmark suite, namely the EMBENCH suite, are converted to SPARK to guarantee the Absence of RunTime Errors (AoRTE). Runtime errors are a well-known source of security-related vulnerabilities, with dozens of system security breaches related to them. Thus, such runtime errors should be eliminated in the context of high assurance systems to avoid potential exploitations that can lead to safety issues. The work not only achieves the above goals but in the process, a plethora of valuable guidelines and best practices are offered to enable others to rapidly adopt the technology for hardening software libraries. SPARK Levels of Software Assurance Adopting SPARK for a new project can seem like an intimidating task, especially when having no prior experience with the technology. The difﬁculty of adopting SPARK depends on the scope of the analysis, such as a whole project or units of a project, and the software assurance level required. Fortunately, there are well-deﬁned guidelines on the adoption of the SPARK technology that simplify the process. These guidelines are offered in the form of ﬁve levels of SPARK assurance, which are incremental in both the effort required and the amount of software assurance they provide. The ﬁve levels are: 1. Stone level - valid SPARK 2. Bronze level - initialization and correct data ﬂow 3. Silver level - Absence of RunTime Errors (AoRTE) 4. Gold level - proof of key integrity properties 5. Platinum level - full functional proof of requirements For this work, we are aiming to reach up to the third level of SPARK assurance. With this level, programs are guaranteed AoRTE, a major source of security vulnerabilities. This also eliminates the need for runtime checks. To move at the highest levels of SPARK requires the precise intent and speciﬁcations of the application to be known, something not always feasible for code that is not written from scratch or well documented, as is the case with the EMBENCH benchmarks conversion. The EMBENCH Benchmark Suite EMBENCH is a recently formed free and open-source embedded benchmark suite, intending to provide benchmarks that reﬂect today’s embedded Internet of Things (IoT) applications’ needs. The initiative for establishing the benchmark suite came from Prof. David A. Patterson, one of the principal ﬁgures behind the RISC-V processor. The main idea is to move away from outdated, artiﬁcial benchmarks, such as the Dhrystone and CoreMark, which are no longer representative of modern embedded applications, and introduce a benchmark suite that will continuously evolve to keep up with the new trends in embedded systems. Currently, EMBENCH includes 19 benchmarks carefully selected to be a representative range of the application space that is typically found in today’s IoT. The suite’s build-system supports both native (x-86, Linux) and cross-compilation of the benchmarks, currently for RISC-V and Arm’s Cortex-M4 based development boards. Selection of Benchmarks For a fair performance comparison between the programming languages C and SPARK, the underlying algorithmic logic and the main program structure of the original EMBENCH benchmarks had to be retained in their new correspondent SPARK versions. This limitation can signiﬁcantly impact the effective use of features and capabilities of the Ada/SPARK languages. For example, casting from one scalar type to another scalar type of smaller range or shifting signed integers are things that you typically do not expect to be part of an Ada or SPARK program. Furthermore, there is a limitation on which level of SPARK assurance can be achieved on each benchmark that stems from the direct translation of the C code to firstly equivalent Ada code and then SPARK code; note that the Ada intermediate stage of translation is needed to facilitate a step-wise process and reduce the complexity of the translation. The original code was not developed with SPARK's formal veriﬁcation technology in mind. This can make it challenging to preserve the original code's logic and structure while enabling the veriﬁcation tools to perform at their optimum level. Finally, in many cases, the lack of sufﬁcient design documentation, particularly in the form of comments within the benchmarks' code, makes it challenging to apply SPARK contracts that capture the semantics of the code. Thus, we can only achieve at best the Silver level of SPARK, AoRTE, which is the default level aimed for high-assurance software. Under the above considerations, eleven benchmarks, shown in the table below, were deemed applicable for this work. The table also shows the level of SPARK assurance achieved for each benchmark. If a benchmark is not at Silver, the number of its sub-modules that reached the Silver level is also given. Benchmark Level Of SPARK No. of Submodules at Silver level aha-mont64 Silver all crc32 Silver all edn Silver all huffbench Silver all matmult-int Silver all nettle-aes Silver all nsichneu Silver all st Bronze 3/4 ud Bronze 0/1 minver Bronze 1/2 nbody Bronze 1/2 Table 1. EMBENCH benchmarks converted to SPARK and their achieved level of SPARK. Achieving the Different SPARK Levels The Implementation Guidance for the Adoption of SPARK manual offers excellent guidance on how to achieve SPARK's several levels of software assurance. These guidelines form the basis for the white paper's work. Its purpose is not to give a comprehensive manual as this is already done in the aforementioned SPARK adoption manual, but rather to give enough context for the reader to understand the process and share the practical experiences of hardening software libraries with Ada and SPARK. Also, any signiﬁcant steps/ﬁndings not found in the SPARK manual used as a starting point, such as a new SPARK feature, are highlighted. The detailed stepwise process followed, and the several challenges and ﬁndings for each step of converting the benchmarks to Ada and SPARK and then achieving the different levels of SPARK software assurance can be found under Section 4.2 of the relevant white paper. One thing I would like to highlight regarding how SPARK can be effectively applied to gain software assurance is the bottom-up modular approach that SPARK technology offers. The modularity is mainly two-fold: 1. Being able to apply the three modes (check, flow analysis, proof) of GNATprove at the line, region, subprogram, and module (ﬁle) level. 2. Each level of SPARK assurance achieved ensures the conformance of all of its lower SPARK assurance levels. It is highly recommended to take a bottom-up modular approach when hardening software libraries with SPARK, such that before moving to a higher level of SPARK, all of the lower levels are achieved. Besides, smaller parts of the code, with the lowest amount of interaction, such as leaf-node subprograms in the call-tree of a program, should be targeted ﬁrst with a gradual move towards the ﬁle level. This signiﬁcantly reduces the effort of achieving the different levels of SPARK assurance. Although for this work SPARK Pro and GNAT Pro versions 21.0w are used, GNAT Community 2021 is also applicable. GNAT Studio 21.0w was the Integrated Development Environment (IDE) used to convert the C benchmarks into SPARK and invoke the GNATprove SPARK static analyser. Evaluation The white paper includes a thorough discussion on the performance evaluation, issues captured by SPARK on the original benchmarks, assessment of the time needed for the completion of the SPARK-related tasks, and some issues found and improved within the SPARK technology. We will cover only the effort/time needed and the performance evaluation for this post. Effort/Time Evaluation One of the aims of this work was to evaluate the effort needed in achieving the absence of runtime errors with the SPARK technology. The completion of this work lasted around 35 working days. The level of experience with the Ada and SPARK technologies was around four months, although the engineer (myself) that carried out the work had an overall programming experience of about 15 years. Taking this into account and that the Silver level of SPARK was achieved within this time frame for the majority of the benchmarks, 7 out of 11, and for most of the functions for the remaining benchmarks, it is fair to say that SPARK technology is easily accessible and its adoption can yield significant benefits for hardening software libraries for security in a short time. Performance Evaluation The performance evaluation was done on an STM32F4DISCOVERY development board. The board is equipped with a 32-bit ARM Cortex-M4 with FPU core, 1-Mbyte Flash memory, 192-Kbyte RAM, and it can run at a maximum frequency of 168MHz. The ARM Cortex-M4 and the RISK-V 32-bit processors are the two embedded processors currently supported in the EMBENCH benchmark suite due to their popularity in the embedded industry. Both in C and SPARK, the benchmarks were compiled with the GNAT Pro 21.0w compiler using the -O2 optimization ﬂag and with link-time optimizations enabled. The link-time optimizations were essential to allow a fairer comparison between the C and the SPARK versions, as the SPARK code for each benchmark is separated from its test harness source ﬁle while the C benchmarks’ code lives in the same ﬁle as their test-harness. Thus, inlining of the benchmark code was feasible in the C versions and not in the SPARK versions. By enabling link-time optimizations, inlining was also enabled for the SPARK code similarly to the C code. Furthermore, Ada’s runtime checks were disabled, using the -gnatp ﬂag, to make the performance comparison fair with C. Nevertheless, subprograms proved at the SPARK Silver level come with an AoRTE guarantee, and thus runtime checks can be safely disabled. This is essential towards certifying at the highest levels of software assurance of safety standards, such as the DO-178C for avionics or the CENELEC EN 50128 for rail systems. This is equally applicable when certifying against security-critical standards and guidelines such as ED-202A/ED-203A and DO-326A/DO-356A for aviation. Benchmark Level Of SPARK SPARK vs C (Performance) aha-mont64 Silver -24.07% crc32 Silver 0.00% edn Silver 8.56% huffbench Silver 2.54% matmult-int Silver 4.27% nettle-aes Silver 10.91% nsichneu Silver 7.55% st Bronze -16.70% ud Bronze 9.49% minver Bronze -0.66% nbody Bronze 38.83% Table 2. Performance comparison of the C and SPARK versions of each benchmark. Note that a positive percentage represents the percentage increase in execution time for a benchmark written in SPARK when compared to the execution time of its corresponding C version, and vice versa, a negative percentage represents the percentage decrease in execution time. Table 2 and Figure 1 show the performance comparison for the C and SPARK versions of the benchmarks. Note that even though the "minver" original benchmark is found to be functionally incorrect, the SPARK level implementation matches that behaviour. Thus, a performance comparison is still valid. With the exception of the "nbody" benchmark, there is no signiﬁcant sacriﬁce in performance when moving from C to SPARK. This is because there were no signiﬁcant intrusive modiﬁcations needed to the code to support SPARK proves. Conclusion The hardening of existing code-bases is expected to become a commonplace security exercise due to the rise of industry-mandated cyber-security standards and guidelines. This is especially predicted within aerospace with the recent adoption by the European Union Aviation Safety Agency (EASA) of the ED-202A/ED-203A security set as the currently only “Acceptable Means of Compliance” for aviation cyber-security airworthiness certiﬁcation. This is expected to be equally true of the Federal Aviation Administration (FAA) and the technically equivalent DO-326A/DO-356A set. This is where SPARK can play an integral role in designing a new security architecture (or when security hardening an existing architecture). The white paper summarized by this post shows that existing systems can beneﬁt from the application of a SPARK hardening approach. Elevation of a security component to SPARK Silver level or higher provides strong evidence to support a security effectiveness assurance argument. This is particularly true when arguing over the effectiveness of security measures against threat scenarios relating to application security vulnerabilities. By design, SPARK aims to eradicate all security bugs, ﬂaws, errors, faults, holes, or weaknesses in software architecture regardless of if threat actors can exploit them. When proven to have achieved Silver level or higher, the guaranteed AoRTE is a powerful countermeasure against cyber-attacks. This work demonstrated that the effort of adopting SPARK is not as hard as perceived; in an arguably short amount of time, a set of benchmarks from the EMBENCH benchmark suite were converted from C to SPARK, and AoRTE was achieved in most of the cases, (see Table 1). In addition, had the complete functional speciﬁcations been available for the remaining, not fully proven benchmarks, the complete set of benchmarks could have achieved the Silver level of SPARK assurance. The SPARK technology was also able to identify a faulty benchmark, namely the "minver", where its implementation was incomplete and produced the wrong results. Furthermore, as demonstrated, the adoption of the SPARK technology did not signiﬁcantly compromise the C-enabled performance. This and the signiﬁcant security beneﬁt of proving AoRTE make SPARK a default choice when it comes to the hardening of software libraries. Finally, the work demonstrates the steps and best practices for adopting SPARK and highlights the use of new features. These, and the provided references to external documentation, can be used as up-to-date guidelines for the easy adoption of the SPARK technology. Acknowledgments The "High-Integrity Complex Large Software and Electronic Systems'' (HICLASS) working group was created to ensure that the UK stays a leading force within civilian aerospace manufacturing. HICLASS is made up of UK academia, tier-one aerospace manufacturers, and associated software development tool providers. This group has been assembled to action the technology strategy set out by the Aerospace Technology Institute (ATI). AdaCore's involvement within this group is to evolve SPARK and associated technologies to ensure they align with the objectives set out within the ever-emerging cyber-security guidelines. HICLASS is supported by the Aerospace Technology Institute (ATI) Programme, a joint Government and industry investment to maintain and grow the UK’s competitive position in civil aerospace design and manufacture, under grant agreement No. 113213. The programme, delivered through a partnership between the ATI, Department for Business, Energy & Industrial Strategy (BEIS), and Innovate UK, addresses technology, capability, and supply chain challenges. More information about the HICLASS project can be found here. ]]> Announcing The First Ada/SPARK Crate Of The Year Award https://blog.adacore.com/announcing-the-first-ada-spark-crate-of-the-year-award Mon, 28 Jun 2021 06:05:00 -0400 Fabien Chouteau https://blog.adacore.com/announcing-the-first-ada-spark-crate-of-the-year-award We're happy to announce our new programming competition, the Ada/SPARK Crate Of The Year Award! We believe the Alire package manager is a game changer for Ada/SPARK, so we want to use this competition to reward the people contributing to the ecosystem. Why “Crate”? This is the name the Alire project uses to designate a software project, library or executable written using the Ada and/or SPARK programming languages and contributed to the Alire ecosystem. The word comes from the Cargo package manager. How does it differ from our previous competition, Make With Ada? • Make with Ada was only for embedded projects: this competition has a prize dedicated to embedded, but the other two prizes are open for any kind of software project. So you can go wild with the topic of your submission. • You can submit a project that you started years ago: it doesn’t have to be developed from scratch during the competition. Timeline The competition is starting today and ends on Friday December 31st 2021 at 23:59 CEST. We'll announce the results in January 2022. As mentioned before, you can submit projects you started before the competition, months or even years ago. The only thing that matters is that your crate has to be available in the Alire community index by the end of the competition. How to enter? The competition is hosted on GitHub. To enter, participants must open an "issue" on the competition repository using the provided template. Read the terms and conditions for more details. Prizes This competition has 3 prizes of2,000 each, for:

• The Ada Crate Of The Year Prize, for best overall Ada crate;
• The SPARK Crate Of The Year Prize, for the best crate written in SPARK and/or contributing to the SPARK ecosystem;
• The Embedded Crate Of The Year Prize, for the best Ada or SPARK crate for embedded software;

Getting started with Alire and Ada/SPARK

You can have a look at the Alire documentation to start your first crate. If you don’t know Ada/SPARK programming, we recommend starting with our interactive online courses here.

We also recommend getting in touch with the Ada/SPARK and Alire community. Here are some links that you may find useful:

Of course, you should also have a look at the existing Alire ecosystem to see if your awesome project idea already exists or to see which existing crates might help in your endeavor.

Have fun, and happy hacking!

]]>
SPARKNaCl with GNAT and SPARK Community 2021: Port, Proof and Performance https://blog.adacore.com/sparknacl-with-gnat-and-spark-community-2021-port-proof-and-performance Fri, 25 Jun 2021 03:54:00 -0400 Roderick Chapman https://blog.adacore.com/sparknacl-with-gnat-and-spark-community-2021-port-proof-and-performance for I in Index_16 loop Do_X; end loop; for I in Index_16 loop Do_Y; end loop;
for I in Index_16 loop
Do_X_And_Y_At_The_Same_Time;
end loop;
declare
T  : GF64_PA;
LT : GF64_Normal_Limb;
begin
T := (others => 0);

for I in Index_16 loop

LT := I64 (Left (I));
T (I) := T (I) + (LT * I64 (Right (0)));
--  and so on for T (I + 1), T (I + 2) ...
declare
subtype U32_Normal_Limb is U32 range 0 .. LMM1;
T  : GF64_PA;
LT : U32_Normal_Limb;
begin
T := (others => 0);

for I in Index_16 loop
LT := U32_Normal_Limb (Left (I));
T (I) := T (I) + I64 (LT * U32_Normal_Limb (Right (0)));
-- and so on...
begin
LT := U32_Normal_Limb (Left (0));

T := GF64_PA'(0  => I64 (LT * U32_Normal_Limb (Right (0))),
1  => I64 (LT * U32_Normal_Limb (Right (1))),
2  => I64 (LT * U32_Normal_Limb (Right (2))),
3  => I64 (LT * U32_Normal_Limb (Right (3))),
4  => I64 (LT * U32_Normal_Limb (Right (4))),
5  => I64 (LT * U32_Normal_Limb (Right (5))),
6  => I64 (LT * U32_Normal_Limb (Right (6))),
7  => I64 (LT * U32_Normal_Limb (Right (7))),
8  => I64 (LT * U32_Normal_Limb (Right (8))),
9  => I64 (LT * U32_Normal_Limb (Right (9))),
10 => I64 (LT * U32_Normal_Limb (Right (10))),
11 => I64 (LT * U32_Normal_Limb (Right (11))),
12 => I64 (LT * U32_Normal_Limb (Right (12))),
13 => I64 (LT * U32_Normal_Limb (Right (13))),
14 => I64 (LT * U32_Normal_Limb (Right (14))),
15 => I64 (LT * U32_Normal_Limb (Right (15))),
others => 0);

--  Iteration "0" is done, so only loop over 1 .. 15 now...
for I in Index_16 range 1 .. 15 loop
--  and so on as before...
]]>
Celebrating Women Engineering Heroes - International Women in Engineering Day 2021 https://blog.adacore.com/international-women-in-engineering-day-2021 Wed, 23 Jun 2021 03:36:00 -0400 Jessie Glockner https://blog.adacore.com/international-women-in-engineering-day-2021

Women make up roughly 38% of the global workforce, yet they constitute only 10–20% of the engineering workforce. In the U.S., numbers suggest that 40% of women who graduate with engineering degrees never enter the profession or eventually leave it. Why? The reasons vary but primarily involve socio-economic constraints on women in general, workplace inequities, and lack of support for work-life balance. Sadly, history itself has often failed to properly acknowledge the instrumental contributions of women inventors, scientists, and mathematicians who have helped solve some of our world's toughest challenges. How can young women emulate their successes if they don't even know about them?

On this International Women in Engineering Day (INWED), we'd like to take the opportunity to celebrate several remarkable women who not only overcame insurmountable challenges to share their exceptional talents but also actively developed some of the most important technologies of humankind. We hope their stories will inspire young women to pursue engineering studies and encourage women engineers to remain in the profession. The world needs you!

Starting off our list is a woman who is particularly near and dear to our hearts here at AdaCore, Lady Ada Lovelace. Enjoy!

Grace Hopper (1906-1992)

Grace Hopper was a computer scientist, programmer, and a rear admiral in the U.S. Navy. She earned a Ph.D. in mathematics from Yale University and was a professor of mathematics at Vassar College. Hopper began her computing career in 1944 as a member of the Harvard Mark I team to develop one of the first computers made for commercial use in the United States. In 1949, she joined the Eckert–Mauchly Computer Corporation, where she was part of the team that developed the UNIVAC I - a commercial data-processing computer designed to replace punched-card accounting machines of the day. She also managed the development of the first COBOL compiler and the COBOL language that is still in use today. Hopper served in the Navy Reserves from 1943 to 1966, but she was recalled to active duty the following year to help standardize the Navy's computer languages. When she retired again in1986, at the age of 79, she was the oldest active-duty commissioned officer in the United States Navy. After she retired from the Navy, she worked as a consultant for Digital Equipment Corporation, sharing her computing experience until her death at age 85. Hopper was elected a fellow of the Institute of Electrical and Electronics Engineers in 1962, named the first computer science Woman of the Year by the Data Processing Management Association in 1969, received the National Medal of Technology in 1991, and the Presidential Medal of Freedom in 2016 (posthumously).

Joan Clarke (1917-1996)

Joan Clarke was a cryptanalyst and numismatist and is known as one of the greatest code-breakers in history. In 1940, while attending Cambridge University, Joan was recruited to join Alan Turing's Hut 8 team at Bletchley Park, best known for breaking the Nazi's Enigma Code and helping end World War II. She was initially placed in an all-women group, referred to as "The Girls," who mainly did routine clerical work. She quickly became the only female practitioner of Banburismus (the cryptanalytic process developed by Alan Turing to decipher German encrypted messages) and deputy head of Hut 8 in 1944. Clarke's work won her many awards and citations, including an appointment as a Member of the Order of the British Empire (MBE) in 1946. After the War, Clarke worked for Government Communications Headquarters (GCHQ). Clarke also developed an interest in numismatics history. She established the sequence of the complex series of gold unicorn and heavy groat coins that were in circulation in Scotland during the reigns of James III and James IV. In 1986, Her research was recognized by the British Numismatic Society when she received the Sanford Saltus Gold Medal. Issue No. 405 of the Numismatic Circular described her paper on the topic as "magisterial." Keira Knightley portrayed Clarke in the film The Imitation Game (2014).

Katherine Johnson (1918-2020)

Katherine Johnson was an American mathematician and the first African-American woman to work as a NASA scientist. Her calculations of orbital mechanics as a NASA employee were critical to the success of the first and subsequent American-crewed spaceflights. Johnson's work included calculating trajectories, launch windows, and emergency return paths for Project Mercury space flights, including those for astronauts Alan Shepard, the first American in space, and John Glenn, the first American in orbit. Her calculations helped synch Project Apollo's Lunar Module with the lunar-orbiting Command and Service Module. She also worked on the Space Shuttle and the Earth Resources Technology Satellite (ERTS, later renamed Landsat), authored or co-authored 26 research reports, and worked on plans for a mission to Mars. She retired in 1986, after 33 years at Langley. In 2015, President Barack Obama awarded Johnson the Presidential Medal of Freedom, America's highest civilian honor. In 2016, she was presented with the Silver Snoopy Award by NASA astronaut Leland D. Melvin and a NASA Group Achievement Award. She was portrayed as a lead character in the 2016 film Hidden Figures. In 2019, Johnson was awarded the Congressional Gold Medal by the United States Congress, and in 2021, she was inducted into the National Women's Hall of Fame.

Frances V. Spence (1922 - 2012)

Frances V. Spence is considered one of the first computer programmers in history. Spence was one of eighty women programmers originally hired by the University of Pennsylvania's Moore School of Engineering to develop the ENIAC project - a classified U.S. Army project designed to construct the first all-electronic digital computer to compute ballistics trajectories during World War 2. In addition to her larger programming duties, Spence was also assigned to a smaller computational development team of six women programmers (called "Computers") to operate an analog computing machine known as a Differential Analyzer, used to calculate ballistics equations. When the War ended, Spence continued working with the ENIAC team and collaborated with other leading mathematicians. In 1997, Spence and the other original ENIAC programmers were inducted into the Women in Technology International Hall of Fame. Their work paved the way for the electronic computers of the future, and their innovation kick-started the rise of electronic computing and computer programming in the Post-World War II era.

Annie J. Easley (1933-2011)

Annie J. Easley was an American computer scientist, mathematician, rocket scientist, and one of the first African-Americans to work as a computer scientist at NASA. Easley began her career doing computations for researchers, analyzing problems, and doing calculations by hand. Her earliest work involved running simulations for the newly planned Plum Brook Reactor Facility. She became an adept computer programmer, using languages like the Formula Translating System (Fortran) and the Symbolic Optimal Assembly Programming (SOAP) to support several of NASA's programs. She developed and implemented code used in researching energy-conversion systems by analyzing alternative power technology, including the battery technology for early hybrid vehicles and the Centaur upper-stage rocket. Later in her career, she became NASA's equal employment opportunity (EEO) counselor. In this role, she helped supervisors address gender, race, and age issues in discrimination complaints at the lowest level and in the most cooperative way possible. Easley retired from NASA in 1989, but she remained an active participant in the Speaker's Bureau and the Business & Professional Women's association. She has inspired many through her enthusiastic participation in outreach programs, breaking down barriers for women and people of color in STEM.

Margaret Hamilton (born 1936)

Margaret Hamilton is an American computer scientist, systems engineer, and business owner. She was director of the Software Engineering Division of the MIT Instrumentation Laboratory, which developed on-board flight software for NASA's Apollo program, which successfully landed the first humans on the Moon. On July 20, 1969, as the lunar module, Eagle, approached the Moon's surface, its computers began flashing warning messages. Fortunately, the software developed by Hamilton and her team was not only informing everyone that there was a hardware-related problem, but the software was compensating for it. And with only enough fuel for 30 more seconds of flight, Neil Armstrong reported, "The Eagle has landed." The achievement was a monumental task, given that computer technology was still in its infancy. The astronauts had access to only 72 kilobytes of computer memory (a 64-gigabyte cell phone today carries almost a million times more storage space), and programmers had to use paper punch cards to feed information into room-sized computers with no screen interface. Hamilton's work guided the remaining Apollo missions that landed on the Moon and benefitted Skylab, the first U.S. space station. In 1972, Hamilton left MIT and started her own company, Higher Order Software. Fourteen years later, she launched another company, Hamilton Technologies, Inc., where she created the Universal Systems Language to make the process of designing systems more dependable. Hamilton has published more than 130 papers, proceedings, and reports about sixty projects and six major programs. She is one of the people credited with coining the term "software engineering." NASA honored Hamilton with the NASA Exceptional Space Act Award in 2003. And in 2016, Hamilton received the Presidential Medal of Freedom from President Barack Obama for her work leading to the development of on-board flight software for NASA's Apollo Moon missions.

Sally K. Ride (1951-2012)

Dr. Sally K. Ride was an American astronaut and physicist and the first American woman to travel into space. On June 18, 1983, she served as a mission specialist aboard the space shuttle Challenger. She also became the first American woman to travel to space a second time when she participated in another Challenger mission on Oct. 5, 1984. Ride served on the accident investigation boards set up in response to the two space shuttle tragedies (Challenger in 1986 and Columbia in 2003). And in 2009, she participated in the Augustine committee that helped define NASA's spaceflight goals. Ride stopped working for NASA in 1987 and joined Stanford University Center for International Security and Arms Control. She later became a professor of physics at the University of California, San Diego. She also served as president of Space.com from 1999 to 2000. Until her death in 2012, Ride was a champion for science education and a role model for women, and especially girls. She wrote books for students and teachers and worked with science programs and festivals around the United States. She also came up with the idea for NASA's EarthKAM project, which lets middle school students take pictures of Earth using a camera on the International Space Station. In 2003, Ride was added to the Astronaut Hall of Fame.

Mae Jemison (born 1956)

Mae Carol Jemison is an American engineer, physician, and former NASA astronaut. Jemison graduated from Stanford University in 1977 (one of the only African American students in her class) with a Bachelor of Science degree in Chemical Engineering and a Bachelor of Arts degree in African and African-American studies. She later attended Cornell and received her Doctorate in Medicine in 1981. Shortly after her graduation, she became an intern at the Los Angeles County Medical Center and then practiced general medicine. Fluent in Russian, Japanese, and Swahili, Jemison joined the Peace Corps in 1983 and served as a medical officer for two years in Africa. After working with the Peace Corps, Jemison opened a private practice as a doctor. However, once Sally Ride became the first American woman in space in 1983, Jemison decided to apply to the astronaut program at NASA. In 1987 she was one of 15 people chosen out of over 2,000 applications for NASA's Astronaut Group 12, the first group selected after the Challenger explosion. When she served as a mission specialist aboard the Space Shuttle Endeavour, she became the first African-American woman to travel into space. Jemison left NASA in 1993 and founded The Jemison Group, a technology research company that encourages science, technology, and social change. She also began teaching environmental studies at Dartmouth College and directed the Jemison Institute for Advancing Technology in Developing Countries. She later formed a non-profit educational foundation and, through the foundation, became the principal of the 100 Year Starship project funded by DARPA. Jemison has written several books for children and appeared on television several times, including in a 1993 episode of Star Trek: The Next Generation. She has received multiple awards and honorary doctorates and serves on the Board of Directors for many organizations. She has been inducted into the National Women's Hall of Fame. National Medical Association Hall of Fame, the Texas Science Hall of Fame, and the International Space Hall of Fame.

Donna Auguste is an African-American engineer, entrepreneur, and philanthropist. Auguste received a Bachelor of Science degree in electrical engineering and computer sciences (EECS) from the University of California at Berkeley, a Master of Science from Regis University, and was the first African-American woman to enter the doctoral program in computer science at Carnegie Mellon. Early in her career, Auguste worked at Xerox and was part of the engineering team at IntelliCorp that introduced some of the world's first commercial artificial intelligence knowledge. She also spent several years at Apple Computer. She was awarded four patents for her innovative engineering work on the Apple Newton Personal Digital Assistant, a forerunner to the Palm Pilot. After receiving her Ph.D., she became the founder and CEO of Auguste Research Group, LLC, involved in research around sensors and actionable data science for IoT. In 1996 she founded Freshwater Software, Inc. to provide companies with tools that would help them monitor and enhance their presence on the Internet. She served as CEO of Freshwater until she sold it in 2000 for $147 million. August also founded the Leave a Little Room Foundation, LLC - a philanthropic organization that helps provide housing, electricity, vaccinations, and improve education and infrastructure to poor communities worldwide. Her current research, DataTip, involves using smartphone sensors to engage non-technical youth and adults in STEM learning to create content relevant to daily living. She was recognized as one of "25 Women Who Are Making It Big in Small Business" by Fortune Magazine. She also won the 2001 Golden Torch Award for Outstanding Women in Technology. Limor Fried (born 1980) Fried is an American electrical engineer and businesswoman. Fried studied at MIT, earning a B.S. in Electrical Engineering and Computer Science (EECS) in 2003 and a Master of Engineering in EECS in 2005. From her dorm room, she created Adafruit Industries (@adafruti). The company designs and resells open-source electronic kits, components, and tools, mainly for the hobbyist market. Adafruit currently ranks #11 of the top 20 USA manufacturing companies and #1 of "fastest-growing private companies in New York City" by Inc. 5000. Fried has been influential in the open-source hardware community. She participated in the first Open Source Hardware Summit and drafted the Open Source Hardware definition. In 2009, she received the Pioneer Award from the Electronic Frontier Foundation for her participation. Fried's numerous accolades continue. She was awarded the Most Influential Women in Technology award in 2011 by Fast Company magazine and became the first female engineer featured on the cover of Wired magazine. In 2012 she was the only female on a list of 15 finalists for entrepreneur's "Entrepreneur of the Year" award. She was named a White House Champion of Change in 2016, became one of Forbes' America's Top 50 Women In Tech in 2018, and received a Women in Open Source Award (Community) by Red Hat in 2019. Known by her moniker ladyada, a homage to Lady Ada Lovelace, she continues to motivate countless girls, young women, and others toward DIY frontiers and in science, technology, engineering, and mathematics. If you'd like to learn about more amazing women, check out INWED's article highlighting five women leaders who have made a considerable impact in the tech world and serve as an inspiration to many. "Women belong in all places where decisions are being made." - Ruth Bader Ginsberg ]]> Going beyond Ada 2022 https://blog.adacore.com/going-beyond-ada-2022 Thu, 03 Jun 2021 10:27:00 -0400 Arnaud Charlet https://blog.adacore.com/going-beyond-ada-2022 As we've seen previously in Ada 2022 support in GNAT, the support for Ada 2022 is now mostly there for everyone to take advantage of. We're now crossing fingers for this new revision to be officially stamped by ISO in 2022. In practice, making new ISO revisions of the language is a long process, which so far happens roughly every ten years. On our side, following the general evolution of the language design culture, and some feedback we've received from our community of users, we're looking to have a shorter and more lively feedback loop. This is why we have started in 2019 a new initiative, centered around the Ada/SPARK RFCs platform, which will allow us to experiment with new language features for Ada. With this platform, we want to give anyone an opportunity to propose language evolutions through RFCs ("request for comments"), discuss the merits of RFCs publicly with the community, select those that are the most promising for prototyping, prototype the features in GNAT (and/or SPARK as appropriate), gather feedback from users of the feature, and depending on that feedback, either abandon the feature, modify it, or keep it. Finally when relevant, propose the features we kept for inclusion in the next version of Ada to the Ada Rapporteur Group, the international body in charge of Ada standardization. In order to assess the needs of current and future users of the Ada programming language, we asked people inside and outside AdaCore what they wish for the future of Ada and SPARK. We got a lot of insights from these answers, coming from programmers with different backgrounds. One of the most common request is to get more compile-time guarantees, in areas such as data initialization before use, access to discriminated fields, dereference of possibly null pointers, dynamic memory management. Note that SPARK already offers such guarantees, at the cost of constraining the language and requiring an analysis which is much more costly than compilation. Here, our goal will be to provide the guarantees above for Ada programs through simpler compilation. Other common requests were: more powerful generics with richer specifications and implicit instantiation, better string handling that properly supports unicode, a more universally available mechanism for data finalization as well as some frequently requested syntax additions. The peculiar object model in Ada turned out to be a contentious issue, with some strong supporters and strong opponents, who advocated for rebuilding it more alike to the dominant model of Java/C++. Thanks to the suggestions received so far from both external contributors and from the language design team at AdaCore, we have started adding some of these new experimental features, with a first implementation available in the latest GNAT Community 2021 release as well as the latest GNAT Pro 22 Continuous Release, under the -gnatX switch and detailed below. Most Wanted Features Let's first start with two "most wanted" features that many Ada users have been asking for years: Fixed Lower Bound Detailed in RFC#38, you can now specify a lower bound for unconstrained arrays that is fixed to a certain value. Use of this feature increases safety by simplifying code, and can also improve the efficiency of indexing operations. For example, a matrix type with fixed lower bounds of zero for each dimension can be declared by the following: type Matrix is array (Natural range 0 .. <>, Natural range 0 .. <>) of Integer; Objects of type Matrix declared with an index constraint must have index ranges starting at zero: M1 : Matrix (0 .. 9, 0 .. 19); M2 : Matrix (2 .. 11, 3 .. 22); -- Warning about bounds; will raise CE Similarly, a subtype of String can be declared that specifies the lower bound of objects of that subtype to be 1: subtype String_1 is String (1 .. <>); If a string slice is passed to a formal of subtype String_1 in a call to a subprogram S, the slice’s bounds will “slide” so that the lower bound is 1. Within S, the lower bound of the formal is known to be 1, so, unlike a normal unconstrained String formal, there is no need to worry about accounting for other possible lower-bound values: procedure Str1 is subtype String_1 is String (1 .. <>); procedure Proc (S : String_1) is begin -- S'First = 1 Put_Line (S); end Proc; S : String_1 := "hello world"; begin Proc (S (7 .. S'Last)); -- sliding on S (7 .. S'Last) occurs automatically when calling Proc, -- so this will pass a String_1 (1 .. 5) whose content is "world" end Str1; Generalized Object.Op Notation Detailed in RFC#34, the so called prefixed-view notation for calls is extended so as to also allow such syntax for calls to primitive subprograms of untagged types. The primitives of an untagged type T that have a prefixed view are those where the first formal parameter of the subprogram either is of type T or is an anonymous access parameter whose designated type is T. This is another "most wanted" feature since the introduction of this notation in Ada 2005! For example: generic type Elem_Type is private; package Vectors is type Vector is private; procedure Add_Element (V : in out Vector; Elem : Elem_Type); function Nth_Element (V : Vector; N : Positive) return Elem_Type; function Length (V : Vector) return Natural; ... end Vectors; package Int_Vecs is new Vectors(Integer); V : Int_Vecs.Vector; ... V.Add_Element(42); V.Add_Element(-33); pragma Assert (V.Length = 2); pragma Assert (V.Nth_Element(1) = 42); Additional "when" Constructs This smaller syntactic addition discussed in RFC#73 adds the ability to use the "when" keyword to "return", "goto" and "raise" statements, in addition to the existing "exit when" control structure. For example: procedure Do_All (Element : access Rec; Success : out Boolean) is begin raise Constraint_Error with "Element is null" when Element = null; Do_1 (Success); return when not Success; Do_2 (Success); return when not Success; Do_3 (Success); return when not Success; Do_4 (Success); end Do_All; Pattern Matching This feature on the other hand (detailed in RFC#50) is a large one and is still being worked on. It provides an extension of case statements to cover records and arrays, as well as finer grained casing on scalar types and will in particular in the future provide more compile time guarantees when accessing discriminated fields. For example, you can match on several scalar values: type Sign is (Neg, Zero, Pos); function Multiply (S1, S2 : Sign) return Sign is (case (S1, S2) is when (Neg, Neg) | (Pos, Pos) => Pos, when (Zero, <>) | (<>, Zero) => Zero, when (Neg, Pos) | (Pos, Neg) => Neg); Matching composite types is currently only supported on records with no discriminants. Support for discriminants and arrays will come later. The selector for a case statement may be of a composite type. Aggregate syntax is used for choices of such a case statement; however, in cases where a “normal” aggregate would require a discrete value, a discrete subtype may be used instead; box notation can also be used to match all values. Consider this example: type Rec is record F1, F2 : Integer; end record; procedure Match_Record (X : Rec) is begin case X is when (F1 => Positive, F2 => Positive) => Do_This; when (F1 => Natural, F2 => <>) | (F1 => <>, F2 => Natural) => Do_That; when others => Do_The_Other_Thing; end case; end Match_Record; If Match_Record is called and both components of X are Positive, then Do_This will be called; otherwise, if either component is nonnegative (Natural) then Do_That will be called; otherwise, Do_The_Other_Thing will be called. If the set of values that match the choice(s) of an earlier alternative overlaps the corresponding set of a later alternative, then the first set shall be a proper subset of the second (and the later alternative will not be executed if the earlier alternative “matches”). All possible values of the composite type shall be covered. In addition, pattern bindings are supported. This is a mechanism for binding a name to a component of a matching value for use within an alternative of a case statement. For a component association that occurs within a case choice, the expression may be followed by “is <identifier>”. In the special case of a “box” component association, the identifier may instead be provided within the box. Either of these indicates that the given identifier denotes (a constant view of) the matching subcomponent of the case selector. Consider this example (which uses type Rec from the previous example): procedure Match_Record2 (X : Rec) is begin case X is when (F1 => Positive is Abc, F2 => Positive) => Do_This (Abc); when (F1 => Natural is N1, F2 => <N2>) | (F1 => <N2>, F2 => Natural is N1) => Do_That (Param_1 => N1, Param_2 => N2); when others => Do_The_Other_Thing; end case; end Match_Record2; This example is the same as the previous one with respect to determining whether Do_This, Do_That, or Do_The_Other_Thing will be called. But for this version, Do_This takes a parameter and Do_That takes two parameters. If Do_This is called, the actual parameter in the call will be X.F1. If Do_That is called, the situation is more complex because there are two choices for that alternative. If Do_That is called because the first choice matched (i.e., because X.F1 is nonnegative and either X.F1 or X.F2 is zero or negative), then the actual parameters of the call will be (in order) X.F1 and X.F2. If Do_That is called because the second choice matched (and the first one did not), then the actual parameters will be reversed. Simpler Accessibility Rules This one (see RFC#47) is still a moving target and would definitely welcome some user experimentation and feedback! It starts with the observation that over the years, the rules that govern accessibility in Ada, that is, what operations on pointers are allowed, have grown to a point where they are barely understood by implementers and even less so by users. So by introducing a new restrictions pragma, we want to both simplify the rules and propose a model where runtime accessibility checks related to the use of anonymous access types are suppressed and replaced by compile time checks, in particular because the runtime accessibility checks are either impossible to implement fully, or worse, may produce false alarms (raising an exception in cases where no dangling access is actually occurring). So if you add as part of your configuration pragmas the following: pragma Restrictions (No_Dynamic_Accessibility_Checks); This will enable this new mode. Currently two variants are implemented: Designated type model The default model when using No_Dynamic_Accessibility_Checks in GNAT Community Edition 2021 although in the more recent GNAT Pro development version, we've switched this model with the other one and this one will be enabled via the -gnatd_b debug switch, in addition to the restriction. In this model, anonymous access types are implicitly declared at the point where the designated type is declared. Point of declaration model Available via the additional use of the -gnatd_b switch in GNAT CE 2021 and by default when using No_Dynamic_Accessibility_Checks in more recent GNAT Pro versions, the anonymous access types are implicitly declared at the point where the anonymous access is used (as part of a subprogram parameter, object declaration, etc...). Both models may be refined further based on the feedback received on actual code and what users would find most useful and practical, so do not hesitate to give it a try and let us know! Next Steps Do you find some of these features useful? Do you want to give them a try and tell us what you think? Do you have some ideas for other new Ada features or other changes to existing features? We encourage you to give it a try, give your feedback, and make new suggestions in the Ada/SPARK RFC platform! On our side we'll continue prototyping other RFCs and refine existing ones, so stay tuned. ]]> GNAT Community 2021 is here! https://blog.adacore.com/gnat-community-2021-is-here Tue, 01 Jun 2021 04:20:00 -0400 Fabien Chouteau https://blog.adacore.com/gnat-community-2021-is-here We are happy to announce that the GNAT Community 2021 release is now available via https://www.adacore.com/download. Here are some release highlights: GNAT Compiler toolchain The 2021 GNAT Community compiler includes tightening and enforcing of Ada rules, performance enhancements, and support for many Ada 2022 features: • Jorvik real-time tasking profile • Support for infinite precision numbers • Declare expressions • Contracts on Access-to-Subprogram • Static expression functions • Iterator Filters • Renames with type inference • Container aggregates There are also some future language features - watch this space for further news on this. The compiler back-end has been upgraded to GCC 10 on all platforms. GNAT Studio This release includes GNAT Studio, our multi-language IDE for Ada, SPARK, C, C++ and Python. Notable features are: • Integration of a new engine - clangd - for C/C++ navigation • Various improvements one Ada/SPARK navigation (handling of dispatching calls, dependency browsers) • Improved Search view (highlighted search area and new preferences) • Various UI improvements (better code folding, a new "toggle comments" action, and more) • Many bug fixes and performance improvements Libadalang Libadalang, a library for parsing and semantic analysis of Ada code, has made a lot of progress in the past year. In this GNAT Community release, you'll find: • Improved name resolution • Improved memory footprint • Many new features accessible through the public APIs (details here) SPARK The possibility to prove that your Ada programs are correct with SPARK now applies to more programs with pointers and to programs using the latest features of Ada. GNATprove messages have been enhanced to be more helpful both on the command-line and inside IDEs. You can also now visualize the generated data flow contracts inside GNAT Studio and verify the termination of recursive functions. SPARK support for pointers was enhanced to: • allow dynamic memory (de)allocation in regular functions • use allocators more liberally inside expressions • support named access-to-constant types • support taking the address of a variable with 'Access • support access-to-subprogram types • provide read-only and read-write access to elements inside formal containers without copy SPARK supports the following features of Ada 2022: declare expressions, delta aggregates, contracts on access-to-subprogram types, the @ symbol, iterated component associations. ]]> An Introduction to Jorvik, the New Tasking Profile in Ada 2022 https://blog.adacore.com/introduction-to-jorvik Wed, 26 May 2021 03:11:00 -0400 Pat Rogers https://blog.adacore.com/introduction-to-jorvik In 2016, AdaCore developed and deployed a new tasking profile based directly on the standard Ravenscar profile, but with some restrictions relaxed or replaced. We presented the new profile in 2017 at Ada Europe [1], providing the justification, additional capabilities, execution-time costs, and resulting schedulability analysis supported. That same year, AdaCore extended SPARK to include the new profile, thus supporting both tasking subsets. The new profile is included in the Ada 2022 draft with some refinements and an official name: Jorvik (pronounced “Yourvick”). Jorvik was the Viking name for a Roman fort/settlement that eventually became known as York, in northern England. Said to be England's "Second City," some well-known cities around the world are named after it. The cover picture shows the city today, with the 13th-century Gothic cathedral and its Rose window in the background. York is not far from the village of Ravenscar, where that profile was introduced. Similarly, Jorvik is not far, technically speaking, from Ravenscar. The differences can be summarized by the small handful of restrictions that are removed or replaced from the Ravenscar list. Everything else in Jorvik remains as it is in Ravenscar. Although the changes are few in number, they are nonetheless significant. We will explore the details below. Before we do, you should understand that Jorvik is not a replacement for Ravenscar. The Ada community can benefit from both profiles. To understand why, let's start with a little background. Ravenscar is intended for applications using tasking in four distinct application domains: • safety-critical systems requiring stringent, exhaustive certification analyses, • high-integrity applications requiring formal static analysis and verification, • hard real-time applications requiring predictability and schedulability analysis, • embedded applications requiring a small memory footprint, high performance, or both. Note the changing requirements. They begin with very expensive, comprehensive and rigorous analyses, shift to less costly forms of analysis and non-functional properties (predictability), and end with only non-functional properties (space and speed). Applications can be in more than one of these domains. A hard-real-time application might also be a high-integrity application, for example, and any of them might also be embedded. The requirements for such applications is then the union of all the applicable domains' requirements. In response to that set of requirements, the Ravenscar profile restrictions remove complexity, both in the application source and in the run-time library (RTL), where Ada tasking is largely implemented. The results address the domains' requirements in three ways: 1. A simplified RTL, and resultingly simpler application code, are less costly to analyze for certification and safety. A subset of any modern language is essential to make these analyses feasible, both technically and economically. 2. At the application level, the tasking subset facilitates the various forms of analysis. The structure of the tasks, especially their possible control and data interactions, enables at the task level the sort of safety analysis previously applied to entire (sequential) programs. That same set of restricted interactions also simplifies schedulability analysis for the application code. 3. A simplified RTL can be far more efficient in both object-code space and speed. For example, abort statements impose distributed costs, i.e., object code size and execution performance penalties whether or not they are actually used in an application. Some other constructs do not impose distributed costs but do require complicated run-time support. Removing support for these constructs reduces object code size and improves speed, dramatically. Ravenscar is designed to maximize simplicity to the degree necessary to meet all of the domains' requirements. Therefore, when certification or formal (e.g., safety) analysis is required, Ravenscar is clearly the right choice. When the smallest possible object code size, and absolutely utmost performance is required, Ravenscar is again the best choice. The cost of a simplified RTL, however, is reduced expressive power at the language level. The most expressive language constructs require RTL support and cannot be made available without it. The requeue statement is a good example, as are task entries and accept statements. A rendezvous is effectively an atomic action with two participants, a very potent facility that requires integrated run-time support for both tasking and exceptions. Both profiles trade away expressive power, but to different extents. The controlling factor is the analyses: the high-integrity and safety-critical domains necessitate stringent analyses, whereas the real-time and embedded domains do not. It follows that they do not, in isolation, require the same degree of simplicity. Consequently, Jorvik is designed to enhance expressive power for applications that are only in the real-time and/or embedded domains. As we said earlier, Ravenscar can be used in this case, and sometimes should be used, but the additional expressive power of Jorvik makes it an attractive alternative. With that understood, let’s explore the Jorvik facilities. The differences between the two profiles can be summarized by the restrictions that are removed or replaced from the Ravenscar restrictions list. The pertinent Ravenscar list is as follows:: • Max_Entry_Queue_Length => 1 • Max_Protected_Entries => 1 • Simple_Barriers • No_Relative_Delay • No_Dependence => Ada.Calendar • No_Dependence => Ada.Synchronous_Barriers • No_Implicit_Heap_Allocations The first two restrictions are the most significant. Removing the first allows multiple callers to be queued simultaneously on a protected entry in Jorvik, rather than at most one. Removing the second restriction allows multiple protected entries per protected object (PO). Note that task schedulability analysis remains possible, and the new restriction Max_Entry_Queue_Length can be used to good effect when performing that analysis. Given those two protected entry relaxations, classic protected type idioms are much more likely allowed in Jorvik. Our example is a concurrent bounded buffer, with two entries and as many queued callers at runtime as necessary: generic type Element is private; package Concurrent_Bounded_Buffers is type Content is array (Positive range <>) of Element; protected type Bounded_Buffer (Capacity : Positive) is entry Put (Item : in Element); entry Get (Item : out Element); function Is_Empty return Boolean; function Is_Full return Boolean; private Values : Content (1 .. Capacity); Next_In : Positive := 1; Next_Out : Positive := 1; Count : Natural := 0; end Bounded_Buffer; end Concurrent_Bounded_Buffers; The other major difference between the two profiles is the content of entry barrier expressions. Ravenscar applies the Simple_Barriers restriction that requires these expressions to consist of either a static expression or a name that statically denotes a Boolean component of the enclosing protected object. Loosely speaking that means either Boolean literals (e.g., “when True”), or single Boolean components (e.g., “when Some_Boolean”). Jorvik replaces Simple_Barriers with a new restriction named Pure_Barriers. The new restriction allows more complex Boolean expressions, within limits. Loosely speaking, Pure_Barriers allows the following content for scalar expressions comprising protected entry barriers. You should assume that there are restrictions in the details that I am glossing over. (See RM clauses D.7 and 4.9 for the full definitions.) • a static expression (numeric literals, named numbers, static constants, static calls to static functions, certain attributes, etc.); • a name for a scalar (i.e., a discrete or real type) component of the enclosing protected unit; • a Count attribute reference for an entry in the enclosing protected unit; • a call to a predefined relational operator or Boolean logical operator; • a membership test; • a short-circuit control form; • a conditional_expression; or • an allowed expression that is enclosed in parentheses. No other language entities are allowed in the barrier expressions. Content is restricted so that side effects, exceptions, and recursion are impossible. Precluding them is important because the language does not specify the number of times a given barrier is evaluated. With these restrictions in place the number of evaluations won’t matter. Given this relaxed content, the typical implementations for our Bounded_Buffer entry bodies are allowed without changes, including especially the entry barriers: protected body Bounded_Buffer is entry Put (Item : in Element) when Count /= Capacity is begin Values (Next_In) := Item; Next_In := (Next_In mod Capacity) + 1; Count := Count + 1; end Put; entry Get (Item : out Element) when Count > 0 is begin Item := Values (Next_Out); Next_Out := (Next_Out mod Capacity) + 1; Count := Count - 1; end Get; function Is_Empty return Boolean is (Count = 0); function Is_Full return Boolean is (Count = Capacity); end Bounded_Buffer; For the sake of comparison, imagine we have some other protected object and are using the Ravenscar profile. There will be only one entry, but let’s reuse entry Get above just for illustration. In Ravenscar, we would have a new Boolean component used solely for the entry barrier. It could be named Not_Empty, would be initialized to False, and then updated in the entry body: entry Get (Item : out Element) when Not_Empty is begin Item := Values (Next_Out); Next_Out := (Next_Out mod Capacity) + 1; Count := Count - 1; Not_Empty := Count > 0; end Get; We must include the negation in the name and value because Simple_Barriers requires static expressions. We could not say “not Empty” in the barrier. But note that the assignment to Not_Empty in the entry body has no barrier-oriented restrictions, so the expression comparing Count to zero is allowed, as would much more complex, potentially non-static references. This approach certainly works, but it isn’t the way one would write an entry body and barrier normally, and we’d like to use implementations without requiring code changes when possible. That won’t always be possible, though, because Pure_Barriers does restrict barrier content. The Ravenscar approach might be used occasionally in Jorvik applications too, in combination with the content Jorvik allows. Note that, in the list of content allowed by Jorvik’s Pure_Barriers, “a name that statically names a scalar subcomponent of the immediately enclosing protected unit” has a specific meaning you need to understand. Remember that the Pure_Barriers restriction doesn’t allow anything that can raise exceptions. Therefore, any part of an expression that has to be checked at run-time is not allowed. For example, maybe you have a discriminated record type with dependent components, a default for the discriminant, and a PO component of this type that takes the default. The dependent record components don’t exist except for specific values of the discriminant, which, thanks to the default, can vary as the program executes. Ada checks to make sure that references to those components are consistent with the current value of the discriminant, raising an exception if the check fails. Such a reference would be rejected by the compiler in an expression required to be consistent with the Pure_Barrier restriction. It is not “pure-barrier-eligible” to use the technical term. Likewise, if you have an object of some array type, the correctness of a variable used as the index must be checked (in certain cases). That usage would not be allowed. In other cases, though, the index need not be checked, because the index value is static -- determinable at compile-time -- and so can be checked then. To make this discussion concrete, let’s change the private part and body of the Bounded_Buffer type so that we use a record object. Currently, the private part is as shown earlier: private Values : Content (1 .. Capacity); Next_In : Positive := 1; Next_Out : Positive := 1; Count : Natural := 0; end Bounded_Buffer; Let’s say that Next_In, Next_Out, and Count are to be components of a record object instead of direct PO components. In this case that wouldn’t really be worth doing, but we’ll use it to illustrate what’s allowed. In real code, though, especially an application specific PO, you might very well compose the PO from various abstract data types declared in their own packages (that being good software engineering, after all). type Management is record Next_In : Positive := 1; Next_Out : Positive := 1; Count : Natural := 0; end record; protected type Bounded_Buffer (Capacity: Positive) is … as before private Values : Content (1 .. Capacity); State : Management; end Bounded_Buffer; The entry barriers then become: entry Put (Item : in Element) when State.Count /= Capacity is … entry Get (Item : out Element) when State.Count > 0 is … The references are slightly more verbose, but the point is that the barriers can reference those record components because they are scalar components, because State is declared immediately within the PO, and because the called functions (the relational operators) are static functions statically called. Your real, application-specific PO may very well contain objects of composite types, and you can reference their components in the barriers as long as they follow the rules. Similarly, and perhaps abandoning realistic code altogether, we could use an array of three Integers in place of the three distinct variables. We could say that Next_In will now be State (1), Next_Out will now be State (2), and Count will be State (3). type Management is array (1 .. 3) of Positive; protected type Bounded_Buffer (Capacity : Positive) is … private Values : Content (1 .. Capacity); State : Management; end Bounded_Buffer; The entry barriers would be like so: entry Put (Item : in Element) when State (3) /= Capacity is … entry Get (Item : out Element) when State (3) > 0 is … That approach would not be an improvement, all other things being equal. But it serves to show that array indexing is allowed when the indexes are static. What would be an improvement, however, is using the two existing functions, Is_Full and Is_Empty, in the barriers. They directly express what the reader must otherwise deduce: entry Put (Item : in Element) when not Is_Full is … entry Get (Item : out Element) when not Is_Empty is … Sadly, those barrier expressions are not allowed by Pure_Barriers because Is_Full and Is_Empty are not static functions. They cannot be made to be static, either. The other restrictions removed in Jorvik are mostly a matter of application developer convenience. We fully expect Jorvik applications to delay periodic tasks with absolute delay statements, just as in Ravenscar applications. Elsewhere, a relative delay statement can be appropriate, and they are now allowed. For example, an electro-mechanical relay may have a requirement that it not be actuated more than N times per second in order to prevent burn-out. The semantics of a relative delay match that requirement nicely. In a way, relative delay statements were already allowed in Ravenscar, via the ugly “hack” of using an absolute delay statement to delay until the value of Clock + some-time-span. Jorvik isn’t really adding much here, but the direct expression is cleaner and simpler. Jorvik removes the restriction prohibiting use of the Ada.Calendar package. This restriction is present in Ravenscar because the Ada.Real_Time package has more appropriate semantics for real-time/embedded applications. However, not all usage of Ada.Calendar is unreasonable, for example time-stamping log messages. That said, Ada.Real_Time will surely remain the primary facility. Jorvik removes the restriction prohibiting use of the Ada.Synchronous_Barriers package. A “barrier” in this case is an abstract data type, not a Boolean expression controlling a protected entry. The semantics are much like visiting a restaurant that requires all members of the dinner party to be present before any are seated. An object of type Synchronous_Barrier has a discriminant that specifies how many tasks are in the “party.” Once that many tasks “arrive” by calling the entry for that object, the entire set of caller tasks are allowed to continue. The obvious implementation of type Synchronous_Barrier is as a protected type that, by definition, must allow multiple callers to queue on a single entry. Jorvik allows that implementation so there was no need to restrict the package. Finally, the restriction No_Implicit_Heap_Allocations is removed. That restriction is most pertinent to the domains requiring certification and/or safety analyses, but Jorvik is not targeted to those domains. Some implicit allocations would not be a problem for Jorvik applications. Nevertheless, related restrictions are needed in this regard. Recall that in both Ravenscar and Jorvik no protected objects or tasks are ever allocated. They are always declared. There are cases, however, in which GNAT would allocate a task or protected object dynamically, transparently, even though an allocation is not visible in the source code. The restriction No_Implicit_Heap_Allocations would catch that, but we’ve removed it in Jorvik. For example, consider a composite object declared in a library package, say a String object. If the bounds of the object are not known at compile-time GNAT will allocate the object, implicitly. with Max_Size; -- a function package P is Name : String (1 .. Max_Size); end P; Now, instead of being a library object, imagine Name is a component of a protected type or protected object: with Max_Size; -- a function package P is protected PO is procedure Q; private Name : String (1 .. Max_Size); end PO; end P; It’s the same problem, except now the compiler will be allocating the enclosing protected object, thus violating the restriction. The same behavior is possible for task objects. Therefore, GNAT adds two new restrictions to Jorvik to prevent these specific cases: No_Implicit_Task_Allocations and No_Implicit_Protected_Object_Allocations. With those two restrictions in place, the above code causes this error message from GNAT: p.ads:7:07: violation of restriction "No_Implicit_Protected_Object_Allocations" These two restrictions are not part of the Jorvik profile in the Ada 2022 standard. They are specific to the GNAT implementation. However, we intend to argue for them in a subsequent update to the standard. In conclusion, if you are unsure when to use one of the profiles, or any subset, there is an applicable maxim, originally expressed for Ravenscar: “If an application cannot be reasonably expressed within the Ravenscar subset, it isn’t a Ravenscar application.” In other words, the application code in these domains, particularly those undergoing rigorous analyses, must be very simple, and, consequently, will be expressible in the subset. Otherwise, the project lead should review whether adhering to Ravenscar is appropriate. That maxim is true for the Jorvik profile as well. If an application “genuinely requires” requeue statements, for example, maybe a larger subset is appropriate. Of course, “genuinely requires” is difficult to define precisely, especially because one can work around some of the two profiles’ restrictions via additional application source code. For example, multiple entry queues in a single protected object can be simulated via multiple protected objects, each with a single entry. This additional application code, in effect, implements in a bespoke manner that which the run-time library would have implemented more generally, had the corresponding restriction not been in place. However, that additional source code injects complexity back into the system under analysis. In a very real sense, the complexity has been “moved” from the run-time library up to the application level. At some point, additional application code complexity argues against use of the profiles. That said, Ravenscar is widely used, and justly so. We think Jorvik will be as well. [1] P. Rogers, J. Ruiz, T. Gingold, and P. Bernardi, A New Ravenscar-Based Profile in Reliable Software Technologies Ada-Europe 2017, Johann Blieberger and Marcus Bader (eds) (2017), LNCS 10300, Springer-Verlag, pp. 169-183. ]]> From Rust to SPARK: Formally Proven Bip-Buffers https://blog.adacore.com/from-rust-to-spark-formally-proven-bip-buffers Wed, 05 May 2021 00:00:00 -0400 Fabien Chouteau https://blog.adacore.com/from-rust-to-spark-formally-proven-bip-buffers function In_Writable_Area (This : Buffer; Offset : Buffer_Offset) return Boolean is (if Is_Inverted (This) then -- Already inverted -- |---W==========R----| -- Inverted (R > W): -- We can write between W .. R - 1 Offset in This.Write .. This.Read - 1 else ( -- |====R---------W=====| -- Not Inverted (R <= W): -- We can write between W .. Size - 1, or 0 .. R - 1 if we invert (Offset in This.Write .. This.Size - 1) or else (Offset in 0 .. This.Read - 1))); function Valid_Write_Slice (This : Buffer; Slice : Slice_Rec) return Boolean is (Valid_Slice (This, Slice) and then In_Writable_Area (This, Slice.From) and then In_Writable_Area (This, Slice.To)); procedure Grant (This : in out Buffer; G : in out Write_Grant; Size : Count) with Post => (if Size = 0 then State (G) = Empty) and then (if State (G) = Valid then Write_Grant_In_Progress (This) and then Slice (G).Length = Size and then Valid_Slice (This, Slice (G)) and then Valid_Write_Slice (This, Slice (G))); -- Request indexes of a contiguous writeable slice of exactly Size elements declare Left : Sample_Array (1 .. 64) := (others => 0); Right : Sample_Array (Left’Range) := (others => 0); Q : aliased Offsets_Only (Left'Length); WG : Write_Grant := BBqueue.Empty; S : Slice_Rec; begin Grant (Q, WG, 8); if State (WG) = Valid then S := Slice (WG); Left (Left'First + S.From .. Left'First + S.To) := (others => 42); Right (Right'First + S.From .. Right'First + S.To) := (others => -42); Commit (Q, WG); end if; declare type My_Data is record A, B, C : Integer; end record; type My_Data_Array is array (Natural range <>) of My_Data; Buf : My_Data_Array (1 .. 64); Q : aliased Offsets_Only (Buf'Length); WG : Write_Grant := BBqueue.Empty; S : Slice_Rec; begin Grant (Q, WG, 8); if State (WG) = Valid then S := Slice (WG); Buf (Buf'First + S.From .. Buf'First + S.To) := (others => (1, 2, 3)); Commit (Q, WG); end if; declare Q : aliased Buffer (64); WG : Write_Grant := Empty; S : Slice_Rec; begin Grant (Q, WG, 8); if State (WG) = Valid then declare B : Storage_Array (1 .. Slice (WG).Length) with Address => Slice (WG).Addr; begin B := (others => 42); end; Commit (Q, WG); end if; declare Q : aliased Framed_Buffer (64); WG : Write_Grant := Empty; RG : Read_Grant := Empty; S : Slice_Rec; begin Grant (Q, WG, 8); -- Get a grant of 8 Commit (Q, WG, 4); -- Only commit 4 Grant (Q, WG, 8); -- Get a grant of 8 Commit (Q, WG, 5); -- Only commit 5 Read (W, RG); -- Returns a grant of size 4 generic type T is mod <>; package Atomic.Generic8 with Preelaborate, Spark_Mode is type Instance is limited private; -- This type is limited and private, it can only be manipulated using the -- primitives below. procedure Add_Fetch (This : aliased in out Instance; Val : T; Result : out T; Order : Mem_Order := Seq_Cst) with Post => Result = (Value (This)'Old + Val) and then Value (This) = Result; ]]> On the Benefits of Families ... (Entry Families) https://blog.adacore.com/on-the-benefits-of-families Wed, 28 Apr 2021 05:04:00 -0400 Pat Rogers https://blog.adacore.com/on-the-benefits-of-families Ada has a concurrency construct known as “entry families” that, in some cases, is just what we need to express a concise, clear solution. For example, let’s say we want to have a notion of “conditions” that application tasks can await, suspending until the specified condition is "signaled." At some point, other tasks will signal that these conditions are ready to be handled by the waiting tasks. Understand that conditions don't have any state of their own, they are more like "events" that have either happened or have not, and may happen more than once. For the sake of discussion let’s generalize this idea to an enumeration type representing four possible conditions: type Condition is (A, B, C, D); These condition names are not very meaningful but they are just placeholders for those that applications would actually define. Perhaps a submersible's code would have conditions named Hatch_Open, Hatch_Closed, Umbilical_Detached, and so on. Responding tasks can suspend, waiting for an arbitrary condition to be signaled, and other tasks can signal the occurrence of conditions, using a “condition manager” that the two sets of tasks call. Here’s the declaration of the condition manager type: type Manager is limited private; The type is limited because it doesn’t make sense to assign one manager to another, nor to compare them via predefined equality. There’s another reason you’ll see shortly. The type is private because that’s the default for good software engineering, and there’s no reason to override that default to make the implementation visible to clients. Our API will provide everything clients require, when combined with the capabilities provided by any limited type (e.g., declaring objects, and passing them as parameters). Tasks can wait for a single condition to be signaled, or they can wait for one of several. Similarly, tasks can signal one or more conditions at a time. Such groups of conditions are easily represented by an unconstrained array type: type Condition_List is array (Positive range <>) of Condition; We chose Positive as the index subtype because it allows a very large number of components, far more than is likely ever required. An aggregate value of the array type can then represent multiple conditions, for example: Condition_List'(A, C) Given these three types we can define a useful API: procedure Wait (This : in out Manager; Any_Of_These : Condition_List; Enabler : out Condition); -- Block until/unless any one of the conditions in Any_Of_These has been -- Signaled. The one enabling condition chosen will be returned in the Enabler -- parameter, and is cleared internally as Wait exits. Any other signaled -- conditions remain signaled. procedure Wait (This : in out Manager; This_One : Condition); -- Block until/unless the specified condition has been Signaled. This -- procedure is a convenience routine that can be used instead of an -- aggregate with only one condition component. procedure Signal (This : in out Manager; All_Of_These : Condition_List); -- Indicate that all of the conditions in All_Of_These are now set. The -- conditions remain set until cleared by Wait. procedure Signal (This : in out Manager; This_One : Condition); -- Indicate that This_One condition is now set. The condition remains set -- until cleared by Wait. This procedure is a convenience routine that can -- be used instead of an aggregate with only one condition component. Here’s a task that waits for either condition A or B, using a global Controller variable of the Manager type: task body A_or_B_Processor is Active : Condition; begin loop Wait (Controller, Any_Of_These => Condition_List'(A, B), Enabler => Active); Put_Line ("A_or_B_Processor responding to condition " & Active'Image); end loop; end A_or_B_Processor; When the call to Wait returns, at least one of either A or B was signaled. One of those signaled conditions is then selected and returned in the Enabler parameter. That selected condition is no longer signaled when the call returns, and will stay that way until another call to procedure Signal changes it. The other condition is not affected, whether or not it was also signaled. A signaling task could use the API to signal one condition: Signal (Controller, This_One => B); or to signal multiple conditions: Signal (Controller, All_Of_These => Condition_List'(A, C, D)); Now let’s consider the Manager implementation. As this is a concurrent program, we need it to be thread-safe. We’ve declared the Manager type as limited, so either a task type or a protected type would be allowed as the type’s completion. (That’s the other reason the type is limited.) There’s no need for this manager to do anything active, it just suspends some tasks and resumes others when called. Therefore, a protected type will suffice, rather than an active thread of control. Clearly, tasks that await conditions must suspend until a requested condition has been signaled, assuming it was not already signaled when the call occurred, so a protected procedure won’t suffice. Protected procedures only provide mutual exclusion. Hence we'll use a protected entry for the waiters to call. As you will see later, there is another reason to use protected entries here. Inside the Manager protected type we need a way to represent whether conditions have been signaled. We can use an array of Boolean components for this purpose, with the conditions as the indexes. For any given condition, if the corresponding array component is True the condition has been signaled, otherwise it has not. type Condition_States is array (Condition) of Boolean; Signaled : Condition_States := (others => False); Thus, for example, if Signaled (B) is True, a task that calls Wait for B will be able to return at once. Otherwise, that task will be blocked and cannot return from the call. Later another task will set Signaled (B) to True, and then the waiting task can be unblocked. Since an aggregate can also contain only one component if desired, we can use a single set of protected routines for waiting and signaling in the Manager protected type. We don't need one set of routines for waiting and signaling a single condition, and another set of routines for waiting and signaling multiple conditions. Here then is the visible part: protected type Manager is entry Wait (Any_Of_These : Condition_List; Enabler : out Condition); procedure Signal (All_Of_These : Condition_List); private … end Manager; Both the entry and the procedure take an argument of the array type, indicating one or more conditions. The entry, called by waiting tasks, also has an output argument, Enabler, indicating which specific entry enabled the task to resume, i.e., which condition was found signaled and was selected to unblock the task. We need that parameter because the task may have specified that any one of several conditions would suffice, and more than one could have been signaled. The bodies of our API routines are then just calls into the protected Manager argument. For example, here are two of the four: procedure Wait (This : in out Manager; Any_Of_These : Condition_List; Enabler : out Condition) is begin This.Wait (Any_Of_These, Enabler); end Wait; procedure Signal (This : in out Manager; This_One : Condition) is begin This.Signal (Condition_List'(1 => This_One)); end Signal; Now let’s examine the implementation of the protected type. It gets slightly complicated, but only a little. Our entry Wait allows a task to request suspension until one of the indicated conditions is signaled, as specified by the entry argument. Normally we’d expect to use the entry barrier to express this so-called “condition synchronization” by querying the conditions’ state array. If one of the requested conditions is True the barrier would allow the call to execute and complete. However, barriers do not have compile-time visibility to the entry parameters, so they cannot be referenced in the barriers. That's also true for the Boolean guards controlling task entry accept statements within select statements. Why not? Ada synchronization constructs are based on “avoidance synchronization,” meaning that 1) the user-written controls that enable/disable the execution of task entry accept statements and protected entry bodies are intended to enable them only when they can actually provide the requested service, and 2) that runtime determination is based on information known prior to the execution of the accept statement or entry body. For example, at runtime, if a bounded buffer is full, that fact can be determined from the buffer's state: is the count of contained items equal to the capacity of the backing array? If so, the controls disallow the operation to insert another value. Likewise, if the buffer is empty, the removal operation is disallowed. When we write the buffer implementation we know beforehand what the operations will try to do, so we can write the controls to disallow them at runtime until they can succeed. Most of the time that’s sufficient, but not always. When we can't know precisely what the operations will do when we write the code, avoidance synchronization doesn't fit the bill. That's the case with the condition manager: we don’t know beforehand which conditions the Wait caller will specify, and we can't refer to the parameters in the barrier, therefore we cannot use the barrier to enable or disable execution of the Wait entry body. To handle cases in which avoidance synchronization is insufficient Ada defines the “requeue” statement. Calling an entry that uses a requeue statement is much like calling a large company on the telephone. Calling the main number connects you to a receptionist (if you're lucky and don't get an annoying menu). If the receptionist can answer your question, they do so and then you both hang up. Otherwise, the receptionist forwards ("requeues") the call to the person you need to speak with. After doing so, the receptionist hangs up, because from their point of view the call is complete. The call is not complete from your point of view, though, until you finish your conversation with the new receiver. And of course you may have to wait to speak to that person. In this metaphor, a task calling entry Wait is the person calling the large corporation. Like the receptionist, Wait must take (execute) the call without knowing the requested conditions, because the entry barrier cannot reference the entry arguments. The specified conditions and their states are only known once the entry body executes. Therefore, Wait may or may not be able to allow the caller to return from the call immediately. If not, it requeues the call and finishes, leaving the call still pending on the requeue target. Because Wait always takes a call, the entry barrier is just hard-coded to True. (That’s always a strong indication that requeue is involved.) Even though this barrier always allows a call, much like a protected procedure, we must use an entry because only protected entries can requeue callers. Inside the entry body the specified conditions’ states are checked, looking for one that is True. If one is found, the entry body completes and the caller returns to continue further, responding to the found condition. If no requested condition is True, though, we cannot let the caller continue. We block it by requeueing the caller on to another entry. Eventually that other entry will allow the caller to return, when an awaited condition finally becomes True via Signal. Here then is the full declaration for the protected type Manager: type Condition_States is array (Condition) of Boolean; protected type Manager is entry Wait (Any_Of_These : Condition_List; Enabler : out Condition); procedure Signal (All_Of_These : Condition_List); private Signaled : Condition_States := (others => False); Prior_Retry_Calls : Natural := 0; entry Retry (Any_Of_These : Condition_List; Enabler : out Condition); end Manager; The private part contains the condition states, a management variable, and the other entry, Retry, onto which we will requeue when necessary. Note that this other entry is only meant to be called by a requeue from the visible entry Wait, so we declare it in the private part to ensure there are no other calls to it. Here’s the body of the entry Wait: entry Wait (Any_Of_These : Condition_List; Enabler : out Condition) when True is Found_Awaited_Condition : Boolean; begin Check_Conditions (Any_Of_These, Enabler, Found_Awaited_Condition); if not Found_Awaited_Condition then requeue Retry; end if; end Wait; The hard-coded entry barrier ("when True") always allows a caller to execute, subject to the mutual exclusion requirement. In the body, we call an internal procedure to check the state of the requested condition(s). If we don’t find one of the specified conditions True, we requeue the caller to the Retry entry. The entry parameters are the same in this case so they go to the Retry entry as usual. On the other hand, if we did find a specified condition True, we just exit the call, Enabler having been set already. Eventually, presumably, an awaited False condition will become True. That happens when Signal is called: procedure Signal (All_Of_These : Condition_List) is begin for C of All_Of_These loop Signaled (C) := True; end loop; Prior_Retry_Calls := Retry'Count; end Signal; After setting the specified conditions' states to True, Signal captures the number of queued callers waiting on Retry. (The variable Prior_Retry_Calls is an internal component declared in the protected type. The value is never presented to callers, but is, instead, used only to manage callers.) At long last, here’s the body of Retry: entry Retry (Any_Of_These : Condition_List; Enabler : out Condition) when Prior_Retry_Calls > 0 is Found_Enabling_Condition : Boolean; begin Prior_Retry_Calls := Prior_Retry_Calls - 1; Check_Conditions (Any_Of_These, Enabler, Found_Enabling_Condition); if not Found_Enabling_Condition then requeue Retry; end if; end Retry; Recall that when a protected procedure or entry exits a protected object, the run-time system re-evaluates all the object’s entry barriers, looking for an open (True) barrier with a caller queued, waiting. If one is found, that entry body is allowed to execute on behalf of that caller. On exit, the evaluation/execution process repeats. This process is known as a “protected action” and is one reason protected objects are so expressive and powerful. The protected action continues iterating, executing enabled entry bodies on behalf of queued callers, until either no barriers are open or no open barriers have callers waiting. Therefore, when procedure Signal sets Prior_Retry_Calls to a value greater than zero and then exits, the resulting protected action allows Retry to execute. Furthermore, Retry continues to execute, attempting to service all the prior callers in the protected action, because its barrier is False only when all those prior callers have been serviced. For each caller, Retry attempts the same thing Wait did: if a requested condition is True the caller is allowed to return from the call. Otherwise, the caller is requeued onto Retry. So yes, Retry requeues the caller onto itself! Doing so is not necessarily a problem, but in this case a caller would continue to be requeued indefinitely when the requested condition is False, unless something prevents that from happening. That’s the purpose of the count of prior callers. Only that number of callers are executed by the body of Retry in the protected action. After that the barrier is closed by Prior_Retry_Calls becoming zero, the protected action ceases when the entry body exits, and any unsatisfied callers remain queued. All well and good, this works, but have you noticed the underlying assumption? The code assumes that unsatisfied callers are placed onto the entry queue at the end of the queue, i.e., in FIFO order. Consequently, they are not included in the value of the Prior_Retry_Calls count and so do not get executed again until Signal is called again. But suppose we have requested that entry queues (among other things) are ordered by caller priority? We’d want that for a real-time system. But then a requeued caller would not go to the back of the entry queue and would, instead, execute all over again, repeatedly. The prior caller count wouldn’t solve that problem. If priority queuing might be used, we must change the internal implementation so that the queuing policy is irrelevant. We’ll still have Wait do a requeue when necessary, but no requeue will ever go to the same entry executing the requeue statement. Therefore, the entry queuing order won't make a difference. The entry family facilitates that change, and rather elegantly, too. An entry family is much like an array of entries, each one identical to the others. To work with any one of the entries we specify an index, as with an array. For example, here’s a requeue to Retry as a member of an entry family, with Active_Retry as the index: requeue Retry (Active_Retry) In the above, the caller uses the value of Active_Retry as an index to select a specific entry in the array/family. The resulting changes to the Manager type are as follows: type Retry_Entry_Id is mod 2; type Retry_Barriers is array (Retry_Entry_Id) of Boolean; protected type Manager is … as before private Signaled : Condition_States := (others => False); Retry_Enabled : Retry_Barriers := (others => False); Active_Retry : Retry_Entry_Id := Retry_Entry_Id'First; entry Retry (Retry_Entry_Id) (Any_Of_These : Condition_List; Enabler : out Condition); end Manager; Our entry family index type is Retry_Entry_Id. We happen to need two entries in this implementation, so a modular type with two values will suffice. Modular arithmetic will also express the logic nicely, as you’ll see. The variable Active_Retry is of this type, initialized to zero. The entry Retry is now a family, as indicated by the entry declaration syntax specifying the index type Retry_Entry_Id within parentheses. Each entry has the same parameters as any others in the family, in this case the same parameters as in the previous implementation. We thus have two Retry entries so that, at any given time, one of the entries can requeue onto the other one, instead of onto itself. An entry family makes that simple to express. At runtime, one of the Retry entries will field requeue calls from Wait, and will also be the entry enabled by Signal. That entry is designated the “active” retry target, via the index held in the local variable Active_Retry. Here’s the updated body of Wait accordingly: entry Wait (Any_Of_These : Condition_List; Enabler : out Condition) when True is Found_Enabling_Condition : Boolean; begin Check_Conditions (Any_Of_These, Enabler, Found_Enabling_Condition); if not Found_Enabling_Condition then requeue Retry (Active_Retry) with abort; end if; end Wait; The body is as before, except that the requeue target depends on the value of Active_Retry. (We'll discuss "with abort" shortly.) When Signal executes, it now enables the “active retry” entry barrier: procedure Signal (All_Of_These : Condition_List) is begin for C of All_Of_These loop Signaled (C) := True; end loop; Retry_Enabled (Active_Retry) := True; end Signal; The barrier variable Retry_Enabled is now an array, using the same index type as the entry family. The really interesting part of the implementation is the body of Retry, showing the expressive power of the language. The entry family member enabled by Signal goes through all its pending callers, attempting to satisfy them and requeuing those that it cannot. But instead of requeuing onto itself, it requeues them onto the other entry in the family. As a result, the order of the queues is immaterial. Again, the entry family makes this easy to express: entry Retry (for K in Retry_Entry_Id) (Any_Of_These : Condition_List; Enabler : out Condition) when Retry_Enabled (K) is Found_Enabling_Condition : Boolean; begin Check_Conditions (Any_Of_These, Enabler, Found_Enabling_Condition); if Found_Enabling_Condition then return; end if; -- otherwise... if Retry (K)'Count = 0 then -- current caller is the last one present Retry_Enabled (K) := False; Active_Retry := Active_Retry + 1; end if; requeue Retry (K + 1) with abort; end Retry; Note the first line: entry Retry (for K in Retry_Entry_Id) as well as the entry barrier (before the reserved word “is”): when Retry_Enabled (K) “K” is the entry family index, in this case iterating over all the values of Retry_Entry_Id (i.e., 0 .. 1). We don’t have to write a loop checking each family member’s barrier; that happens automatically, via K. When a barrier at index K is found to be True, that corresponding entry can execute a prior caller. Slick, isn’t it? Ada is a very powerful language. Note the last statement, the one performing the requeue: requeue Retry (K + 1) with abort; Like the Active_Retry variable, the index K is of the modular type with two possible values, so K + 1 is always the “other” entry of the two. The addition wraps around, conveniently. As a result, the requeue is always onto the other entry, never itself, so the entry queue ordering makes no difference. The “with abort” is important but is not a controlling design factor. In a nutshell, it means that task abort is enabled for the requeued task. The significance is that an aborted task that is suspended on an entry queue is removed from that queue. That’s allowable in this case because we are not using the count of prior callers to control the number of iterations in the protected action, unlike the FIFO implementation. In that other implementation we could not allow requeued tasks to be aborted because the count of prior callers would no longer match the number of queued callers actually present. The protected action would await a caller that would never execute. In this implementation that cannot happen so it is safe to allow aborted tasks to be removed from the queue. Note that we do still check the count of pending queued callers, we just don't capture it and use it to control the number of iterations in the protected action. If we’ve processed the last caller for member K, we close member K’s barrier immediately, and then set the active retry index to the other entry member so that Wait will requeue to the new “active retry” entry and Signal will, eventually, enable it. Because we did not make the implementation visible to the package’s clients, our internal changes will not require users to change any of their code. Note that both the Ravenscar and Jorvik profiles allow entry families, but Ravenscar only allows one member per family because only one entry is allowed per protected object. Such an entry family wouldn't be much use. Jorvik allows multiple entry family members because it allows multiple entries per protected object. However, neither profile allows requeue statements, for the sake of simplifying the underlying run-time library implementation. For more on tasking and topics like this, see the book by Burns and Wellings, Concurrent and Real-Time Programming In Ada, Cambridge University Press, 2007. Yes, 2007, but it is still excellent and directly applicable today. Indeed, this solution to the condition manager is based on their Resource_Controller example and was supplied to a customer this year. Thanks to Andrei Gritsenko (Андрей Гриценко@disqus_VErl9jPNvR) for suggesting a nice simplification of the FIFO version of the facility. This blog entry has been updated accordingly. The full code for the entry-family approach follows. Note that we have used a generic package so that we can factor out the specific kind of conditions involved, via the generic formal type. As long as the generic actual type is a discrete type the compiler will be happy. That’s essential because we use the condition type as an index for an array type. -- This package provides a means for blocking a calling task until/unless any -- one of an arbitrary set of "conditions" is signaled. -- NOTE: this implementation allows either priority-ordered or FIFO-ordered -- queuing. generic type Condition is (<>); package Condition_Management is type Manager is limited private; type Condition_List is array (Positive range <>) of Condition; procedure Wait (This : in out Manager; Any_Of_These : Condition_List; Enabler : out Condition); -- Block until/unless any one of the conditions in Any_Of_These has been -- Signaled. The one enabling condition will be returned in the Enabler -- parameter, and is cleared internally as Wait exits. Any other signaled -- conditions remain signaled. procedure Wait (This : in out Manager; This_One : Condition); -- Block until/unless the specified condition has been Signaled. This -- procedure is a convenience routine that can be used instead of an -- aggregate with only one condition component. procedure Signal (This : in out Manager; All_Of_These : Condition_List); -- Indicate that all of the conditions in All_Of_These are now set. The -- conditions remain set until cleared by Wait. procedure Signal (This : in out Manager; This_One : Condition); -- Indicate that This_One condition is now set. The condition remains set -- until cleared by Wait. This procedure is a convenience routine that can -- be used instead of an aggregate with only one condition component. private type Condition_States is array (Condition) of Boolean; type Retry_Entry_Id is mod 2; type Retry_Barriers is array (Retry_Entry_Id) of Boolean; protected type Manager is entry Wait (Any_Of_These : Condition_List; Enabler : out Condition); procedure Signal (All_Of_These : Condition_List); private Signaled : Condition_States := (others => False); Retry_Enabled : Retry_Barriers := (others => False); Active_Retry : Retry_Entry_Id := Retry_Entry_Id'First; entry Retry (Retry_Entry_Id) (Any_Of_These : Condition_List; Enabler : out Condition); end Manager; end Condition_Management; package body Condition_Management is ---------- -- Wait -- ---------- procedure Wait (This : in out Manager; Any_Of_These : Condition_List; Enabler : out Condition) is begin This.Wait (Any_Of_These, Enabler); end Wait; ---------- -- Wait -- ---------- procedure Wait (This : in out Manager; This_One : Condition) is Unused : Condition; begin This.Wait (Condition_List'(1 => This_One), Unused); end Wait; ------------ -- Signal -- ------------ procedure Signal (This : in out Manager; All_Of_These : Condition_List) is begin This.Signal (All_Of_These); end Signal; ------------ -- Signal -- ------------ procedure Signal (This : in out Manager; This_One : Condition) is begin This.Signal (Condition_List'(1 => This_One)); end Signal; ----------- -- Event -- ----------- protected body Manager is procedure Check_Conditions (These : Condition_List; Enabler : out Condition; Found : out Boolean); ---------- -- Wait -- ---------- entry Wait (Any_Of_These : Condition_List; Enabler : out Condition) when True is Found_Enabling_Condition : Boolean; begin Check_Conditions (Any_Of_These, Enabler, Found_Enabling_Condition); if not Found_Enabling_Condition then requeue Retry (Active_Retry) with abort; end if; end Wait; ------------ -- Signal -- ------------ procedure Signal (All_Of_These : Condition_List) is begin for C of All_Of_These loop Signaled (C) := True; end loop; Retry_Enabled (Active_Retry) := True; end Signal; ----------- -- Retry -- ----------- entry Retry (for K in Retry_Entry_Id) (Any_Of_These : Condition_List; Enabler : out Condition) when Retry_Enabled (K) is Found_Enabling_Condition : Boolean; begin Check_Conditions (Any_Of_These, Enabler, Found_Enabling_Condition); if Found_Enabling_Condition then return; end if; -- otherwise... if Retry (K)'Count = 0 then -- current caller is the last one present Retry_Enabled (K) := False; Active_Retry := Active_Retry + 1; end if; requeue Retry (K + 1) with abort; end Retry; ---------------------- -- Check_Conditions -- ---------------------- procedure Check_Conditions (These : Condition_List; Enabler : out Condition; Found : out Boolean) is begin for C of These loop if Signaled (C) then Signaled (C) := False; Enabler := C; Found := True; return; end if; end loop; Found := False; end Check_Conditions; end Manager; end Condition_Management; ]]> Showing Global contracts with GNAT Studio https://blog.adacore.com/showing-global-contracts-with-gnat-studio Fri, 26 Feb 2021 00:30:00 -0500 Simon Buist https://blog.adacore.com/showing-global-contracts-with-gnat-studio In SPARK, data-flow analysis performs 2 steps: (1) verifying the user-defined data-flow contracts (e.g. Global / Depends) and (2) generating them when they are missing. Accurate data-flow analysis is a necessary prerequisite for proof of absence of run-time-errors (AoRTE) checks. SPARK 2005 only did the first step (assuming null global data in the absence of user-defined global contracts). The second step, generation, is useful for users that want to get the benefits of flow-analysis, including proven AoRTE, without bothering to annotate their code with contracts. However, the global contracts that get generated may not meet expectations / requirements. For example, you may write to a variable that should only be read from. In this case, flow analysis and AoRTE proof could pass, but your code would not meet its requirements. So there is a need to see what global contracts SPARK generates. The generated global contracts were previously hidden from the user. They could be exposed with the switch --flow-show-gg, but then SPARK would output the generated global contracts to the console, which made them hard to utilise effectively. In the integrated development environment, GNAT Studio, there is now a plugin that inserts the generated global contracts inline with the code. The GNAT Studio contextual menu Here, we have a screenshot of some Ada source code being edited in GNAT Studio. The code uses two globals, G1 and G2. Function Potato reads both. Function Kitty reads G1. The user has right-clicked to select the new plugin, “Show generated Global contracts”. Once the user clicks “Show generated Global contracts”, the generated global contracts get inserted into the editor window, so that the user can see what SPARK data-flow analysis detects as globals. The user can then inspect the globals, and if they wish, copy-paste from this into their code, to add contracts. From this point, they can check whether the system fulfills global data-flow requirements. We see this as a great learning tool for beginners to SPARK. Tokeneer Let’s test the plugin on the Tokeneer project - a codebase that makes extensive use of Globals. Tokeneer has been fully verified in SPARK, and has a comprehensive set of user-written Global contracts, so we need to hide these contracts from our plugin by adding the following directives to the file we want to test: pragma Ignore_Pragma (Global); pragma Ignore_Pragma (Refined_Global); pragma Ignore_Pragma (Depends); pragma Ignore_Pragma (Refined_Depends);  This will allow us to compare the user-written Global contracts with the ones displayed by our plugin. The results: File enclave.adb has a procedure named ValidateAdminToken: We can see that the user-supplied Global contract (white background) closely matches the generated one (grey background). GNATprove has resolved Output => Status conservatively as In_Out => Enclave.Status. The plugin correctly summarises the effect on individual variables by the aggregated effect on the corresponding abstract state, when the variables belong to one. We chose to prefix Global variable names with their full path, even when they are inside the current program unit. We decided to do this to disambiguate in such situations where we have two different variables sharing the same name, for example where we have Global variable X in package Outer, and Global variable X in package Inner, with Inner nested inside Outer: Coming back to the Tokeneer code, we also ran the plugin on file keystore.adb, which has a function named GetBlock: The plugin has correctly identified that there is no Global state. It’s useful to annotate this function with Global => null, so that we know it is not modifying any Global state. We’d love to know how you use this feature and if you see any useful enhancements. One feature we could add is the ability to automatically insert a generated contract in the code if the user wishes so. ]]> Doubling the Performance of SPARKNaCl (again...) https://blog.adacore.com/doubling-the-performance-of-sparknacl-again Thu, 18 Feb 2021 00:00:00 -0500 Roderick Chapman https://blog.adacore.com/doubling-the-performance-of-sparknacl-again function Product_To_Seminormal (X : in Product_GF) return Seminormal_GF with Pure_Function, Global => null; -- "LM" = "Limb Modulus" -- "LMM1" = "Limb Modulus Minus 1" LM : constant := 65536; LMM1 : constant := 65535; -- "R2256" = "Remainder of 2**256 (modulo 2**255-19)" R2256 : constant := 38; -- "Maximum GF Limb Coefficient" MGFLC : constant := (R2256 * 15) + 1; -- "Maximum GF Limb Product" MGFLP : constant := LMM1 * LMM1; subtype Product_GF is GF with Dynamic_Predicate => (for all I in Index_16 => Product_GF (I) >= 0 and Product_GF (I) <= (MGFLC - 37 * GF_Any_Limb (I)) * MGFLP); ---------------------------------------------------------------------- -- A "Seminormal GF" is the result of applying a single -- normalization step to a Product_GF -- -- Least Significant Limb ("LSL") of a Seminormal GF. -- LSL is initially normalized to 0 .. 65535, but gets -- R2256 * Carry added to it, where Carry is (Limb 15 / 65536) -- The upper-bound on Limb 15 is given by substituting I = 14 -- into the Dynamic_Predicate above, so -- (MGFLC - 37 * 14) * MGFLP = 53 * MGFLP -- See the body of Product_To_Seminormal for the full -- proof of this upper-bound subtype Seminormal_GF_LSL is I64 range 0 .. (LMM1 + R2256 * ((53 * MGFLP) / LM)); -- Limbs 1 though 15 are in 0 .. 65535, but the -- Least Significant Limb 0 is in GF_Seminormal_Product_LSL subtype Seminormal_GF is GF with Dynamic_Predicate => (Seminormal_GF (0) in Seminormal_GF_LSL and (for all I in Index_16 range 1 .. 15 => Seminormal_GF (I) in GF_Normal_Limb)); ]]> Performance analysis and tuning of SPARKNaCl https://blog.adacore.com/performance-analysis-and-tuning-of-sparknacl Tue, 09 Feb 2021 05:29:00 -0500 Roderick Chapman https://blog.adacore.com/performance-analysis-and-tuning-of-sparknacl #define FOR(i,n) for (i = 0;i < n;++i) #define sv static void typedef long long i64; typedef i64 gf[16]; sv car25519(gf o) { int i; i64 c; FOR(i,16) { o[i]+=(1LL<<16); c=o[i]>>16; o[(i+1)*(i<15)]+=c-1+37*(c-1)*(i==15); o[i]-=c<<16; } } if (i == 15) o[0] += c-1+37*(c-1); else o[i+1] += c-1; -- returns equivalent of X >> 16 in C, doing an arithmetic -- shift right when X is negative, assuming 2's complement -- representation function ASR_16 (X : in I64) return I64 is (To_I64 (Shift_Right_Arithmetic (To_U64 (X), 16))) with Post => (if X >= 0 then ASR_16'Result = X / LM else ASR_16'Result = ((X + 1) / LM) - 1); function Car (X : in GF) return GF is Carry : I64; R : GF; begin R := X; for I in Index_16 range 0 .. 14 loop Carry := ASR_16 (R (I)); R (I + 1) := R (I + 1) + Carry; R (I) := R (I) mod LM; end loop; Carry := ASR_16 (R (15)); R (0) := R (0) + R2256 * Carry; R (15) := R (15) mod LM; return R; end Car; is Carry : I64; R : GF with Relaxed_Initialization; begin Carry := ASR_16 (X (0)); R (0) := X (0) mod LM; R (1) := X (1) + Carry; pragma Assert (R (0)'Initialized and R (1)'Initialized); for I in Index_16 range 1 .. 14 loop Carry := ASR_16 (R (I)); R (I) := R (I) mod LM; R (I + 1) := X (I + 1) + Carry; pragma Loop_Invariant (for all K in Index_16 range 0 .. I => R (K)'Initialized); pragma Loop_Invariant (R (I + 1)'Initialized); end loop; pragma Assert (R'Initialized); -- as before... function Scalarmult (Q : in GF_Vector_4; S : in Bytes_32) return GF_Vector_4 with Global => null; function Scalarmult (Q : in GF_Vector_4; S : in Bytes_32) return GF_Vector_4 is CB : Byte; Swap : Boolean; LP, LQ : GF_Vector_4; begin LP := (0 => GF_0, 1 => GF_1, 2 => GF_1, 3 => GF_0); LQ := Q; for I in reverse U32 range 0 .. 255 loop CB := S (I32 (Shift_Right (I, 3))); Swap := Boolean'Val (Shift_Right (CB, Natural (I and 7)) mod 2); CSwap (LP, LQ, Swap); -- Note user-defined "+" for GF_Vector_4 called here LQ := LQ + LP; LP := LP + LP; CSwap (LP, LQ, Swap); end loop; return LP; end Scalarmult; -- For each byte of S, starting at the MSB for I in reverse Index_32 loop -- For each bit, starting with bit 7 (the MSB) for J in reverse Natural range 0 .. 7 loop CB := S (I); Swap := Boolean'Val (Shift_Right (CB, J) mod 2); CSwap (LP, LQ, Swap); LQ := LQ + LP; LP := LP + LP; CSwap (LP, LQ, Swap); end loop; end loop; -- For each byte of S, starting at the MSB for I in reverse Index_32 loop CB := S (I); -- For each bit of CB, starting with bit 7 (the MSB) for J in reverse Natural range 0 .. 7 loop Swap := Boolean'Val (Shift_Right (CB, J) mod 2); CSwap (LP, LQ, Swap); LQ := LQ + LP; LP := LP + LP; CSwap (LP, LQ, Swap); end loop; end loop; function "*" (Left, Right : in Normal_GF) return Normal_GF is T : GF_PA; -- 31 digits begin T := (others => 0); -- "Textbook" ladder multiplication for I in Index_16 loop for J in Index_16 loop T (I + J) := T (I + J) + (Left (I) * Right (J)); end loop; end loop; T (I) := T (I) + (Left (I) * Right (0)); T (I + 1) := T (I + 1) + (Left (I) * Right (1)); T (I + 2) := T (I + 2) + (Left (I) * Right (2)); -- and so on... LT := Left (I); T (I) := T (I) + (LT * Right (0)); T (I + 1) := T (I + 1) + (LT * Right (1)); T (I + 2) := T (I + 2) + (LT * Right (2)); -- and so on... function DL64 (X : in Bytes_8) return U64 is U : U64 := 0; begin for I in X'Range loop U := Shift_Left (U, 8) or U64 (X (I)); end loop; return U; end DL64; W := (0 => DL64 (Bytes_8 (M (CB + 0 .. CB + 7))), 1 => DL64 (Bytes_8 (M (CB + 8 .. CB + 15))), 2 => DL64 (Bytes_8 (M (CB + 16 .. CB + 23))), 3 => DL64 (Bytes_8 (M (CB + 24 .. CB + 31))), -- and so on... 15 => DL64 (Bytes_8 (M (CB + 120 .. CB + 127)))); function DL64 (X : in Byte_Seq; I : in N32) return U64 with Global => null, Pre => X'Length >= 8 and then I >= X'First and then I <= X'Last - 7; function DL64 (X : in Byte_Seq; I : in N32) return U64 is LSW, MSW : U32; begin -- Doing this in two 32-bit groups avoids the need -- for 64-bit shifts on 32-bit machines. MSW := Shift_Left (U32 (X (I)), 24) or Shift_Left (U32 (X (I + 1)), 16) or Shift_Left (U32 (X (I + 2)), 8) or U32 (X (I + 3)); LSW := Shift_Left (U32 (X (I + 4)), 24) or Shift_Left (U32 (X (I + 5)), 16) or Shift_Left (U32 (X (I + 6)), 8) or U32 (X (I + 7)); -- Just one 64-bit shift and an "or" is now required return Shift_Left (U64 (MSW), 32) or U64 (LSW); end DL64; W := (0 => DL64 (M, CB), 1 => DL64 (M, CB + 8), 2 => DL64 (M, CB + 16), -- and so on... 15 => DL64 (M, CB + 120)); A := F (B, C); function "+" (Left, Right : in GF_Vector_4) return GF_Vector_4 is A, B, C, D, E, F, G, H : Normal_GF; function Double (X : in Normal_GF) return Normal_GF is (X + X) with Global => null; begin A := (Left (1) - Left (0)) * (Right (1) - Right (0)); B := (Left (0) + Left (1)) * (Right (0) + Right (1)); C := (Left (3) * Right (3)) * GF_D2; D := Double (Left (2) * Right (2)); E := D - C; F := D + C; G := B - A; H := B + A; return GF_Vector_4'(0 => G * E, 1 => H * F, 2 => F * E, 3 => G * H); end "+"; e = sparknacl."-" (&d, &c); f = sparknacl."+" (&d, &c); g = sparknacl."-" (&b, &a); h = sparknacl."+" (&b, &a); R.6 = sparknacl."*" (&g, &e); [return slot optimization] <retval>[0] = R.6; R.7 = sparknacl."*" (&h, &f); [return slot optimization] <retval>[1] = R.7; R.8 = sparknacl."*" (&f, &e); [return slot optimization] <retval>[2] = R.8; R.9 = sparknacl."*" (&g, &h); [return slot optimization] <retval>[3] = R.9; return <retval>; e = sparknacl."-" (&d, &c); [return slot optimization] f = sparknacl."+" (&d, &c); [return slot optimization] g = sparknacl."-" (&b, &a); [return slot optimization] h = sparknacl."+" (&b, &a); [return slot optimization] <retval>[0] = sparknacl."*" (&g, &e); [return slot optimization] <retval>[1] = sparknacl."*" (&h, &f); [return slot optimization] <retval>[2] = sparknacl."*" (&f, &e); [return slot optimization] <retval>[3] = sparknacl."*" (&g, &h); [return slot optimization] return <retval>; ]]> AdaCore at FOSDEM 2021 https://blog.adacore.com/adacore-at-fosdem-2021 Thu, 04 Feb 2021 03:55:00 -0500 Fabien Chouteau https://blog.adacore.com/adacore-at-fosdem-2021 Like previous years, AdaCore will participate in FOSDEM. This time the event will be online only, but this won’t prevent us from celebrating Open Source software. AdaCore engineers will give two talks in the Safety and Open Source devroom, a topic at the heart of AdaCore since its inception: Hope to see you (virtually) at FOSDEM this week-end! ]]> Mini SAM M4 Ada BSP https://blog.adacore.com/mini-sam-m4-ada-bsp Tue, 02 Feb 2021 03:40:00 -0500 Fabien Chouteau https://blog.adacore.com/mini-sam-m4-ada-bsp $ alr get minisamd51_example
$cd minisamd51_example*$ alr build
$bash convert_to_uf2.sh ]]> How To: GNAT Pro with Docker https://blog.adacore.com/how-to-gnat-pro-with-docker Fri, 22 Jan 2021 00:00:00 -0500 Léo Germond https://blog.adacore.com/how-to-gnat-pro-with-docker Using GNAT Pro with containerization technologies, such as Docker, is so easy, a whale could do it! In this article I will show you how to get started setting up a Docker image with the GNAT Pro tools. But first... What is Docker? Maybe a better question: What isn’t Docker these days… Am I right? Over the last decade, DevOps methodology has revolutionized software engineering as we know it. A fundamental concept, Infrastructure as Code, has helped us build software that matters in a reliable, repeatable, and performant way, providing an actionable solution to one of the hardest IT problems: How can I ensure my build process is stable and repeatable? Docker is one such tool (one we really like here at AdaCore) that allows us to define our IT infrastructure as code in a highly portable way. It uses an exhaustive approach to configuration using layers, which means it is easier than ever to set up build environments that share dependencies, a common occurrence for build systems. Furthermore, its decentralized architecture will work independently from any external infrastructure, an argument that has gained weight as fast as a whale recently. Lastly, a small piece of advice: keep in mind that Docker instances are containers. While powerful, some things aren’t set up out of the box, such as sharing a USB port or a host system directory. In order to do these things you'll need to do some host configuration work. Rule of thumb: if you need more than a ssh to the tools inside the container, e.g. to run a JTAG debugger through Docker, it will require some leg work. Getting Started As a prerequisite, you must have a recent version of Docker, a GNAT Pro release package for Linux x86-64, and a Python 3 install. The procedure and script works with the current stable native compiler 20.2 version, as well as in-stabilization 21.1 and wavefront 22.0. Older versions or cross toolchains may require some additional work. First, let’s start with a sample Dockerfile. An example can be found in the AdaCore GitHub's GNAT Docker repository. The repository contains two directories: • gnatpro-deps/ Resources for building the GNAT Pro base image. • gnatpro/ Resources for building the full GNAT Pro toolsuite image. For technical reasons, the actual build is performed in two steps: we set up a build environment in gnatpro-deps/ then use it to build the GNAT Pro release, that we install in gnatpro/. All of this is done by the create_image script. Creating a Docker Image Using create_image The create_image Python script starts by creating a gnatpro:deps image containing the necessary tooling to build GNAT Pro from a release package. Then using a provided GNAT Pro release archive, it builds a second image with GNAT Pro actually installed. Finally it tags this image as gnatpro:NN.N, with NN.N being the GNAT Pro version number, so images of several GNAT Pro versions can coexist side-by-side. If we try to run the create_image Python script, we can see that the GNAT Pro release package file must be provided, and an optional GNAT Pro version number. When provided, the version number will be used to tag the GNAT Pro image. Otherwise, it will be gathered from the filename, when possible. usage: create_image [-h] [--verbose] [--gnat_version GNAT_VERSION] gnatpro_release positional arguments: gnat_release GNAT Pro release package file optional arguments: -h, --help show this help message and exit --verbose, -v Display commands as they are run --gnat_version GNAT_VERSION GNAT Pro version number for automatic tagging and archive search. Leave empty for the script to infer it. In our example we're using a 20.2 stable GNAT Pro version: • The --gnat_version argument should therefore be 20.2, and we provide the path to gnatpro-20.2-x86_64-linux-bin.tar.gz which contains the GNAT Pro installer. • If we set the --verbose flag, we can see in real time the commands as they are called by the script. From the gnatpro-deps/ directory, it builds the gnatpro:deps image, which takes no arguments. Then, it builds the full GNAT Pro image. This is done in two steps: First it copies the release package into the gnatpro/ directory (for technical reasons relative to Docker’s “copy” command). It then calls docker build. This command does the heavy lifting by running the building steps described in the file gnatpro/Dockerfile. The format of this file is defined in the Docker builder doc. It accepts arguments, and in our case one is mandatory: gnat_release, which should point to the GNAT Pro release package so the create_image script can use it as an argument to the docker build command. Once this step is performed, the image is complete, and is tagged as gnatpro:20.2. ./create_image --verbose --gnat_version=20.2 gnatpro/gnatpro-20.2-x86_64-linux-bin.tar.gz Docker for build dependencies: image gnatpro:deps > docker build -t gnatpro:deps gnatpro-deps Sending build context to Docker daemon 2.048kB Step 1/2 : FROM ubuntu:18.04 ---> 56def654ec22 Step 2/2 : RUN set -xe && DEBIAN_FRONTEND=noninteractive [...] Get:1 http://archive.ubuntu.com/ubuntu bionic InRelease [242 kB] [...] ---> 89afd2c07618 Successfully built 89afd2c07618 Successfully tagged gnatpro:20.2 GNAT Pro image built successfully You can open a shell on it with the command docker run --entrypoint bash -it gnatpro:20.2 Do you want to build and run the GNAT example ? [yN] The script finishes by asking to test the newly built image. Input Y to start the test. The test builds all the GNAT Pro examples and checks that they compile and run without error. Do you want to build and run the GNAT example ? [yN] y > docker run --entrypoint make -t gnatpro:20.2 -C /usr/gnat/share [...] make: Entering directory '/usr/gnat/share/examples/gnat' make -C starter make[1]: Entering directory '/usr/gnat/share/examples/gnat/starter' gnatmake -g hello gcc -c -g hello.adb gnatbind -x hello.ali gnatlink hello.ali -g ./hello Hello World. Welcome to GNAT gnatmake -g demo1 [...] Bind [gprbind] use_of_import.bexch [Ada] use_of_import.ali Link [archive] libimport_from_c.a [index] libimport_from_c.a [link] use_of_import.adb import_from_c/use_of_import I am now in the imported function make[1]: Leaving directory '/usr/gnat/share/examples/gnat/other_languages' make: Leaving directory '/usr/gnat/share/examples/gnat' gcc -c -g demo1.adb gcc -c -g gen_list.adb gcc -c -g instr.adb gnatbind -x demo1.ali gnatlink demo1.ali -g Notice how the examples are running on the container. At this point, we have a Docker image complete with GNAT Pro! Using the Docker CLI, we can see the gnatpro:deps and gnatpro:20.2 images. $ docker image ls
REPOSITORY   TAG       IMAGE ID       CREATED             SIZE
gnatpro      20.2      89afd2c07618   4 minutes ago       1.19GB
gnatpro      deps      9f1bdb3dbef0   15 minutes ago      87.9MB

Only the gnatpro:20.2 image is necessary for using the GNAT Pro toolset. The gnatpro:deps image can be removed if you don't need to build images for other GNAT Pro versions.

How to Use This Example

Using these resources as a template, you'll be able to quickly get a working GNAT Pro toolchain running in a Docker container.

On top of that foundation, it is possible to build on-demand CI matrices, scalable compilation jobs, and on-demand analysis services with tools like CodePeer, GNATcoverage and SPARK Pro.

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Ada on any ARM Cortex-M device, in just a couple minutes https://blog.adacore.com/ada-on-any-arm-cortex-m-device-in-just-a-couple-minutes Mon, 11 Jan 2021 00:00:00 -0500 Fabien Chouteau https://blog.adacore.com/ada-on-any-arm-cortex-m-device-in-just-a-couple-minutes <device Dname="STM32F407VG"> <memory id="IROM1" start="0x08000000" size="0x00100000" startup="1" default="1"/> <memory id="IRAM1" start="0x20000000" size="0x00020000" init ="0" default="1"/> <memory id="IRAM2" start="0x10000000" size="0x00010000" init ="0" default="0"/>
project Hello is

for Languages use ("Ada", "ASM_CPP"); -- ASM_CPP to compile the startup code
for Source_Dirs use ("src");
for Object_Dir use "obj";

for Target use "arm-eabi";

--  generic ZFP run-time compatible with our MCU

--  Linker script generated by startup-gen

package Device_Configuration is

--  Name of the CPU core on the STM32F407
for CPU_Name use "ARM Cortex-M4F";

for Float_Handling use "hard";

--  Number of interrupt lines on the STM32F407
for Number_Of_Interrupts use "82";

--  List of memory banks on the STM32F407
for Memories use ("SRAM", "FLASH", "CCM");

--  Specify from which memory bank the program will load
for Boot_Memory use "FLASH";

--  Specification of the SRAM
for Mem_Kind ("SRAM") use "ram";
for Size ("SRAM") use "128K";

--  Specification of the FLASH
for Mem_Kind ("FLASH") use "rom";
for Size ("FLASH") use "1024K";

--  Specification of the CCM RAM
for Mem_Kind ("CCM") use "ram";
for Size ("CCM") use "64K";

end Device_Configuration;
end Hello;
with Ada.Text_IO;

procedure Hello is
begin
end Hello;
project Prj is

type Boot_Mem is ("flash", "sram");
Boot : Boot_Mem := external ("BOOT_MEM", "flash");

package Device_Configuration is

for Memories use ("flash", "sram");

for Boot_Memory use Boot;

--  [...]
end Device_Configuration;
end Prj;
project Prj is

type Board_Kind is ("dev_board", "production_board");
Board : Board_Kind := external ("BOARD", "dev_board");

package Device_Configuration is

for Memories use ("flash", "sram");

case Board is
when "dev_board" =>
for Size ("sram")     use "256K";
when "production_board" =>
for Size ("sram")     use "128K";
end case;

--  [...]
end Device_Configuration;
end Prj;
]]>

Some of you may recall an AdaCore blog post written in 2017 by Thales engineer Lionel Matias titled "Leveraging Ada Run-Time Checks with Fuzz Testing in AFL". This insightful post took us on a journey of discovery as Lionel demonstrated how Ada programs, compiled using GNAT Pro and an adapted assembler pass can be subjected to advanced fuzz testing. In order to achieve this Lionel demonstrated how instrumentation of the generated assembly code around jump and label instructions, could be subjected to grey-box (path aware) fuzz testing (using the original AFL v2.52b as the fuzz engine). Lionel explained how applying the comprehensive spectrum of Ada runtime checks, in conjunction with Ada's strong typing and contract based programming, enhanced the capabilities of fuzz testing beyond the abilities of other languages. Ada's advanced runtime checking, for exceptions like overflows, and the scrutiny of Ada's design by contract assertions allow corner case bugs to be found whilst also utilising fuzz testing to verify functional correctness.

The success of Thales' initial research work and the obvious potential for this technology (see the blog post for real world examples of the bugs Lionel found), coupled with the evolution of fuzzing tools from the wider community and the impressive bug finding success of AFL, lead to an ongoing AdaCore research and development campaign around advanced security hardening through vulnerability testing.

This blog post provides the reader with an update into the science of vulnerability testing and describes some of the projects AdaCore has undertaken around fuzzing technologies as we work towards an industrial grade fuzz testing solution for Ada applications.

In addition this blog post celebrates the recent inclusion of AdaCore's GCC plugin into the publicly available latest release of AFLplusplus (version 3.00c).

A Brief Introduction to Fuzzing

A modern-day definition of fuzzing would be a testing capability designed to automatically create and inject inputs into a software process under test, while monitoring and recording detected faults. When considered, as part of a software development strategy, fuzzing is widely regarded as ‘negative testing’ and tends to fall under the remit of ‘software robustness’. More recently, however, it has been widely associated with ‘cyber-security’ and proven to be a suitable mechanism for finding code vulnerabilities that could otherwise lead to malicious exploitation. In this context, it is often known as ‘vulnerability testing’. In addition the term 'Refutation Testing', used in the context of aerospace security airworthiness and used to describe the act of 'refuting that our system is not secure', can also be used to describe fuzz testing.

Fuzzing, as a robustness strategy, has evolved exponentially since its initial inception back in the 1980s. Although the goal of fuzzing is always the same, tool development has forked in many directions such that different fuzzing tools use different test strategies and vary significantly in sophistication and complexity. Simplistic ‘black box’ fuzzers rely on brute force and speed of test case generation and execution. Whilst others instrument the code under test before, or during, compilation to allow the fuzzing engine to understand execution flow around decision points and to help guide the mutation engine onto diverging paths.

One of the main features of fuzz testing is rapid and automated test case generation and, as is the case within AFLplusplus, this is often achieved by a series of mutation phases over a starting corpus of test case files. The mutations tend to be 'non-structure aware' and instead alter the test case at the binary level by performing bit flips. That being said AFLplusplus does support an impressive custom mutation algorithm API that we will talk about later. Mutated test cases are then injected into the system under test while the fuzzer simultaneously monitors the application to detect a hung process or core dump. Test cases that resulted in a software fault are then placed in a separate directory for later manual scrutiny. This approach provides flexibility as the engine doesn't need to understand the format of the input data.

Instrumentation Guided Genetic Algorithms

AFLplusplus understands, by using test instrumentation applied during code compilation, when a test case has found a new path (increased coverage) and places that test case onto a queue for further mutation, injection and analysis. The fuzzing driver sets up a small shared memory area for the tested program to store execution path signatures. This structure is a basic array of counters incremented (or set) as execution reaches instrumentation points. Each instrumentation point is assigned, at compile time, a pseudo-random identifier. The array element to be incremented at run time is determined by hashing ordered pairs of identifiers of previous and current instrumentation points, to count the traversal of edges.

Though the driver can spawn (fork and exec) a new test process for each execution, instrumentation of the test program introduces a fork server mode that maps the shared memory and forks a new test process when requested by the driver. This avoids the overhead of exec and of running global constructors.

The aim of the instrumentation is to focus on non-intrusive manipulation; the fuzzer does not rely on complex code analysis or other forms of constraint solving. This form of 'path awareness' fuzzing, using an instrumentation-guided genetic algorithm, focuses the mutations on the test cases that discover path divergence and allows the testing to explore deeper into the control flow.

Figure 1 presents a basic representation of a Grey Box Fuzz Testing Architecture.

Mutation Strategies - "One size doesn't fit all"

Mutators applied to test inputs can be deterministic or nondeterministic. This flexibility encompasses changes suitable for both text and binary file formats. Mutators involve flipping single or multiple contiguous bits, arithmetic add and subtract at various word widths and endiannesses; overwrites, inserts and deletes. Deterministic mutators iterate over all bits or bytes of an input file, generating one or multiple new potential input tests at each iteration point. Nondeterministic approaches may perform multiple random mutations at random points and with random operands. In addition, chunks from another test may be spliced into the current before another round of random mutations. Some arithmetic mutators use a selection of “interesting” values, while string mutators use blocks of constant values. These are copies of chunks from elsewhere in the test input, fuzzer-identified tokens, and user-supplied keywords.

In figure 2 we can see a basic example of a deterministic mutation strategy and a non-deterministic strategy. The deterministic strategy involves flipping each bit within the test case to produce a new test case whilst the non-deterministic approach shows random sequences of bits being swapped around whilst other sequences are flipped at the same time. However in most cases the mutation engines will actually utilise pseudo-random number generators which are seeded to ensure patterns are repeatable.

An Enhanced GCC Instrumentation Approach

American Fuzzy Lop (AFL) was a revolutionary game changer for fuzzing technologies. For many security engineers and researchers it quickly became the de facto standard for how to implement a generic grey box fuzzing engine. AFL paved the way for the genetic algorithm approach to test data mutation through execution path awareness and without compromise of test execution frame rates. Coupled with an easy to understand interface and high configurability and the fact that AFL has always been 'free software' (in that users have the freedom to run, copy, distribute, study, change and improve the software), lead to the discovery and patching of multiple vulnerabilities within released, and widely used, commercial software applications. AFL can also support full black box testing by utilising modern emulator technology which allows vulnerabilities to be discovered on binary executables when the source code and / or build environment is not available.

Although AFL v2.52b offered a post-processor to instrument assembly output from the compiler, it was target architecture specific, and introduced significant run-time overhead. AFL v2.52b did offer up a pre-existing plugin for LLVM with much lower-overhead instrumentation, but GCC users had to tolerate higher instrumentation overheads. In order to ensure that GCC compiled applications could benefit from a low-overhead and optimizable AFL instrumentation AdaCore took inspiration from an existing plugin, originally developed for instrumenting programs written in 'C', to create a new feature rich and modern GCC plugin for the AFL fuzzing engine.

Building the program with instrumentation for AFL amounts to inserting an additional stage of processing whereby the compiler's assembly code output is modified, before running the actual assembler. This additional build stage introduces global variables, a fork-server and edge counting infrastructure in the translation unit that defines the main function and calls to an edge counting function after labels and conditional branches. In order to not disturb the normal execution of the program, each such call must preserve all registers modified by the edge counter, and by the call sequence, e.g. to pass the identifier of the current instrumentation point as an argument to the edge counter. The edge counter loads the previous instrumentation point identifier from a global variable and updates it, in addition to incrementing the signature array element associated with the edge between them.

Integrating the additional instrumentation stages into the compiler enables several improvements, as demonstrated by the pre-existing plugin for LLVM, and a newly introduced one for GCC. Besides portability to multiple target architectures, this move enables the instrumentation to be integrated more smoothly within the code stream. For example, while the assembler post-processor, with its limited knowledge, has to conservatively assume several registers need to be preserved across the edge counter calls, the compiler can take the calls into account when performing register allocation or, even better, avoid the call overhead entirely and perform the edge counting inline.

Using compiler infrastructure also makes it easier to use a thread-local variable to hold the previous instrumentation point identifier, so that programs with multiple threads do not confuse the fuzzer with cross-thread edges. The sooner in the compiler pipeline instrumentation is introduced, the more closely it will resemble the source code structure, and the more opportunities to optimize the instrumentation there will be. Conversely, the later it is introduced, the more closely it will resemble the executable code blocks after optimization, so duplicated blocks, e.g. resulting from loop unrolling, get different identifiers, and merged blocks, e.g. after jump threading, may end up with fewer instrumentation points.

AFLplusplus

AFLplusplus takes up the mantle from where AFL v2.52b left off adding various enhancements to the capability whilst also offering up a highly customisable, tool chain-able and extendable approach through it's modular architecture.

One particular feature of AFLplusplus of interest is the ability to add custom mutator algorithms which can either extend, sanitise or completely replace the AFL internal mutation engine.

This, along with other features and a strong interest in industry lead by Thales, guided AdaCore towards ensuring our GCC plugin was compatible with AFLplusplus and ensuring that new features, developed into the existing AFLplusplus GCC plugin, were reimplemented in the AdaCore developed plugin.

This work has now been completed and in partnership with Thales and with support from the AFLplusplus maintainer team AdaCore have since upstreamed the work into the publicly available AFLplusplus repository.

Or to quote the AFLplusplus public github repo:

Programming languages come with a variety of runtime libraries which are responsible for abstracting away fundamental architecture required to execute program instructions. By design each runtime is matched to the predicted role of the programming language and Ada, which is aimed at safety and security systems, is no exception. Runtimes, like the Ada runtime, that support advanced constraint checking are particularly well suited to fuzz testing. Languages, like 'C', with runtimes that allow some constraints to go unchecked are less well suited.

The following examples explain this better:

C program that receives a password as a command line argument and displays secret information if the password = "A*^%bd0eK":

#include <string.h>
#include <stdio.h>

{
int Privileged = 0;

if (strcmp (Password_Buff, "A*^%bd0eK") == 0)
{
Privileged = 1;
}

if (Privileged != 0)
{
printf ("Shhhh... the answer to life the universe and everything is 42!\n");
} else {
printf ("Pah. There is no getting passed this system's hardened security layer...\n");
}
}

int main(int argc, char **argv)
{
return 0;
}

Ada equivalent version (disclaimer: implemented to closely represent the C code above and I'm not in anyway advocating that this is the correct way to write Ada - because it clearly isn't!)

with Ada.Text_IO; use Ada.Text_IO;

pragma Warnings (Off);
procedure API_Fuzzing_Example is

Privileged : Integer := 0;
Password_Buff : String (1 .. 9);
begin

Privileged := 1;
end if;

if Privileged /= 0 then
Put_Line ("Shhhh... the answer to life the universe and everything is 42");
else
Put_Line ("Pah. There is no getting passed this system's hardened security layer...");
end if;

begin
end API_Fuzzing_Example;

Both code examples above contain an array overflow bug. In particular the assignment to variable 'Password_Buffer' will overflow if the command line argument contains a String greater than 9 characters.

For the examples above let's assume that the encompassing applications have system security requirements around the protection of the identified security asset: "the answer to life the universe and everything".

Let's now consider what happens when both of these programs execute with the following command line arguments:

1. "A*^%bd0eK"

This is the correct password. Both programs correctly expose the security asset by displaying: "Shhhh... the answer to life the universe and everything is 42!"

2. "123456789"

This is an incorrect password. Both programs correctly protect the security asset by displaying: "Pah. There is no getting passed this system's hardened security layer..."

This is an incorrect password. The Ada runtime raises a constraint error on line 14 "Password_Buff := Password.all;" stating "length check failed". The exception is unhandled so the program terminates - the security asset is protected.

The C program displays: "Pah. There is no getting passed this system's hardened security layer..." - the security asset is also protected.

4. "A*^%bd0eKKKK"

This is an incorrect password. The Ada runtime raises a constraint error on line 14 "Password_Buff := Password.all;" stating "length check failed". The exception is unhandled so the program terminates - the security asset is protected.

The C program displays: "Shhhh... the answer to life the universe and everything is 42!". The security asset is now exposed and can be exploited by the attacker.

But why?

You will have likely spotted the reason the C program failed to protect the asset which is obviously due to the memcpy overflow writing data into the neighbouring 'int' variable 'Privileged'. This type of stack overflow exploitation is a widely recognised security issue within C programs and various counter measures exist to prevent this type of vulnerability. For example random stack memory allocation can help ensure the overflow leads to a core dump rather than continuing to execute with corrupted state.

The issue is exasperated when you consider that fuzz testing the C program may or may not reveal the bug. The overflow error could eventually cause a segmentation fault however, this is not guaranteed. Fuzz testing the Ada program however would quickly identify the problem and allow the developer to fix the bug with some basic defensive programming and therefore remove the vulnerability altogether.

Fuzzing into the Future

Within the UK AdaCore is actively researching and developing a fuzz testing solution for Ada based software as part of the “High-Integrity, Complex, Large, Software and Electronic Systems” (HICLASS) Aerospace Technology Institute (ATI) supported research and development programme. HICLASS was created to enable the delivery of the most complex, software-intensive, safe and cyber-secure systems in the world and is best described in an earlier AdaCore blog post titled "AdaCore for HICLASS - Enabling the Development of Complex and Secure Aerospace Systems".

Stage one of this work involved the development of an Ada library to abstract away the complexity involved in fuzz testing Ada programs at the unit level. Unit level fuzz testing involves considering the test case files as in-memory representations of identified test data. The test data is any software variables the test engineer would like the fuzzer to drive and this is typically made up of 'in mode' parameters for the function under test, global state, stub return values and even private state manipulated through test injection points or memory overlays. Once the test data has been identified a series of generic packages are instantiated with the Ada types associated with the test data. These packages are responsible for marshalling and unmarshalling the starting corpus and mutated test cases between binary files and Ada variables; we call this collection of test data information an "AFL Map". A simple test harness is then constructed which receives the test case from the fuzzer via a String file path within a command line parameter. The unmarshalled data is then used to setup and execute the test. In addition a top level exception handler is placed in the test harness to capture unexpected exceptions raised by the runtime, these are converted into core dumps to signal a fault to the fuzzer.

Figure 3 shows a high level overview of the architecture of unit level fuzz testing for Ada.

In addition, this work also involved researching and developing a prototype 'Starting Corpus Generator' to create the initial set of test cases that the fuzzer will mutate and a 'Test Case Printer' to convert binary test cases into a human readable syntax. Figure 4 below shows the content of a binary test case created using an AFL Map made up of test case data with the types: Integer, Boolean, Character, Float and String and the human readable version created using the Test Case Printer.

Stage two of this work required the construction of a separate Ada application to automate the building and executing of Ada fuzz tests. In addition this application allows the test engineer to build up a set of stop criteria rules that, when satisfied, will terminate the fuzzing session. The rules are based around the exposed state of AFLplusplus which the fuzzing engine broadcasts every thirty seconds. This tool also features an incorporated coverage capability through GNATcoverage allowing for dynamic full MCDC coverage analysis that can also be included within the stop criteria rules. Finally this tool allows for additional AFLplusplus companion applications to easily be utilised by Ada fuzzing sessions - including starting corpus minimisation and utilising maximum available processor cores to run multiple afl-fuzz instances simultaneously.

Figure 5 below shows this tool in action with an active AFLplusplus fuzzing session.

Stage 3 involved the building of a fuzzer benchmark testing suite used internally to perform comparisons between versions of the capability as new prototype enhancements are developed.

Stage 4 is where we research and develop a prototype structure aware custom mutation engine for Ada fuzzing sessions. This application is being developed as a standalone dynamic library that implements the AFLplusplus custom mutation API. The library can then be registered with an active AFLplusplus session, to unmarshall the received byte array representations of the test case data into their associated Ada types before being sent to type specific custom mutation packages for manipulation.

Stage 5 is looking into random automated test case generation. This capability will form the basis of the starting corpus of test cases by automatically generating wide ranging sets of test case data. In addition it will be utilised by the custom mutation library to mutate the data whilst ensuring the data remains correctly constrained.

Stage 6 involves further automation in the form of automated test harness creation and test data identification.

Final Note...

Fuzz testing is seen by AdaCore as a key tool in the software development process for asset protection within security related and security critical applications and continued effort will be applied to ensure the capability matches the needs of our customers. In addition when combining Fuzzing with other security based tools such as static code analysers like CodePeer and formal verification tools like SPARK Pro, and particularly when the tools come with extensive documentation showing mitigation against the top CWEs, the result can be a cyber-secure software development environment fit for the modern cyber-warfare battle ground.

To conclude: the future may be fuzzy but we're ok with that! ;-)

]]>

Charles Villard, Cyril Etourneau, Thomas Delecroix, Louise Flick worked together in the ADArrose project. It won the student prize in the Make with Ada 2019/20 competition. This project was originally posted on Hackster.io here. For those interested in participating in the 2020/21 competition, registration is now open and project submissions will be accepted until Jan 31st 2021, register here.

Story

This project intends to make an automated sprinkler based on a STM32F429 board. This project will use ADA language and SPARK verification system. We all know the trouble to keep a plant. Especially in Paris, where the light is low and where our lives goes to 100km/h. We don’t have time to water plants. On the eventuality that we manage to find the necessary time, we will most likely give it too much water or not enough and by the time we find out the plant will be dead. Using the secure ADA language and contract verification, we will be able to create an automatic sprinkler to keep our plants alive and have a beautiful green environment without worrying about it.

Features

Our automatic sprinkler comes with numerous features. First of all it can be set in different activity modes to best respond to your needs.

• Continuous mode:
The automatic sprinkler is always active. Once every hour, the soil humidity around the plant is checked with a sensor. If the humidity percentage isn’t high enough, the sprinkler will water the plant with the right amount of water to keep it well hydrated. Perfect for fragile plants.

• Economic mode:
The automatic sprinkler is always active. When the luminosity is low, the soil humidity around the plant is checked once every hour with a sensor. If the humidity percentage isn’t high enough, the sprinkler will water the plant with the right amount of water to keep it well hydrated. This is done to avoid losing water due to evaporation. Perfect to minimize the water consumption. Use on plants without special needs.

• Planned mode:
The automatic sprinkler always active, but can only water the plant during the periods specified by the user. The humidity of the soil is checked at the start of the period and then once every hour. If the humidity percentage isn’t high enough, the sprinkler will water the plant with the right amount of water to keep it well hydrated. A last check is done at the end of the period to ensure the well being of the plant. Useful for a tailored experience.

• Punctual mode:
The automatic sprinkler does one humidity check when starting. If the humidity percentage isn’t high enough, the sprinkler will water the plant with the right amount of water to keep it well hydrated. When the process is finished, the system signals to the user that it can be shut down. For those who want to keep some contact with their green friends. Use to reduce the electricity bill.

Numerous data sets about the plant are collected when the sprinkler is in the first 3 modes. The soil humidity and surrounding luminosity are recorded once every hour and the user can have access to these data over the last 24 hours. Whenever something abnormal is detected in the plant consumption or in the environment, messages will warn him about the peculiar situation. If a problem arises in the system, the user will also be notified.

• Access the project schematics here.
• Access the project code here.
]]>

Laboratorio Gluon's project won a finalist prize in the Make with Ada 2019/20 competition. This project was originally posted on Hackster.io here. For those interested in participating in the 2020/21 competition, registration is now open and project submissions will be accepted until Jan 31st 2021, register here.

Story

Introduction

In this project I will show you the why and the how of my idea for the MakeWithAda contest. I also made a video, which is in spanish but I subtitled it to english so all you can understand, explaining the truck structure and building.

All the code, images, design and videos of this contest are made by myself, except the Console library, which I made, with a friend, for the previous MakeWithAda Contest in this project. Everything else is new creation for this contest.

End Disclaimer;

The problem

With the unstoppable increase in world temperatures mainly due to Global Warming, the news of fires devastating vast areas of the rainforest and populated areas, will be the usual. We can already see them in a daily basis, and will increase in the following years if we do not stop it.

These fires worldwide, are one of the greatest dangers for many species, including the human race. But as the most dangerous specie for the world, we need to implement all we can to stop this devastation.

Also, we are experiencing one notorious development of unmanned systems, from UAVs(Unmanned Aerial Vehicle) to UGVs(Unmanned Ground Vehicle), using this technology we can develop systems that increment the effectiveness of the fire extinguishing methods.

The solution

In this project I would like to introduce you AFT (the Autonomous FireTruck) which is my prototype for a solution that could be implemented in large scale, and for which the technology is already available. This proyect will consists in two main parts:

• The guidance system, which is the one that detect the position of the AFT and analyze the fire propagation and position
• The vehicle that puts out fires, which is in constant communication with the guidance system, to know exactly his own position and the fire's one.

The guidance system can be: from a mesh of antennas and sensors around the mountains, or a UAV flying over the fires, to an space satellite measuring the fire and vehicles. On the other hand, the vehicles can be anaything from an autonomous firetruck, or an autonomous airplane, or even small robots, for smaller fires.

This solution would safe life of many people and animals, directly and indirectly. Directly, by avoiding sending people to dangerous fires, and indirectly with a faster response and tireless systems that can work 24/7, increasing the efficiency of the current fire extinguishing systems.

Prototype

Since the scope of the proyect can be anywhere from a simple simulation to a fully working system, I am developing something in between. That is:

• A small guidance system which will be an image processing software running in a computer, that sends the position of the fire and the vehicle information.
• The vehicle will be a small truck that carries a small water deposit and tower to aim the waterjet.

The main reasons behind this prototype is the technology testing and the availability for a real world implementation. The guidance system (a camera detecting fire and trucks) can be mounted in poles in the forest, in an UAV or in satellites. Also, the technology in Unmanned vehicles' navigation system allow for a safe autonomous guidance that can drive them to the objective marked by the guidance system.

To sum up, this prototype tries to implement a small scale prototype of a guidance system based in image, and a small truck that actually tries to put out fires. This system can be also evolved to a real world vehicles and technologies.

Technologies

In order to achieve our goal of a fully functional prototype we need modern hardware and software libraries. In this section I will cover the main ones I am using in this project.

OpenCV

The first one I will talk about it 'OpenCV' a computer vision and machine learning software library. It has been ported to many languages so one can use it with the one he feels more confortable. In my case I am using openCV with Python for a basic system. As I said earlier, this is not the main part of the project, so the software will be as simple as possible.

STM32f407

As the main brain on the AFT I will use the STM32F4DISCOVERY. I used this board in previous contests, and I could use some of the code that I implemented for it. Furthermore, I will keep expanding the library of drivers for this board, which will make the next projects easier to implement.

For this prototype we will be using the following interfaces of the board:

• SPI: For communication with the NRF24
• Digital I/O: For the water pump activation
• PWM: For the motors and servos.

The embedded software which will be run on the STM32 board will be programmed in Ada. I though this proyect would be a nice choice since Ada is one of the main languages in the development of critical systems, where the life of human beings are at risk. So, since this is a prototype of the software on a "real truck" or "real UAV", the use of Ada is very convenient.

Also, the use of Ada as the programming language helps in the development, detecting the most common errors and problem during developing. This makes that every release of the software is reliable with even low testing. We will be using the Ada Library so most of the interfaces of the STM32's modules are already available. However we will be implementing some library for:

• Servo control: A servo library that allows to initialize and configure the PWM port, adjust the calibration and allow for limited range of movement.
• CarControl: A library that can be used to speak with a L298N (or equivalent H-bridge) and allows to control the speed and direction.
• RF24 library: To initialize de SPI port, configure the pins, and allows for simple RF24 startup and use as a receiver.
nRF24L01

Is a cheap wireless communication module in 2.4Ghz that can be used to send and receive data. It is connected through an SPI port, and while the configuration and initiazation is a bit complex, once it is configured it is easy to use. Therefore, once this library is implemented and included in any project, the nRF24L01 module can be used seamlessly in any other project.

Guidance System

The Guidance system is the one in charge of detect the fire and its position from the vehicle, and send this information to the vehicle. Since this project is aimed for the Make With Ada Contest, I will not focus too much in this section as it is out of the scope of the contest. However, it will be fully functional, yet can be highly improved.

This system consists in two parts:

• The PC software that analyze the image from a camera
• The Arduino connected to the with the communication system attached.

The implementation of the guidance software was done in multiple iterations. I choose openCV + Python as the base, and since it was my first time with openCV I really needed several steps until I get a functional prototype. Basicly the workflow has been:

• Learn how to calibrate the camera with openCV.
• Learn how to detect simple things (in my case, colors): In this step I started detecting a color pattern in the back of the vehicle.
• Calculate the position of the vehicle from the Camera (with solvePnP function).
• Implementation of mask filtering in order to improve segmentation of colors.
• Detect background and filter constant background from the image. I have written a small tutorial, but it is in spanish
• Detect the fire/objective.
• Send all this information to the vehicle.

The Arduino software is a basic interface between the PC and the NRF24, so it basically send the data received by the USB to the NRF24.

Vehicle

The AFT (Autonomous FireTruck) is the main part of the proyect, is the one implemented in Ada, and in where I put most of my effords. This truck is build around a cardboard box where all the components are attached. This system contains:

• A 3S-LiPo battery pack
• Two main motors with wheels
• An L298N module to control the motors.
• The STM32 as the main board of the AFT.
• A deposit with a Water Pump.
• An adaptation board for the Water Pump activation.
• Two servos for a 2 degrees of freedom control of a Hose.
• The Hose who connects the water deposit and the control tower.

Part 1 : Communications

The first part was establishing the communication between the computer and the STM32 board, so I can advance from there. With the communication link working, the next steps would be much easier. I did not found any nRF24L01 library for STM32 written in Ada, so I started working on my own version, following a little bit the interface of the Arduino's version, so it can be easily followed and upgraded. This part was a big headache, making the nRF24L01 start receiving data took me a lot of time. The nRF24L01 has a lot of configurations and registers, which makes the device hard to turn on. However, once it has been initialized it is really easy to use and intregate in any other project.

For the data in the communication I use a record to store the command type and the command data, and then I parse each binary packet received from the nRF24L01 and convert it to the Command. Then in the main loop, I update all the systems according to the new data received.

The Command types I have defined for this project were:

• TEST_LED: For testing purposes, it toggles a LED in the board so I can check that the code it is still running.
• SET_DIRECTION: Configure the CarController to set the direction of movement.
• SET_SPEED: Set the speed for the wheel of the CarController. The CarController has already the direction from the previous command.
• SET_SERVO: Sets the angle of the servos that control the hose.
• SET_PUMP: De-/Activate the water pump.
• SET_MAIN_STATUS: Currently not used, but was created in order to command different states to the truck, i.e.: STOP, MOVING_TO_TARGET, ...
• INFO_TARGET: Is used to receive the data with the position and distance of the objective.

As you would expect the TEST_LED command was the first one implemented, with this command, I started testing the communications and the full message parsing functions. Since in this prototype we are only using one way commands, with no responses, the nRF24L01 library will not implement the send functions.

In this part I learned a lot about the nRF24L01 and all the internal states and configuration, also improved my hability handling big and complex datasheets. Also, since the communication with the nRF24L01 is using the SPI interface, I learned how to use it for the first time.

About the library of the nRF24L01, it is not complete, as there are many configuration options that were out of scope for this project. However, I tried to do it as scalable as possible, with a clean and structured code.

For example to add a new register to the processing one has to:

• Define the new record and its sizes
• Create the FROM/TO unchecked Conversion
• Use the register!

Part 2: Movements

The next natural step is making it moving with the L298N and two motors. In this step I implemented the SET_DIRECTION command, which makes the car move in the desired direction. I implemented the CarController library that allows one to control a L298N board, using the digital pins to set the direction of rotation of the wheels, and a PWM signal generated to set the speed of rotation.

This library is pretty simple, it just set the values for the GPIOs that are connected to the L298N. In this part I learned about the PWM signal generation, which I will use later in the Servo controller.

Part 3: Pumping water!

The programming of the Water pump is pretty straightforward, it just ONE or ZERO!!! , almost everything done in 2-3 lines of code :D. However, the intesting part of this section was the electronic design to activate the water pump which worked at 7V, while the STM32 works at 3.3V. So the first iteration of the code was to do a two-stage circuit: the first part was to rise de 3.3V from the GPIO to 7V but with low power consumption, this 7V would activate a MOSFET to control the actual water pump.

However, this solution was really complex for the problem that was addressing. The water pump just draws 50mA of current at full power, so it can be controlled with just a single BJT transistor, leaving us a pretty simple circuit.

Part 4: Aim and ... shoot!

The last part of the truck is the building of the tower that controls the direction of the water splash. The tower is build around two servos: one for vertical and the second one for horizontal aiming. There are two main problems here, the servo calibration, and the aiming.

For the servo calibration, we have to take into account that, since these are cheap servos, they do not follow the standard protocol exactly, that is why my servo library allows for a simple calibration. This calibration methods allow to set the value of the PWM for 0, 90 and 180 degrees.

So, once the PWM signal was configured for the 20ms period, I did a testbench in order to get the values of microseconds that makes the servos go from 0 to 180 degrees.

Then with these values, the servo library has two methods to move them: one with degrees as input, and another one with microseconds. If they are calibrated, the first one is recomended. Also, limits for the movement of the servos can be configured, therefore if the servo is commanded to move further the limit, the servo will stop in the limit.

Until now, I can aim the tower, however the aiming was not related to the information received by the "guidance system", in order to fix the aiming. First I have to do some experiments to calculate the relation between the angle of the vertical servo, and the distance to where the water hits the ground.

The results were as follow:

With this values, we can now interpolate, with a simple linear interpolation, the angle of the vertical servo. So finally, we can calculate the orientation of the vertical and horizontal servos, based on the distance and angle to the target!

All together

We have seen each part separately, but the fun part comes when they get connected together. The next image represent how all the items are conected, the voltages of each one, and all the interfaces needed.

The Embedded code

I tried to do the code as self-explanatory as possible, but the big picture would be:

• First we call all the initialization for each component. Each one has the Init/Config function in order to configure the GPIOs and internal registers (timers, ...)
• Then in the main loop I just get all the new commands from the nRF24L01 wifi device and update the commands. Then, only new commands are executed, each one in its component. For example, if a "SET_SERVO" arrives, in that loop the software will set the servo position and mark the command as "old", so it is not processed again. However, every command is stored until the next one arrives.

In the following picture we can see the full software flow. On the left column we have the entry point of the software, while the other 4 columns are the components of the project and what they do in each step of the main program.

The OpenCV Code

The folder 'openCV' of the project contains the code for the python and openCV. There we have 5 files:

• primerIntento.py it was a playground to learn how to start detecting things with openCV, so this is not mandatory for the project, and hence it is not clean nor optimized
• Calibration.py: is the file that handles the calibration process of the Camera, it only has to be run once.
• ApagaFuegos.py is the main file of the software. This piece of software initialize the camera, the detection algorithm and execute it in the main loop
• Comms.py contains all the code to initialize the serial, compose the data packets and send it to the Arduino with the nRF24
• Detector.py Has all the logic of the detection algorithm. It can work in three modes: Simple (which is the simplest one and now it try to go to the target using only 2 elements), the tester (which is used to get the HSV filter in realtime) and the normal mode (which is the one used in the AFT)

How to use it

In order to replicate this project the following steps shall be followed:

• Assemble the hardware following the previous notes
• Flash the Embedded AdaCore source in the stm32F407
• Flash the Arduino Code to send data over the WiFi
• Adapt the openCV Comms script to open the correct Serial port.
• Setup the webcam and connect to the PC.
• Launch the "ApagaFuegos" in the openCV folder
• Once the scenario is clear, press "b" to take a picture of the background
• Add the truck and the fire to the scenario
• Press "m" to activate the movement in the Truck
• Wait and enjoy! :D

Conclusion and future

As we can see in the video, the truck work as expected, and actually moves to the target and aim to it. However, as this is a prototype is has a strong contrains in the enviroment, due to the openCV algorithm, which is where this project is highly improbable.

• Access the project schematics here.
• Access the project code here.
]]>

Guillermo Perez's project won a finalist prize in the Make with Ada 2019/20 competition. This project was originally posted on Hackster.io here. For those interested in participating in the 2020/21 competition, registration is now open and project submissions will be accepted until Jan 31st 2021, register here.

With this smart robot, we will use a neural network to follow the light and stop when it reaches its goal.

Story

INTRODUCTION

If the robots had to have everything pre-programmed it would be a problem. They simply could not adapt to changes, a "programmer" would have to change the code for the robot to work every time the environment changes. The way to solve it's by allowing the robot to learn and artificial intelligence is a necessary alternative if we want the robots not only to do what we need them to do, but to help us find even a better solution. The perceptron is an algorithm based on the functioning of a neuron. It was first developed in 1957 by Frank Rosenblatt, and here you can check the whole story: https://en.wikipedia.org/wiki/Perceptron

The main goal of this project is to develop a robot car that follows the light emitted by a lamp and through the use of neural networks.

Particular goals:

• We will use three inputs (including the BIAS), four hidden neurons and four outputs (each one connected to the motors pins). It's important to clarify that some programmers prefer that the robot simultaneously calculate the weights of the neurons and execute the action, and this is the cause a small delay of time. We prefer to calculate the weights using software to avoid these errors.
• And also we include an input to stop the robot when it reaches its target from a threshold of light, since many robots continues the move without stopping.

Next, you will find all the information in the following order: Hardware, Neural Network, GPS Project, Assembling the Chassis, Test and Conclusion.

HARDWARE

(Timing: 2 hrs.)

The electrical diagram of this project I show you in the figure below.

Note: All parts are commercial and easy to get. I recommend that when you assemble the motors do it carefully, since you can connect them in the opposite direction. So make the connections, and then do the tests.

STM32F429I-DISC1

This is the board that I will use for my project, which has many useful tools and a great AdaCore library support.

• STM32F429ZIT6 microcontroller featuring 2 Mbytes of Flash memory, 256 Kbytes of RAM in an LQFP144 package
• USB functions: Debug port, Virtual COM port, and Mass storage.
• Board power supply: through the USB bus or from an external 3 V or 5 V supply voltage
• 2.4" QVGA TFT LCD
• 64-Mbit SDRAM
• L3GD20, ST-MEMS motion sensor 3-axis digital output gyroscope
• Six LEDs: LD1 (red/green) for USB communication, LD2 (red) for 3.3 V power-on, two user LEDs: LD3 (green), LD4 (red) and two USB OTG LEDs: LD5 (green) VBUS and LD6 (red) OC (over-current)
• Two push-buttons (user and reset)

Tools, software, and resources on: https://www.st.com/en/evaluation-tools/32f429idiscovery.html

LDR (Light Dependent Resistor)

Two cadmium sulphide(cds) photoconductive cells with spectral responses similar to that of the human eye. The cell resistance falls with increasing light intensity. Applications include smoke detection, automatic lighting control, batch counting and burglar alarm systems.

NEURAL NETWORK

(Timing: 8 hrs.)

A neural network is a circuit of neurons, or an artificial neural network, composed of artificial neurons or nodes. Thus a neural network is either a biological neural network, made up of real biological neurons, or an artificial neural network, for solving artificial intelligence problems. The connections of the biological neuron are modeled as weights. A positive weight reflects an excitatory connection, while negative values mean inhibitory connections. All inputs are modified by a weight and summed. This activity is referred as a linear combination. Finally, an activation function controls the amplitude of the output. For example, an acceptable range of output is usually between 0 and 1, or it could be −1 and 1. https://en.wikipedia.org/wiki/Neural_network

Bibliographic reference:

The theoretical information you can find on several websites, in my case I support my work with the following publication in Spanish: https://www.aprendemachinelearning.com/crear-una-red-neuronal-en-python-desde-cero/

Here the author shows us the theory, and an example to calculate the weights of a neural network with Python and applied to Arduino board. In my case I will make a brief review of how I use this information to do my project with AdaCore. For example I've modified the codes in Python programming language according to my needs, and I'm the author of the GPS codes.

Lets starts

In our robot, what we want to achieve is that it moves, using a light sensor, in a directed way towards the light. The robot will only have two types of movement: move forward, and rotate to the left. We want the car to learn to move forward when its movement brings it closer to the light, and to turn for a time of 100 milliseconds when it moves away from the light. The reinforcement will be as follows:

• If the movement improves the light, with respect to its previous position, then it rewards: move forward
• And if the movement results in a less light position, then the movement is penalized for not doing so again: turn to the left

I show you the neural network that I designed with three inputs in the figure below, four neurons and four outputs. We can appreciate all the connections that each one represents the hidden weights that we will calculate later using software.

We want our robot to learn to follow the light with only one light sensor. We are going to build a system that takes two inputs: 1) the received light (a value between 0 to 4095) before the movement, and 2) the received light after the movement. In the figure below we can see the table of values of how our neural network would work.

We can appreciate that there are only four options:

• If the current light is greater than or equal to the previous light and we have a light emission of less than 3750, the robot moves forward for 100 milliseconds.
• If the current light is less than the previous light and we have a light emission less than or equal to 3750, the robot turns to the left for 100 milliseconds.
• In the other two remaining options the robot stops because the light emitted by the lamp is greater than 3750 bits. This means that it's very close to the light.

The Complete code of the "Neural network" with Backpropagation is as follows:

import numpy as np

# We create the class
class NeuralNetwork:

def __init__(self, layers, activation='tanh'):
if activation == 'sigmoid':
self.activation = sigmoid
elif activation == 'tanh':
self.activation = tanh

# Initialize the weights
self.weights = []
self.deltas = []
# Assign random values to input layer and hidden layer
for i in range(1, len(layers) - 1):
r = 2*np.random.random((layers[i-1] + 1, layers[i] + 1)) -1
self.weights.append(r)
# Assigned random to output layer
r = 2*np.random.random( (layers[i] + 1, layers[i+1])) - 1
self.weights.append(r)

def fit(self, X, y, learning_rate=0.2, epochs=100000):
# I add column of ones to the X inputs. With this we add the Bias unit to the input layer
ones = np.atleast_2d(np.ones(X.shape[0]))
X = np.concatenate((ones.T, X), axis=1)

for k in range(epochs):
i = np.random.randint(X.shape[0])
a = [X[i]]

for l in range(len(self.weights)):
dot_value = np.dot(a[l], self.weights[l])
activation = self.activation(dot_value)
a.append(activation)
#Calculate the difference in the output layer and the value obtained
error = y[i] - a[-1]
deltas = [error * self.activation_prime(a[-1])]

# We start in the second layer until the last one (A layer before the output one)
for l in range(len(a) - 2, 0, -1):
deltas.append(deltas[-1].dot(self.weights[l].T)*self.activation_prime(a[l]))
self.deltas.append(deltas)

# Reverse
deltas.reverse()

# Backpropagation
# 1. Multiply the output delta with the input activations to obtain the weight gradient.
# 2. Updated the weight by subtracting a percentage of the gradient
for i in range(len(self.weights)):
layer = np.atleast_2d(a[i])
delta = np.atleast_2d(deltas[i])
self.weights[i] += learning_rate * layer.T.dot(delta)

if k % 10000 == 0: print('epochs:', k)

def predict(self, x):
ones = np.atleast_2d(np.ones(x.shape[0]))
a = np.concatenate((np.ones(1).T, np.array(x)), axis=0)
for l in range(0, len(self.weights)):
a = self.activation(np.dot(a, self.weights[l]))
return a

def print_weights(self):
print("LIST OF CONNECTION WEIGHTS")
for i in range(len(self.weights)):
print(self.weights[i])

def get_weights(self):
return self.weights

def get_deltas(self):
return self.deltas

# When creating the network, we can choose between using the sigmoid or tanh function
def sigmoid(x):
return 1.0/(1.0 + np.exp(-x))

return sigmoid(x)*(1.0-sigmoid(x))

def tanh(x):
return np.tanh(x)

return 1.0 - x**2

########## CAR NETWORK

nn = NeuralNetwork([2,2,4],activation ='tanh') # no incluir la bias aqui porque si la esta en los calculos

X = np.array([[1,1],   # light_Current >= light_Before & light_Current <= 3750
[-1,1],   # light_Current < light_Before & light_Current <= 3750
[1,-1],   # light_Current >= light_Before & light_Current > 3750
[-1,-1],   # light_Current < light_Before & light_Current > 3750
])
# the outputs correspond to starting (or not) the motors
y = np.array([[1,0,1,0], # go forward
[0,1,1,0], # turn to the left
[0,0,0,0], # stop
[0,0,0,0], # stop
])
nn.fit(X, y, learning_rate=0.03,epochs=15001)

def valNN(x):
return (int)(abs(round(x)))

index=0
for e in X:
prediccion = nn.predict(e)
print("X:",e,"expected:",y[index],"obtained:", valNN(prediccion[0]),valNN(prediccion[1]),valNN(prediccion[2]),valNN(prediccion[3]))
index=index+1

You can run this code with Python 3.7.3, or with Jupiter Notebook, and the result is as follows:

With 15 thousand periods it was enough to get minimal errors. We have to add the following code to see the cost graphic:

########## WE GRAPH THE COST FUNCTION

import matplotlib.pyplot as plt

deltas = nn.get_deltas()
valores=[]
index=0
for arreglo in deltas:
valores.append(arreglo[1][0] + arreglo[1][1])
index=index+1

plt.plot(range(len(valores)), valores, color='b')
plt.ylim([0, 1])
plt.ylabel('Cost')
plt.xlabel('Epochs')
plt.tight_layout()
plt.show()

Finally we can see the hidden weights obtained and the output weights of the connections, as these values will be the ones we will use in the final network on our AdaCore code:

########## WE GENERATE THE GPS CODE
def to_str(name, W):
s = str(W.tolist()).replace('[', '{').replace(']', '}')
return 'float '+name+'['+str(W.shape[0])+']['+str(W.shape[1])+'] = ' + s + ';'

# We get the weights trained to be able to use them in the GPS code
pesos = nn.get_weights();

print('// Replace these lines in your arduino code:')
print('// float HiddenWeights ...')
print('// float OutputWeights ...')
print('// With trained weights.')
print('\n')
print(to_str('HiddenWeights', pesos[0]))
print(to_str('OutputWeights', pesos[1]))

I repeat it, I modified these codes and ran them with Python and you can get in the download section. Now we are going to implement these coefficients in our AdaCore code.

GPS PROJECT

(Timing: 8 hrs.)

GNAT Programming Studio (GPS) is a free multi-language integrated development environment (IDE) by AdaCore. GPS uses compilers from the GNU Compiler Collection, taking its name from GNAT, the GNU compiler for the Ada programming language. Released under the GNU General Public License, GPS is free software. The download link is as follows: https://www.adacore.com/download

In my opinion, this is a powerful software with mathematical tools that allow us to develop high-precision projects as I explain below. First step, we have to download the next library: https://github.com/AdaCore/Ada_Drivers_Library

I'm the author of the following code, and to develop this GPS project, I used the following two examples: demo_gpio_direct_leds, and demo_adc_polling.

How does it work?

To start my project I have used the sample code: demo_adc_gpio_polling.adb; in my project I'm going to measure the analog signal generated by the LDR sensor through the analog port PA5, we will have 4095 values with its 12-bit CAD.

Converter     : Analog_To_Digital_Converter renames ADC_1;
Input_Channel : constant Analog_Input_Channel := 5;
Input         : constant GPIO_Point := PA5;

By default, this code has already configured and initialized the digital output ports of the green and red LEDS, ie PG13 and PG14. I have only added ports PD12 and PD13.

LED1       : GPIO_Point renames PG13; -- GREEN LED
LED2       : GPIO_Point renames PG14; -- RED LED
LED3       : GPIO_Point renames PD12;
LED4       : GPIO_Point renames PD13;

After the digital and analog ports have been configured, in the declaration of variables, I've loaded the weights of the neural network calculated in the previous section:

-------------------------------------------
--- NEURONAL NETWORKS' WEIGHTS
-------------------------------------------
HiddenWeights_1_1 := -0.8330212122953609;
HiddenWeights_1_2 := 1.0912649297996564;
HiddenWeights_1_3 := -0.6179969826549335;
HiddenWeights_1_4 := -1.0762413280914194;
HiddenWeights_2_1 := -0.7221015071612642;
HiddenWeights_2_2 := -0.3040531641938827;
HiddenWeights_2_3 := 1.424273154914373;
HiddenWeights_2_4 := 0.5226012446435597;
HiddenWeights_3_1 := -1.3873042452980089;
HiddenWeights_3_2 := 0.8796185107005765;
HiddenWeights_3_3 := 0.6115239126364166;
HiddenWeights_3_4 := -0.6941384010920131;
OutputWeights_1_1 := 0.4890791000431967;
OutputWeights_1_2 := -1.2072393706400335;
OutputWeights_1_3 := -1.1170823069750404;
OutputWeights_1_4 := 0.08254392928517773;
OutputWeights_2_1 := 1.2585395954445326;
OutputWeights_2_2 := 0.7259701403757809;
OutputWeights_2_3 := 0.05232196665284013;
OutputWeights_2_4 := 0.5379573853597585;
OutputWeights_3_1 := 1.3028834913318572;
OutputWeights_3_2 := -1.3073304956402805;
OutputWeights_3_3 := 0.1681659447995217;
OutputWeights_3_4 := -0.016766185238717802;
OutputWeights_4_1 := -0.38086087439361543;
OutputWeights_4_2 := 0.8415209522385925;
OutputWeights_4_3 := -1.527573567687556;
OutputWeights_4_4 := 0.476559350936026;

Once the program starts working, we calculate and load the input values, using the variables 'eval' and 'lux':

----------------------------------------------
-- CHARGE INPUT VALUES...
----------------------------------------------
if light_Current >= light_Before then -- If light_Current is greater than or equal to, move forward
eval := 1.0;
else -- else, it turns counterclockwise
eval := -1.0;
end if;
if light_Current <= 3750 then -- If light_Current is less than, move forward or turn to the left
lux := 1.0;
else -- else, "Stop"
lux := -1.0;
end if;
---------

I experimentally calibrated the lux value of 3750 with several models of LDR sensors... with this value, I found a satisfactory response to position the robot close to the light source and without impacting the lamp or losing it in its movement.

We multiply the input matrix by the matrix of the hidden weights, and to each accumulated result we calculate the hyperbolic tangent to obtain values between 1 and -1:

----------------------------------------------
-- Input * HiddenWeights
-- We use the Tanh to get values between 1 and -1
----------------------------------------------
Hidden_Layer_1_1 := Tanh(1.0*HiddenWeights_1_1 + eval*HiddenWeights_2_1 + lux*HiddenWeights_3_1);
Hidden_Layer_1_2 := Tanh(1.0*HiddenWeights_1_2 + eval*HiddenWeights_2_2 + lux*HiddenWeights_3_2);
Hidden_Layer_1_3 := Tanh(1.0*HiddenWeights_1_3 + eval*HiddenWeights_2_3 + lux*HiddenWeights_3_3);
Hidden_Layer_1_4 := Tanh(1.0*HiddenWeights_1_4 + eval*HiddenWeights_2_4 + lux*HiddenWeights_3_4);

Again we multiply the matrix of the hidden weights by the matrix of the output weights and to each accumulated result we also calculate the hyperbolic tangent.

----------------------------------------------
-- Hidden_Layers * OutputWeights
-- We use the Tanh to get values between 1 and -1
----------------------------------------------
Output_Layer_1_1 := Tanh(Hidden_Layer_1_1*OutputWeights_1_1 + Hidden_Layer_1_2*OutputWeights_2_1 + Hidden_Layer_1_3*OutputWeights_3_1 + Hidden_Layer_1_4*OutputWeights_4_1);
Output_Layer_1_2 := Tanh(Hidden_Layer_1_1*OutputWeights_1_2 + Hidden_Layer_1_2*OutputWeights_2_2 + Hidden_Layer_1_3*OutputWeights_3_2 + Hidden_Layer_1_4*OutputWeights_4_2);
Output_Layer_1_3 := Tanh(Hidden_Layer_1_1*OutputWeights_1_3 + Hidden_Layer_1_2*OutputWeights_2_3 + Hidden_Layer_1_3*OutputWeights_3_3 + Hidden_Layer_1_4*OutputWeights_4_3);
Output_Layer_1_4 := Tanh(Hidden_Layer_1_1*OutputWeights_1_4 + Hidden_Layer_1_2*OutputWeights_2_4 + Hidden_Layer_1_3*OutputWeights_3_4 + Hidden_Layer_1_4*OutputWeights_4_4);

For each value of the output matrix we calculate the absolute value to handle positive values and round the result:

----------------------------------------------
-- We charge absolute and integer values at the outputs
----------------------------------------------
Output_1 := Integer (abs (Output_Layer_1_1));
Output_2 := Integer (abs (Output_Layer_1_2));
Output_3 := Integer (abs (Output_Layer_1_3));
Output_4 := Integer (abs (Output_Layer_1_4));

Finally we use these output values to activate and deactivate the digital output pins, which in turn feed the L298N driver (IN1, IN2, IN3 and IN4). This process is repeated every 100 milliseconds.

----------------------------------------------
-- Activate the outputs according to the calculations of the neural network
----------------------------------------------
if Output_1 = 1 then
LED1.Set;
else
LED1.Clear;
end if;
if Output_2 = 1 then
LED2.Set;
else
LED2.Clear;
end if;
if Output_3 = 1 then
LED3.Set;
else
LED3.Clear;
end if;
if Output_4 = 1 then
LED4.Set;
else
LED4.Clear;
end if;

We can see the output status of these values on our LCD screen:

----------------------------------------------
-- Print the outputs of the neural network
----------------------------------------------
Print (0, 25, Output_1'Img);
Print (0, 50, Output_2'Img);
Print (0, 75, Output_3'Img);
Print (0, 100, Output_4'Img);

You must successfully compile and flash the code on your STM32F429I board. You can get the complete code in the download section.

Note: The code works well on the STM32F429I board, when it's powered by the PC. However, when the board is powered independently, the code doesn't run and the screen turns white. To correct this error you must update the firmware of your board... How? Simply download and update the firmware with the "STM32 ST-LINK Utility" application that you can download at: https://www.st.com/en/development-tools/stsw-link004.html

ASSEMBLING THE CHASSIS

(Timing: 4 hrs.)

The two pieces that I printed with a 3D printer are the following: the Chassis is where we are going to adjust the two gear-motors, the L298N driver, the battery, and the LDR sensor module.

In figures below we can see several views of the assembly of our autonomous robot.

TEST

(Timing: 2 hrs.)

Below you can see a demonstration of how this prototype works.

CONCLUSION

In this project I learned and confirmed some theoretical concepts, and I'm contributing with an application of neural networks to the Adacore contest, and with an autonomous robot that showed us its effectiveness in developing its task. I didn't have to program all the situations or combinations that the robot would possibly face, and only 3 bits of information in the inputs were enough. This leads me to deduce that it's possible to add more bits of information to develop more complex tasks. The robot did find its target and also stopped when was necessary. I also demonstrated this smart device can be guided with a moving light, and that GPS libraries can make precision mathematical calculations with real-life problems. I recommend you do this kind of thing very carefully and verify every step.

• Access the project schematics here.
• Access the project code here.
]]>

Welcome to the Ada for micro:bit series where we look at simple examples to learn how to program the BBC micro:bit with Ada.

In this eighth part we will see how to play some simple music with the micro:bit and a piezo buzzer.

In the analog output example we said that the analog output was actually a Pulse Width Modulation (PWM) signal. When a piezo buzzer is connected to a PWM a note is produced, changing the frequency of the PWM signal means changing the note that the buzzer will play.

Wiring Diagram

For this example we will need a couple of extra parts:

• A piezo buzzer
• A couple of wires to connect them

Interface

The package MicroBit.Music provides different ways to play music. We are going to use the procedure that plays a melody. A Melody is an array of Note, each Note is made of a pitch and a duration in milliseconds. This package also provides declaration of pitches for the chromatic scale (e.g C4, GS5).

So we declare our little melody like so:

My_Little_Melody : constant MicroBit.Music.Melody :=
((C4,   400),
(G3,   800),
(B3,   400),
(Rest, 400),
(A3,   400),
(G3,   400));

Then we simply have to call the procedure Play of the package MicroBit.Music.

procedure Play (Pin : Pin_Id; M : Melody)
with Pre => Supports (Pin, Analog);

Arguments:

• Pin : The id of the pin that the melody will play on
• M : The melody that we want to play

Precondition:

• The procedure Play has a precondition that the pin must support analog IO.

Here is the full code of the example:

with MicroBit.Music; use MicroBit.Music;

procedure Main is

My_Little_Melody : constant MicroBit.Music.Melody :=
((C4,   400),
(G3,   800),
(B3,   400),
(Rest, 400),
(A3,   400),
(G3,   400));
begin

--  Loop forever
loop

--  Play the little melody on pin 0
MicroBit.Music.Play (0, My_Little_Melody);
end loop;
end Main;

Following the instructions of Part 1 you can open this example (Ada_Drivers_Library-master\examples\MicroBit\music\music.gpr), compile and program it on your micro:bit.

Don't miss out on the opportunity to use Ada in action by taking part in the fifth annual Make with Ada competition! We're calling on developers across the globe to build cool embedded applications using the Ada and SPARK programming languages and are offering over $9,000 in total prizes. Find out more and register today! ]]> First beta release of Alire, the package manager for Ada/SPARK https://blog.adacore.com/first-beta-release-of-alire-the-package-manager-for-ada-spark Fri, 30 Oct 2020 00:00:00 -0400 Fabien Chouteau https://blog.adacore.com/first-beta-release-of-alire-the-package-manager-for-ada-spark Ada 202x support in GNAT https://blog.adacore.com/ada-202x-support-in-gnat Thu, 29 Oct 2020 04:37:00 -0400 Arnaud Charlet https://blog.adacore.com/ada-202x-support-in-gnat News from the Ada front The next revision of the Ada standard is now almost ready, so it's time for a status update on where GNAT and AdaCore stand on this front! This new standard, called Ada 202x for now, is currently getting the final touches at the ARG (Ada Rapporteur Group) before official standardization by the relevant ISO bodies (WG9, SC22 and JTC1). If you want to know more about these groups, you can visit this page. In all likelihood, Ada 202x will become the new official version of Ada by the end of 2021 or early 2022, so may become Ada 2022. In any event, we'll call it Ada 202x here, and GNAT Pro 21 will provide support for many of the new features under the -gnat2020 and -gnatX switches as detailed below. The 21.0 preview has just be released to our customers and the official 21.1 release will be available in February 2021. Ada 202x contains many useful features that complement nicely the current Ada standard, in particular those related to expressiveness of the language and with a focus on programming by contract, introduced with Ada 2012. We'll detail some of these in this blog post. Assessing Ada 202x and making some tough choices In the past year or so, we have been working hard assessing and implementing most of these Ada 202x changes (called AIs: Ada Issues in ARG terms). The implementation work and feedback from first users allowed us to identify that a few of these features would need additional time and attention. This led us to make a difficult decision - in order to allow for more investigation and to avoid users to start to rely on constructs that may need to change or be replaced, we decided to put on hold the implementation of some of the changes in language. Of course, we’re currently engaged with the ARG to discuss these. The main set of features that AdaCore and GNAT are putting on hold are related to the support for parallel constructs. While the overall vision is an exciting and promising one, we realized when looking at the state of the art and gathering user requirements that there were a lot more aspects to consider on top of those currently addressed by the AIs. Some of these are related to GPGPU (General Purpose GPU) support as well as their future CPU counterparts, and include topics such as control of memory transfer, precise allocation of tasks and memory on the hardware layout, target-aware fine tuning options as well as various other parametrization needs. These capabilities happen to be fundamental to obtain actual performance benefits from parallel programming, and providing them may require profound changes in the language interface. Consequently, we’re putting all parallel AIs on hold, including support for the Global and Nonblocking aspects beyond the current support in SPARK. Note also as a reminder that GNAT Pro already takes full advantage of multicore environments on all its supported targets using Ada tasking, including on bare metal platforms via its Ravenscar and now Jorvik (see below) runtimes. Ada 202x features already supported in GNAT Pro 21 So back to the Ada 202x support offered in GNAT Pro 21... We have already implemented over 200 AIs, including the following new features: Jorvik profile Jorvik is a subset of the Ada tasking capabilities, similar to Ravenscar and which imposes fewer restrictions, removing the following ones compared to the Ravenscar profile: • No_Implicit_Heap_Allocations • No_Relative_Delay • Simple_Barriers • Max_Entry_Queue_Length => 1 • Max_Protected_Entries => 1 • No_Dependence => Ada.Calendar • No_Dependence => Ada.Synchronous_Barriers The configuration pragma pragma Profile now supports Jorvik as a possible value to enforce the restrictions and is available as part of the ravenscar-full runtimes on bare metal platforms. Improvements to the 'Image attribute A number of improvements have been done in the way the ‘Image attribute works. In particular, this attribute can be used directly on objects and now applies to any type, and not just scalar types. A new attribute and aspect Put_Image has been introduced, allowing a custom implementation for any type as a replacement of the default supplied one. The exact form of the user supplied Put_Image procedure is still under finalization at the ARG and is provided in an intermediate form in GNAT Pro 21 which will likely change in release 22. Atomic Operations Four new packages, System.Atomic_Operations.Exchange, System.Atomic_Operations.Test_And_Set, System.Atomic_Operations.Integer_Arithmetic and System.Atomic_Operations.Modular_Arithmetic now support accessing processor-specific atomic operations, allowing users to write thread-safe concurrent code without the use of system locks. Support for volatile and atomic objects are also further refined via the Full_Access_Only aspect to ensure that objects are always read and written entirely. Support for infinite precision numbers Two new packages Ada.Numerics.Big_Numbers.Big_Integers and Ada.Numerics.Big_Numbers.Big_Reals provide support for unbounded integer and real numbers with arithmetic operations implemented in software. User-Defined Literals Literals for user types (Integer, Real, String) can also be specified and are supported by default for the infinite precision number types. Variadic C function import Importing variadic functions was not portable and not easily done in practice without resorting to writing C wrappers. It is now supported via a new convention, C_Variadic_N, where N is the number of fixed parameters in the C profile: procedure printf (format : String; optional_param1 : int) with Import, Convention => C_Variadic_1; printf ("value is %d" & LF & NUL, 20); Improved expression and contract expressiveness Declare expressions Ada 202x now allows declaring constants and renamings inside a declare expression, which facilitates writing more complex preconditions and postconditions: Val : Integer := (declare X : constant Integer := Random_Value; begin X + X); Delta aggregates This Ada feature replaces the SPARK 'Update attribute and allows modifying partially the copy of an object: Record_Object := (Record_Object with delta Field1 => X, Field2 => Y); Contracts on Access-to-Subprogram Aspects Pre and Post can now be specified on access to subprogram types. As a consequence, when a call is made from such a type, the contract of the type will be executed, together with the specific contracts of the called subprogram if any. Static expression functions Ada 202x defines a new aspect Static that can be specified on expression functions. Such an expression function can be called in contexts requiring static expressions when the actual parameters are all static, allowing for greater abstraction in complex static expressions. For example: function Half_Size (S : Integer) return Integer is (S / 2) with Static; type T is range 0 .. 10 with Size => Half_Size (Integer'Size); Iterator Filters Iterators can now be provided with an optional filter. This can be used in loops or in the new container aggregates and reduce expressions. For example: for E of Some_Array when E /= 0 loop Put_Line (E'Image); end loop; Indexes in array aggregate Array aggregates now support a for-loop like syntax: (for Index in 1 .. Count => Function_Returning_Limited (Index)) Assignment target name @ Ada 202x provides a convenient shortcut to refer to the left hand side of an assignment statement, as in: Some_Very_Long.And_Complex (Expression) := @ + 1; Another_Very_Long.And_Complex (Expression) := Function_Call (@); Renames with type inference The type information in a renames clause now becomes optional, as in: X renames Y; This means also that named numbers can now be renamed as well: PI : constant := 3.1415926; PI2 renames PI; 'Reduce A new attribute 'Reduce is available for experimentation under the -gnatX switch. It supports performing a map/reduce operation over the values of a container. For example: X : Integer := (1, 2, 3)'Reduce ("+", 0); Will add 1, 2 and 3, and store the result (6) in X. Container aggregates You can now initialize a container object via the aggregate notation, e.g: V : Vector := (1, 2, 3); Next Steps In GNAT Pro 22, we will complete the implementation of all the relevant AIs, and at the same time have started a language prototyping and experimentation effort to prepare future Ada (and SPARK) revisions, including many exciting or most wanted features such as a simplified model for accessibility checks and anonymous access types, generalized case statements on any type (aka pattern matching), simplified and universal storage pools, more static guarantees (e.g. on object initialization), improved string processing in the standard library, simplified finalization support, implicit generic instantiations, ... If you are interested, you can follow and give your input and ideas on this effort via the Ada & SPARK RFC open platform. ]]> Make With Ada 2020: High Integrity Sumobot https://blog.adacore.com/make-with-ada-2020-high-integrity-sumobot Thu, 29 Oct 2020 00:00:00 -0400 Juliana Silva https://blog.adacore.com/make-with-ada-2020-high-integrity-sumobot Blaine Osepchuk's project won a finalist prize in the Make with Ada 2019/20 competition. This project was originally posted on Hackster.io here. For those interested in participating in the 2020/21 competition, registration is now open and project submissions will be accepted until Jan 31st 2021, register here. This document contains all the instructions you’ll need to build your own High Integrity Sumobot. This fully functional mini-sumobot is an advanced-level project programmed in Ada/SPARK and Arduino (C++). Story I created the High Integrity Sumobot using Ada/SPARK and high integrity software engineering techniques. I wanted to make it easy for people interested in Ada/SPARK to see how all the pieces fit together in a small embedded project. I did my best to write simple, clean, and maintainable code throughout. It’s extensively commented and it’s open source so you can use it almost any way you want. This document contains all the instructions you’ll need to build your own High Integrity Sumobot. System overview I started with a Zumo 32U4 sumobot. But instead of programming it directly as a microcontroller like most people do, I turned it into an I2C slave device, which I control with a microbit I programmed in Ada/SPARK. The zumo continuously collects sensor data and stores it in a buffer. The microbit periodically requests sensor data from the zumo over I2C and the zumo sends the data it most recently collected. The microbit validates the sensor data, decides how the zumo should move, and sends motor commands back to the zumo (also over I2C). The zumo validates the motor commands and then executes them. And then the process repeats (at least 50 times per second). The main job of a sumobot is to fight in sumo competitions. So the bulk of the code on the microbit is related to the fighting algorithm. Here’s a simplified state table for it: The 5 x 5 display on the microbit shows a character representing the current top-level state of the system (see the “Display” column of the state table above): You can see the most important zumo parameters on the zumo’s LCD, which is helpful for development and debugging. It shows the following information: My plan for achieving high integrity I: • followed a relaxed version of the personal software process (PSP). I did most the processes prescribed by PSP but because I am new to Ada/SPARK it didn’t make sense to do the tracking and analysis steps • carefully implemented the smallest amount of functionality required to create the High Integrity Sumobot • chose to do as much processing as possible on the microbit • programmed the microbit in SPARK as much as possible, and fell back to Ada only where necessary • programmed the microbit using simple constructs, constrained types, contracts, extensive data validation, formal proofs, and unit tests. I even proved functional correctness for some simple subprograms • programmatically generated test cases for the fighting algorithm to help ensure I tested all the important paths through the code Parts 1 x microbit microcontroller1 x SparkFun micro:bit Breakout (with Headers)1 x Zumo 32u4 robot assembled with 75:1 motors. Note: the older, cheaper version of this robot (found here) will not work with this project1 x Breadboard – Self-Adhesive (White)1 x SparkFun Logic Level Converter – Bi-Directional (needed to convert the I2C signals back and forth from 3.3 to 5 volts because the microbit operates at 3.3 volts and the zumo operates at 5 volts)4 x AA NiMH batteries1 x Break Away Headers – Straight1 x Female Headers2 x Break Away Female Headers – Swiss Machine Pin (used for mounting the breadboard)2 x USB C cables1 x sumo ring (which you can buy or build yourself)misc parts: scavenged bits of wire, hot glue, solder, electrical tape, etc. Getting up and running I tested all of the steps below on Windows 7 and Windows 10 and the 2019 versions of GNAT Community and ARM ELF. Other platforms and versions may work but I have not tested them. The ring and training dummy • you’ll need some kind of ring to train your High Integrity Sumobot in. The specs for mini sumo rings can be found here. You can purchase a ring online but the shipping cost is probably going to be prohibitive. I didn’t have anything as big as the spec requires so I used a old piece of furniture and made the border with masking tape • you’ll also need a training dummy, unless you have another robot to compete against. I made mine out of Lego but you could just as easily build one out of cardboard and hot glue Arduino development environment and Zumo 32U4 setup • install batteries in the zumo and turn it on. It comes pre-loaded with a sketch that demonstrates some of the features of the bot. Ensure your zumo works • download and install the Arduino IDE. Note: I chose the windows installer, not the windows app (I’m not saying there’s anything wrong with the app, I just have no experience with it) • start the Arduino IDE. You may be prompted to update your boards and libraries. Go ahead and do that • install the windows drivers for the zumo • configure the Arduino IDE to program the zumo 32U4. Make sure you can upload the blink demo to your zumo and confirm that it works • in tools -> manage libraries install the following Arduino libraries: Zumo32U4 and ArduinoUnit • normal uploading won’t work once the zumo and microbit are linked by I2C (because the I2C interrupts on the zumo interfere with uploading). So you’ll need use an alternative uploading method. You have to edit a file called boards.txt on your PC to enable the “Pololu A-Star 32U4 (bootloader port)” option in the tools -> boards menu. Follow the directions here under the “The bootloader-before-uploading method” to edit the boards.txt file. I found mine in “C:\Users\<username>\AppData\Local\Arduino15\packages\pololu-a-star\hardware\avr\4.0.2\boards.txt” on Windows 10 (your location might be different). Set your board to “Pololu A-Star 32U4 (bootloader port)” • turn on “verbose output during upload” so you can see why your upload failed (if it fails). Open file -> preferences and check the appropriate checkbox: • try to upload the “face towards opponent” sketch onto the zumo using “The reset” and “The bootloader-before-uploading method” instructions described here. If that works, you are good to go Ada/SPARK development environment setup • setup your dev environment for Ada: https://blog.adacore.com/ada-on-the-microbit • newer microbits are not recognized properly and will not flash. You can fix the problem with a one line change by following the directions here • create a copy of the scrolling text demo project and flash it to your microbit as per the instructions here • “build all” and “flash to board” as shown in the screenshot below • you should see “Make with Ada!” scroll across the 5 x 5 display on your microbit • clone or download the Ada drivers library. Put the code wherever you want. You’ll need it for the next step Create a copy of the project's code base • clone the High Integrity Sumobot’s code base from GitHub • edit the first line of the “microbit/high_integrity_sumobot.gpr” file to point to the location of your Ada drivers library code • open the project in GNAT Programming Studio (GPS) • set scenario in GPS: “ADL_BUILD: debug”, “ADL_BUILD_CHECKS: enabled”. These settings turn on the compiler switches you need for development and production (turning contract checking off doesn’t increase the frame rate so you might as well leave it on so it can catch errors) • build and flash the code to your microbit. The letter ‘U’ should show on the 5 x 5 display indicating an error (because the microbit is not in communication with the zumo yet) Building your sumobot Zumo headers • cut 3 female headers into groups of two. Use a box cutter to make a deep groove into the next channel on both sides of the header, break the header on the cut marks, then trim the sharp edges and discard the pin you cut free • trim the ends of the headers so it’s impossible for them to touch the batteries once they are soldered into place on the zumo (ruptured batteries are a safety risk) • cut 2 female swiss machine pin headers into groups of two • solder the headers to the zumo. Three of the headers are for wiring (see the wiring diagram) and the other two are the rear mount points for the bread board Note: you can ignore the headers that are crossed out in the image above (they are not used in this project). Level converter headers • cut 2 female headers into groups of six and solder them to the side of the level converter containing the electronics • set the level converter aside LCD wiring harness • make a wiring harness to free the LCD from the zumo main board. You’ll need two sets of male headers on one side and two sets of female headers on the other side. Make your wires 10 cm long. Solder all that together • install it in the zumo and turn it on to ensure the LCD still works • remove the wiring harness from the zumo, hot glue all the solder joints so they can never short together, and then reinstall it between the zumo and the LCD Zumo wires • cut the five wires that connect to the zumo (5V, 3.3V, ground, I2C clock, and I2C data). Each wire should be 15 cm long • strip both ends of each wire • cut 3 male headers into groups of two • solder the 3.3V and ground to a male header • solder the I2C lines to the next male header • and solder the 5V wire to the last male header (the other pin on that header isn’t soldered to anything). Note: a clothespin holds headers nicely while you solder them • wrap the individual wires soldered to the male headers in electrical tape to prevent shorts • insert the wires into the appropriate locations on the zumo • position the ends of each wire away from the rear of the zumo so they are out of the way for the next steps Breadboard • remove the power rail from the breadboard closest to column A • hot glue the level converter to the breadboard. Ensure the high voltage side is facing the rear of the bot (towards column J). Also ensure it is all the way to the edge or it will conflict with the microbit breakout board when you install it • use female swiss machine pin headers and hot glue to make the rear mounting points for the breadboard. Trim the pins on the long header as shown below to make more room for your wires. You might have to experiment to get the breadboard to a sufficient height from the zumo to allow all your wiring to fit between the zumo and the breadboard, depending on the thickness and length of your wiring • bend the LCD wiring harness so that the LCD sits just above the breadboard • cut two female swiss machine pin headers into groups of 8 to make the front mount points for the breadboard • test mount the breadboard on the zumo. If you’re happy with that glue one half of the front mount to the zumo • install (but do not glue) the breadboard headers on the zumo so you can simply place the breadboard on top of your mount points • if you’re happy with how your breadboard sits on the mount points remove the breadboard, put hot glue on top of the mounting headers, and then gently lower the breadboard onto the mount points. Note: do not glue all the mount points yet. You should be able to install and remove the breadboard easily with only gravity and friction holding it to the zumo. • put a row of hot glue on the rear mount point where you cut the pins off in a previous step to prevent the pins from rubbing your wiring. The bottom of your breadboard should look like this: Populating the breadboard • hot glue a single female swiss machine pin header under the left and right sides of the breakout board to hold it level against the breadboard • insert the microbit breakout board into the breadboard from B9 to B30 • insert the microbit into the breakout board with the display and buttons facing up • mount the breadboard to the zumo mount points • make the wires that connect from the breadboard to the level converter, leaving the wires longer than required and stripping the ends • connect all the wires as per the wiring diagram Uploading code and testing your sumobot • configure the zumo portion of the project to run its unit tests (see instructions in the loop function of zumo.ino) upload the code to the zumo and open the serial monitor to see the results • turn off the unit tests and upload the code to the zumo again • run “prove all” on the microbit code in SPARK. Command template: “gnatprove -P%PP -j0 %X –level=0 –ide-progress-bar –no-subprojects –mode=all” (if gnatprove complains about ‘uncompiled dependencies’ add “-f” to the command and try again) • configure the microbit portion of the project to run its unit tests (see instructions in main.adb) and flash them to the microbit. The number of passed assertions will scroll across the display if all the unit tests passed. A filename and line number will scroll across the display for the first failed test encountered • turn off the microbit unit tests and flash the code to the microbit • put your bot in your ring by itself and confirm it drives around searching for targets without leaving the ring • also confirm that it can find and push your training dummy out of the ring • you can also run the end-to-end tests if you like Finishing up the build (optional) If your High Integrity Sumobot is working but you want to make it a little more durable you can do these steps: • hot glue the breadboard mounting points more permanently • shorten the wires and hot glue them to the breadboard and level converter so they don’t come loose or snag during a fight • hot glue the LCD wiring harness to the side of the breadboard • give your bot another test alone in the ring and then with the training dummy to confirm it is operating correctly Information for the Make with Ada judges • I created this project as a solo effort (no teammates) • all of the code in the GitHub for the High Integrity Sumobot was created by me for this contest. However, the algorithm for this bot was inspired by a demo in the zumo library, which you can find here. I modified the code extensively but it still retains some of the state machine, variable names, and comments from the original demo. And the body of the last chance handler is roughly based on some code I found online, the source of which is documented in that file • none of the code in this project has been submitted in a previous Make with Ada competition • I realize that you won’t be able to easily run the code for this project without actually building the High Integrity Sumobot. But you can run the microbit unit tests if you have a microbit. And you can also run the zumo unit tests by uploading the code to almost any Arduino compatible microcontroller. After that you have gnatprove, other kinds of static analysis, and code reading • I plan to continue working on this project after the contest close date so please make sure you checkout the correct version of the code. I created release 0.1.0 just for you • I wanted to make sure you didn’t miss the written requirements and end-to-end tests in this 3, 200 word document. I think they are important but have never seen them emphasized in a project write-up Finally, I just wanted to make my case for why this project should score well for the “buzz effect” judging criteria. I spent a great deal of time brainstorming projects that might have “buzz effect” and reviewing all the other projects I could find that have used the Ada Drivers Library and this project is appealing on a number of fronts. First, many universities, schools, and clubs already have mini-sumo competitions because fighting robots are a fun and easy way to learn embedded software development. And for an extra ~$30 over the cost of the sumobot they were already going to buy, people can build this project, improve upon the code, and actually use Ada/SPARK in a project they care about. Therefore, I believe a sumobot programmed in Ada/SPARK should be relatively easy to market and promote.

Secondly, this project’s code base (excluding tests) is actually pretty small and the objectives of the sumobot are easy to understand. So, diving into the code base and making changes shouldn’t be too intimidating as compared to the Certyflie project, for example. And yet, the opportunities for improvements are vast. I can imagine a class of students starting with this code base, then having everyone making different improvements, and then competing to see who does the best.

Thirdly, this project is my first with Ada/SPARK and I couldn’t find any accessible examples of how to use the high dependability features of these languages or much guidance for how I should use them in the context of a full project. The book Building High Integrity Applications with SPARK was the best resource I found but even it was lacking in several areas. I’m sure I misused some features of Ada/SPARK and you guys are going to get a chuckle out of it but this is the kind of example I looked for and couldn’t find when I started this competition.

Fourthly, I tried to use not just Ada/SPARK but an entire suite of processes (within the time I had available) that would help ensure a high quality, low defect project. I think that’s where the benefits of Ada/SPARK really kick in and it reinforces the case for possibly using this project as a learning example.

Future improvements

• get gnattest working and port the homemade unit tests to gnattest
• replace the breadboard and wires with a custom-made circuit board (less likely to malfunction or suffer damage in competition)
• 3D print a cover (to protect the hardware on top of the zumo)
• write proofs for more of the code
• use the motor encoders to compensate for the effect battery voltage has on the sumobot’s speed (speed differences mess with the optimal time limits in the fighting algorithm)
• improve the fighting algorithm/add new algorithms (there is plenty of unused RAM and flash available on the microbit)

Resources

Final thoughts

I loved working on this project. It was interesting switching back and forth between C++ on the zumo and Ada/SPARK on the microbit. I’ve got about 10 years of experience building projects on the Arduino platform so I was able to jump right in on that side. Whereas the learning curve in SPARK and Ada were pretty steep.

But, once I got into it, I started to resent the Arduino compiler for not catching all the stupid ways C++ allows you to write incorrect programs that the Ada compiler would have caught. I also came to appreciate how clearly the Ada compiler told me exactly what I did wrong when I did make a mistake. And, by the end of the project, I had a lot more confidence in the correctness of the Ada/SPARK code as compared to the Arduino code. I can totally see why the AdaCore people are so excited about Ada/SPARK.

If you decide to build your very own High Integrity Sumobot or have questions about it, I’d love to hear from you in the comments below. Finally, if you find any bugs in this project please report them here.

Thanks for reading all the way to the end. Cheers.

• Access the project schematics here.
• Access the project code here.
]]>

Welcome to the Ada for micro:bit series where we look at simple examples to learn how to program the BBC micro:bit with Ada.

In this seventh part we will see how to use the accelerometer of the micro:bit. The accelerometer can, for instance, be used to know which way the micro:bit is oriented.

Interface

To get the acceleration value for all axes, we will just call the functionMicroBit.Accelerometer.Data. This function returns a record with X, Y and Z field giving the value for each axis.

function Data return MMA8653.All_Axes_Data;
--  Return the acceleration value for each of the three axes (X, Y, Z)

Return value:

• A record with acceleration value for each axis

Note that the type used to store the values of the accelerometer is declared in the package MMA8653 (the driver), so we have to with and use this package to have access to the operations for this type.

We can use the value in the record to get some information about the orientation of the micro:bit. For example, if the Y value is below -200 the micro:bit is vertical.

Here is the full code of the example:

with MMA8653; use MMA8653;

with MicroBit.Display;
with MicroBit.Display.Symbols;
with MicroBit.Accelerometer;
with MicroBit.Console;
with MicroBit.Time;

use MicroBit;

procedure Main is

Data : MMA8653.All_Axes_Data;

Threshold : constant := 150;
begin

loop

Data := Accelerometer.Data;

--  Clear the LED matrix
Display.Clear;

--  Draw a symbol on the LED matrix depending on the orientation of the
--  micro:bit.
if Data.X > Threshold then
Display.Symbols.Left_Arrow;

elsif Data.X < -Threshold then
Display.Symbols.Right_Arrow;

elsif Data.Y > Threshold then
Display.Symbols.Up_Arrow;

elsif Data.Y < -Threshold then
Display.Symbols.Down_Arrow;

else
Display.Display ('X');

end if;

--  Do nothing for 100 milliseconds
Time.Sleep (100);
end loop;
end Main;

Following the instructions of Part 1 you can open this example (Ada_Drivers_Library-master\examples\MicroBit\accelerometer\accelerometer.gpr), compile and program it on your micro:bit.

See you next week for the last Ada project of this series.

Don't miss out on the opportunity to use Ada in action by taking part in the fifth annual Make with Ada competition! We're calling on developers across the globe to build cool embedded applications using the Ada and SPARK programming languages and are offering over $9,000 in total prizes. Find out more and register today! ]]> Make with Ada 2020: CHIP-8 Interpreter https://blog.adacore.com/make-with-ada-2020-chip-8-interpreter Wed, 21 Oct 2020 00:00:00 -0400 Juliana Silva https://blog.adacore.com/make-with-ada-2020-chip-8-interpreter Laurent Zhu's and Damien Grisonnet's project won a finalist prize in the Make with Ada 2019/20 competition. This project was originally posted on Hackster.io here. For those interested in participating in the 2020/21 competition, registration is now open and project submissions will be accepted until Jan 31st 2021, register here. CHIP-8 language interpreter for STM32F429 Discovery board Story Motivations This project was accomplished for the EPITA Ada courses and the Make With Ada contest. Originally this project was supposed to be a GBA emulator. However, instead of implementing one from scratch in ADA we wanted to port an existing one on the STM32F429 Discovery and write bindings in ADA. But, while trying to port the emulator we noticed that there was not much to do in ADA and that the project would be mostly written in C instead of ADA. We thought that it was a pity since we were supposed to do a project in ADA. Then, we decided to switch to another project that would allow us to write more ADA. We could have written our own GBA emulator in ADA but it was too big of a challenge to write one on time for the challenge. Thus, we decided to write this CHIP-8 emulator instead which involves the same coding challenges as the GBA one but it is way faster to implement. CHIP-8 The first step of the project was to understand how the emulator is working: Memory Memory size of 4K with the first 512 bytes of the memory space reserved for the CHIP-8 interpreter. It is common to use this reserved space to store font data. Registers CHIP-8 has 16 8-bit data registers named V0 to VF. Stack The stack is only used to store return addresses when subroutines are called. In modern implementations stacks can store up to 16 elements. Timers CHIP-8 has two timers. They both count down at 60 hertz, until they reach 0. • Delay timer: This timer is intended to be used for timing the events of games. Its value can be set and read. • Sound timer: This timer is used for sound effects. When its value is nonzero, a beeping sound is made. Inputs Input is done with a hex keyboard that has 16 keys ranging 0 to F. This keyboard is displayed on the bottom part of the screen of the STM32F429 Discovery. Graphics Original CHIP-8 Display resolution is 64×32 pixels, and color is monochrome. Graphics are drawn to the screen solely by drawing sprites, which are 8 pixels wide and may be from 1 to 15 pixels in height. Sprite pixels are XOR'd with corresponding screen pixels. In other words, sprite pixels that are set flip the color of the corresponding screen pixel, while unset sprite pixels do nothing. The carry flag (VF) is set to 1 if any screen pixels are flipped from set to unset when a sprite is drawn and set to 0 otherwise. This is used for collision detection. Since the STM32F429 Discovery screen resolution is 320x240, the display of the ROM was scale 5 times to improve the user experience and to match the platform. Sound A beeping sound is supposed to be played when the value of the sound timer is nonzero. However, since the STM32F429 Discovery does not have any audio module, no sound are played. Opcode Table CHIP-8 has 35 opcodes, which are all two bytes long and stored big-endian. CHIP-8 Interpreter The different steps of the interpreter: • The screen, the touch panel and the layers are initialized • We draw the keyboard on the bottom of screen with the first layer, by using the CHIP-8 sprites. In order to do that, we iterate through all the existing keys and from their position in the font set table, we can draw it easily • A ROM is loaded with the Load_Rom procedure. The ROMs are located in the file that we generate with a python script (scripts/gen_rom.py). It generates all the Ada arrays from all the ROMs located in the roms/ directory. • Then, we have our main loop: Main Loop • An opcode, consisting of 2 bytes, is fetched from the memory at the program counter position • We call the right function to execute by looking at the 4 first bits of our opcode. Some instruction will not increment the program counter, some will increment it and some will skip the next instruction by incrementing 2 times more • At the end of the loop we read the touch screen inputs and we update the list of pressed keys accordingly Setup the project git clone https://github.com/laurentzh/CHIP-8.git cd CHIP-8 git clone --recursive https://github.com/AdaCore/Ada_Drivers_Library.git source env.sh python2 Ada_Drivers_Library/scripts/install_dependencies.py Compile the project gprbuild --target=arm-eabi -d -P chip8.gpr -XLCH=led -XRTS_Profile=ravenscar-sfp -XLOADER=ROM -XADL_BUILD_CHECKS=Disabled src/main.adb -largs -Wl,-Map=map.txt Flash the project arm-eabi-objcopy -O binary objrelease/main objrelease/main.bin st-flash --reset write objrelease/main.bin 0x8000000 Add a ROM cp$ROM roms/
./scripts/gen_rom.py roms/ src/roms.ads

Change ROM

We encountered a few memory problems with the implementation of the menu. So, in order to choose the ROM, you need to change the argument of the call to Load_Rom in the main.adb file.

Examples

Simple ROM that does not require user interaction:

Demonstrations

• Access the project schematics here.
• Access the project code here.
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Welcome to the Ada for micro:bit series where we look at simple examples to learn how to program the BBC micro:bit with Ada.

In this sixth part we will see how to read the analog value of a pin. This means reading a value between 0 and 1023 that tells the voltage applied to the pin. 0 means 0 volts, 1023 means 3.3 volts.

Wiring Diagram

For this example we will need a couple of extra parts:

• An LED
• A 470 ohm resistor
• A potentiometer
• A couple of wires to connect them all

For this example we start from the same circuit as the pin output example, and we add a potentiometer. The center pin of the potentiometer is connected to pin 1 of the micro:bit the other two pins are respectively connected to GND and 3V.

Interface

To read the analog value of the IO pin we are going to use the function Analogof the package MicroBit.IOs.

function Analog (Pin : Pin_Id) return Analog_Value
with Pre => Supports (Pin, Analog);
--  Read the voltagle applied to the pin. 0 means 0V 1023 means 3.3V

Arguments:

• Pin : The id of the pin that we want read the analog value from

Precondition:

• The function Analog has a precondition that the pin must support analog IO.

In the code, we are going to write an infinite loop that reads the value of pin 1, and set pin 0 to the same value.

This means that you can control the brightness of the LED using the potentiometer.

Here is the full code of the example:

with MicroBit.IOs;

procedure Main is

Value : MicroBit.IOs.Analog_Value;
begin

--  Loop forever
loop

--  Read analog value of pin
Value := MicroBit.IOs.Analog (1);

--  Write analog value of pin 0
MicroBit.IOs.Write (0, Value);
end loop;
end Main;

Following the instructions of Part 1 you can open this example (Ada_Drivers_Library-master\examples\MicroBit\analog_in\analog_in.gpr), compile and program it on your micro:bit.

See you next week for another Ada project on the micro:bit.

Don't miss out on the opportunity to use Ada in action by taking part in the fifth annual Make with Ada competition! We're calling on developers across the globe to build cool embedded applications using the Ada and SPARK programming languages and are offering over $9,000 in total prizes. Find out more and register today! ]]> Ada for micro:bit Part 5: Analog Output https://blog.adacore.com/ada-for-microbit-part-5-analog-output Tue, 13 Oct 2020 08:25:00 -0400 Fabien Chouteau https://blog.adacore.com/ada-for-microbit-part-5-analog-output Welcome to the Ada for micro:bit series where we look at simple examples to learn how to program the BBC micro:bit with Ada. If you haven't already, please follow the instructions in Part 1 to setup your development environment. In this fifth part we will see how to write an analog value to a pin. The micro:bit doesn't have a real digital to analog converter, so the analog signal is actually a Pulse Width Modulation (PWM). This is good enough to control the speed of a motor or the brightness of an LED. There is a limit of three analog (PWM) signals on the micro:bit, if you try to write an analog value to more than three pins an exception will be raised. Wiring Diagram For this example we use the same circuit as the pin output example. Interface To write an analog value to the IO pin we are going to use the procedure Write of the package MicroBit.IOs. procedure Write (Pin : Pin_Id; Value : Analog_Value) with Pre => Supports (Pin, Analog); Arguments: • Pin : The id of the pin that we want to read as digital input • Value : The analog value for the pin, between 0 and 1023 Precondition: • The procedure Write has a precondition that the pin must support analog IO. In the code, we are going to write an loop with a value that goes from 0 to 128 and set write this value to pin 0. We could go from 0 to 1023 but since the LED doesn't get brighter after 128, there is no need to go beyond that value. We also use the procedure Delay_Ms of the package MicroBit.Time to stop the program for a short amount of time. Here is the full code of the example: with MicroBit.IOs; with MicroBit.Time; procedure Main is begin -- Loop forever loop -- Loop for value between 0 and 128 for Value in MicroBit.IOs.Analog_Value range 0 .. 128 loop -- Write the analog value of pin 0 MicroBit.IOs.Write (0, Value); -- Wait 20 milliseconds MicroBit.Time.Delay_Ms (20); end loop; end loop; end Main; Following the instructions of Part 1 you can open this example (Ada_Drivers_Library-master\examples\MicroBit\analog_out\analog_out.gpr), compile and program it on your micro:bit. See you next week for another Ada project on the micro:bit. Don't miss out on the opportunity to use Ada in action by taking part in the fifth annual Make with Ada competition! We're calling on developers across the globe to build cool embedded applications using the Ada and SPARK programming languages and are offering over$9,000 in total prizes. Find out more and register today!

]]>

Welcome to the Ada for micro:bit series where we look at simple examples to learn how to program the BBC micro:bit with Ada.

In this fourth part we will see how to read the digital state of a pin. This means reading if the pin is at 0 volts (low) or 3.3 volts (high).

Wiring Diagram

For this example we will need a couple of extra parts:

• An LED
• A 470 ohm resistor
• A push button
• A couple of wires to connect them all

We start from the same circuit as the part three example, and we add a push button.

Interface

To control the IO pin we are going to use the function Set of the package MicroBit.IOs.

function Set (Pin : Pin_Id) return Boolean
with Pre => Supports (Pin, Digital);

Arguments:

• Pin : The id of the pin that we want to read as digital input

Precondition:

• The procedure Set has a precondition that the pin must support digital IO.

As you can see, the function Set to read the pin has the same name as the procedure set that we used to control the pin in the output example. It is called overloading, two subprograms with the same name that provide different services.

In the code, we are going to write an infinite loop that reads the state of pin 1. If it is high, it means the button is not pressed so we turn off the LED on pin 0. It if it is low, it means the button is pressed so we turn on the LED on pin 0.

Here is the full code of the example:

with MicroBit.IOs;

procedure Main is
begin

--  Loop forever
loop

--  Check if pin 1 is high
if MicroBit.IOs.Set (1) then

--  Turn off the LED connected to pin 0
MicroBit.IOs.Set (0, False);
else

--  Turn on the LED connected to pin 0
MicroBit.IOs.Set (0, True);
end if;
end loop;
end Main;

Following the instructions of Part 1 you can open this example (Ada_Drivers_Library-master\examples\MicroBit\digital_in\digital_in.gpr), compile and program it on your micro:bit.

See you next week for another Ada project on the micro:bit.

Don't miss out on the opportunity to use Ada in action by taking part in the fifth annual Make with Ada competition! We're calling on developers across the globe to build cool embedded applications using the Ada and SPARK programming languages and are offering over $9,000 in total prizes. Find out more and register today! ]]> AdaCore Code of Conduct https://blog.adacore.com/adacore-code-of-conduct Mon, 05 Oct 2020 10:19:00 -0400 Fabien Chouteau https://blog.adacore.com/adacore-code-of-conduct Starting today, AdaCore has put in place a Code of Conduct (CoC) to ensure a positive environment for everyone willing and wanting to interact with us. With the development of this blog, our twitter accounts, and our GitHub corporate account, there is more and more communication between AdaCore and a number of communities. In this Code of Conduct we want to explain how we are going to moderate the AdaCore-maintained community spaces with the goal of maintaining a welcoming, friendly environment. The full Code of Conduct can be found on our website: https://www.adacore.com/company/code-of-conduct Here is the introduction of the document: We expect this code of conduct to be followed by anyone who contributes to AdaCore-maintained community spaces such as Github repositories, public and private mailing lists, issue trackers, wikis, blogs, Twitter, and any other communication channel maintained by AdaCore, and by anyone who participates to an activity organised by AdaCore. It applies equally to users, moderators, administrators, AdaCore staff, partners, and community members. This code is not exhaustive or complete. It serves to distill our common understanding of a collaborative, shared environment and goals. We expect it to be followed in spirit as much as in the letter, so that it can enrich all of us and the technical communities in which we participate. Diversity Statement AdaCore welcomes and encourages participation by everyone. We are committed to being a community that everyone feels good about joining. Although we may not be able to satisfy everyone, we will always work to treat everyone well. No matter how you identify yourself or how others perceive you: we welcome you. ]]> Ada for micro:bit Part 3: Pin Output https://blog.adacore.com/ada-for-microbit-part-3-pin-output Tue, 29 Sep 2020 08:06:00 -0400 Fabien Chouteau https://blog.adacore.com/ada-for-microbit-part-3-pin-output Welcome to the Ada for micro:bit series where we look at simple examples to learn how to program the BBC micro:bit with Ada. If you haven't already, please follow the instructions in Part 1 to setup your development environment. In this third part we will see how to control the output state of a micro:bit pin by lighting an LED. Wiring Diagram For this example we will need a couple of extra parts: • A breadboard • An LED • A 470 ohm resistor • A couple of wires to connect them all Wiring the LED directly from the output pin to ground will make it burn, so we have to add a resistor to limit the flow of current. Interface To control the IO pin we are going to use the procedure Set of the package MicroBit.IOs. procedure Set (Pin : Pin_Id; Value : Boolean) with Pre => Supports (Pin, Digital); Arguments: • Pin : The id of the pin that we want to control as digital output • Value : A Boolean that says if we want the pin to be high (True) or low (False) Precondition: • The procedure Set has a precondition that the pin must support digital IO. We also use the procedure Delay_Ms of the package MicroBit.Time to stop the program for a short amount of time. Here is the full code of the example: with MicroBit.IOs; with MicroBit.Time; procedure Main is begin -- Loop forever loop -- Turn on the LED connected to pin 0 MicroBit.IOs.Set (0, True); -- Wait 500 milliseconds MicroBit.Time.Delay_Ms (500); -- Turn off the LED connected to pin 0 MicroBit.IOs.Set (0, False); -- Wait 500 milliseconds MicroBit.Time.Delay_Ms (500); end loop; end Main; Following the instructions of Part 1 you can open this example (Ada_Drivers_Library-master\examples\MicroBit\digital_out\digital_out.gpr), compile and program it on your micro:bit. See you next week for another Ada project on the micro:bit. Don't miss out on the opportunity to use Ada in action by taking part in the fifth annual Make with Ada competition! We're calling on developers across the globe to build cool embedded applications using the Ada and SPARK programming languages and are offering over$9,000 in total prizes. Find out more and register today!

]]>

In my current job and in my previous job, I was involved in customer support where a user of our product would basically say “your tool does not handle my code correctly.” Whether the problem really is in our tool, or actually with their code, we need to come up with a simplified version of the “incorrect” code. This is not always a trivial task. The customer may be dealing with a fairly large chunk of code, and they may not understand their codebase or our tool well enough to narrow down the problem. It is usually faster for the customer to send as much code as possible, and let the support engineer try to shrink it to a manageable size, as they likely have a better understanding of the tool and what is happening to the example code.

So now we are onto the next problem – how does the support engineer get a copy of the problematic code? In most instances, the customer is very protective of their code – proprietary information, algorithms, or data structures; sometimes even object names can convey information! For example, suppose you had a constant:

Constant_Max_Velocity : constant Velocity_T := 12.3;

That name and value tells you something important about the widget being built.

One could always use some search-and-replace mechanisms to mask some names, but that gets into some complicated matching algorithms, and consistency issues. In just the simple example above, how would you replace “constant” in one place and not the other, and do you replace “Velocity” with the same token or different tokens. What we really want is a tool that understands the language, and can intelligently replace tokens based on their context, not just their textual value.

So I took on the task of writing a tool to “obfuscate” a code base. One of the best ways to write a tool like this for Ada is using Libadalang. This library (available in Ada or Python) allows the user to parse Ada files and gather semantic information. I could then use the semantic information to track every name definition and replace the name wherever it was referenced. With Libadalang, I did not need to worry about the above problems with “constant” and “velocity”.

I chose to write this tool in Ada, because Ada gave me access to the GNATcoll bindings. With these bindings, I could parse GNAT Project Files to find all the source files I needed. This allowed me to obfuscate an entire codebase – once modified, the new codebase should compile correctly and even perform correctly, although none of the variable names would make any sense!

As a simple example, I wrote a small program that solve the “N-Queens” problem (wikipedia) There is not much “proprietary” information in the code, but looking at the original code and obfuscated code side-by-side shows the information given just by changing names.

Even in this simple example, we see the information “loss”. On the left, we know we are working with a Board made of Rows and Columns; on the right, all we know is we are looping through some counter.

When I started writing this application, I chose to use SPARK subset of the Ada language rather than the full language. With this, I could use the SPARK provers to help ensure the robustness and correctness of the code. My plan was to get all of the code to SPARK “Silver” level (proving the absence of run-time errors), with some of the code reaching “Gold” level (proving functions actually do what they are supposed to).

The first iteration of coding was just writing the application to be SPARK-compliant (“Stone” level). Even this low level required some re-thinking of coding practices. For example: I have a subprogram that basically returns the next available obfuscated name. In Ada, I would write this as a function that increments a counter and returns the next item. In SPARK, functions cannot have side effects (modifying global data) so this function had to be re-written as a procedure. Not a difficult task, but a language constraint (that makes your code a little safer!)

Next, I moved onto flow analysis (“Bronze” level). This involved adding Global dependency contracts (showing which global data was read/modified by which subprogram) and turning off SPARK for some subprograms that had to deal with non-SPARK compliant run-time code. With SPARK, an interface can be in SPARK mode while its implementation is not in SPARK mode. I used this to wrap some non-SPARK code (like some run-time packages) and make my SPARK analysis happier.

In my mind, the most important step was making sure I didn’t have any of the typical overflow issues – an absence of run-time errors (“Silver”) level. This is not as easy as it seems, especially when dealing with strings. The simple act of concatenating two strings raises a lot of flags in proving there are no run-time errors.

procedure Do_Something (Y : String; Z : String) with
SPARK_Mode
is
X : constant String :=
Y (Y'first + 1 .. Y'last) & Z (Z'first .. Z'last - 1);
begin

The most obvious concern is what happens when concatenating the two strings creates a string that is too long for X to hold. You would need to add preconditions on the lengths of Y and Z. But, because the index range of a string is integer, then Y’first could be integer’last – and adding one to that can generate a constraint error (similarly for Z’last and integer’first). As a human, you “know” these indices are not going to happen, but for SPARK to accept it, you have to prove the indices are not going to happen – either through preconditions, better type definitions, or redefining the concatenation operator.

Finally, I started implementing some contracts on subprogram behavior to prove that the subprogram did what I thought it should do (“Gold”) level. For some subprograms that is easy, some of them take a little more thought, and for some, the effort involved in proving correctness (SPARK) is more than the effort involved in showing correctness (testing). This process is still ongoing, and others who work with this application are welcome to contribute!

This application is available on GitHub for use or investigation. This code obfuscator is still a work in progress and will benefit from future Libadalang evolutions in order to support a more complete set of Ada features. But it is a good start to help those of us in the customer support world help our customers!

]]>

Welcome to the Ada for micro:bit series where we look at simple examples to learn how to program the BBC micro:bit with Ada.

In this second part we will see how to use the two buttons of the micro:bit.

Interface

To know if a button is pressed or not, we will use the function State of the MicroBit.Buttons package.

type Button_State is (Pressed, Released);

type Button_Id is (Button_A, Button_B);

function State (Button : Button_Id) return Button_State;

Arguments:

• Button : The Id of the button that we want to check. There are two Ids: Button_A or Button_B.

Return value:

• The function State return the Button_State that can be either Pressed or Released.

We will use this function to display the letter A if the button A is pressed, or the letter B if the button is pressed.

Here is the full code of the example:

with MicroBit.Display;
with MicroBit.Buttons; use MicroBit.Buttons;
with MicroBit.Time;

procedure Main is
begin

loop
MicroBit.Display.Clear;

if MicroBit.Buttons.State (Button_A) = Pressed then
--  If button A is pressed

--  Display the letter A
MicroBit.Display.Display ('A');

elsif MicroBit.Buttons.State (Button_B) = Pressed then
--  If button B is pressed

--  Display the letter B
MicroBit.Display.Display ('B');
end if;

MicroBit.Time.Delay_Ms (200);
end loop;
end Main;

Following the instructions of Part 1 you can open this example (Ada_Drivers_Library-master\examples\MicroBit\buttons.gpr), compile and program it on your micro:bit.

See you next week for another Ada project on the micro:bit.

First program

Start the GNATstudio development environment that you installed earlier, click on "Open Project" and select the file "Ada_Drivers_Library-master\examples\MicroBit\text_scrolling/text_scrolling.gpr" from the archive that you extracted earlier.

Click on the "Build all" icon in the toolbar to compiler the project.

Plug your micro:bit using a USB micro cable.

And finally click on the "Flash to board" icon in the toolbar to run the program on the micro:bit.

You should see a text scrolling on the LEDs of the micro:bit:

That's it for the setup of your Ada development environment for the micro:bit. See you next week for another Ada project on the micro:bit.

Don't miss out on the opportunity to use Ada in action by taking part in the fifth annual Make with Ada competition! We're calling on developers across the globe to build cool embedded applications using the Ada and SPARK programming languages and are offering over $9,000 in total prizes. Find out more and register today! ]]> GNATcoverage: getting started with instrumentation https://blog.adacore.com/gnatcoverage-getting-started-with-instrumentation Thu, 10 Sep 2020 08:34:33 -0400 Pierre-Marie de Rodat https://blog.adacore.com/gnatcoverage-getting-started-with-instrumentation This is the second post of a series about GNATcoverage and source code instrumentation. The previous post introduced how GNATcoverage worked originally and why we extended it to support source instrumentation-based code coverage computation. Let’s now see it in action in the most simple case: a basic program running on the host machine, i.e. the Linux/Windows machine that runs GNATcoverage itself. Source traces handling Here is a bit of context to fully understand the next section. In the original GNATcoverage scheme, coverage is inferred from low level execution trace files (“*.trace”) produced by the execution environment. These traces essentially contain a summary of program machine instructions that were executed. We call these “binary traces”, as the information they refer to is binary (machine) code. With the new scheme, based on the instrumentation of source code, it is instead the goal of each instrumented program to create trace files. This time, the information in traces refers directly to source constructs (declarations, statements, IF conditions, …), so we call them “source traces” (“*.srctrace” files). The data stored in these files is conceptually simple: some metadata to identify the sources to cover and a sequence of booleans that indicate whether each coverage obligation is satisfied. However, for efficiency reasons, instrumented programs must encode this information in source traces files using a compact format, which is not trivial to produce. To assist instrumented programs in this task, GNATcoverage provides a “runtime for instrumented programs” as a library project: gnatcov_rts_full.gpr, for native programs which have access to a full runtime (we will cover embedded targets in a future post). Requirements First, the GNATcoverage instrumenter needs a project file that properly describes the closure of source files to instrument as well as the program main unit. This is similar to what a compiler needs: access to all the dependencies of a source file in order to compile it. Next, this blog series assumes the use of a recent GPRbuild (release 20 or beyond), for the support of two switches specifically introduced to facilitate building instrumented sources without modifying project files. What the new options do is conceptually simple so it would be possible to build without this, just less convenient. Then the source constructs added by the instrumentation expect an Ada 95 compiler. The instrumenter makes several compiler-specific assumptions (for instance when handling Pure/Preelaborate units), so for now we recommend using a recent GNAT compiler. Finally, users need to build and install the “runtime for instrumented programs” described in the previous section. To make sure the library code can be linked with the program to analyze, the library first needs to be built with the same toolchain, then installed: # Create a working copy of the runtime project. # This assumes that GNATcoverage was installed # in the /install/path/ directory.$ rm -rf /tmp/gnatcov_rts
$cp -r /install/path/share/gnatcoverage/gnatcov_rts /tmp/gnatcov_rts$ cd /tmp/gnatcov_rts

# Build the gnatcov_rts_full.gpr project and install it in
# gprinstall’s default prefix (most likely where the toolchain is installed).
$gprbuild -Pgnatcov_rts_full$ gprinstall -Pgnatcov_rts_full

Note that depending on your specific setup, the above may not work without special filesystem permissions, for instance if the toolchain/GPRbuild was installed by a superuser. In that case, you can install the runtime to a dedicated directory and update your environment so that GPRbuild can find it: add the --prefix=/dedicated/directory argument to the gprinstall command, and add that directory to the GPR_PROJECT_PATH environment variable.

A first example

Now that prerequisites are set up, we can now go ahead with our first example. Let’s create a very simple program:

--  example.adb
procedure Example is
begin
[…]
end Example;

-- example.gpr
project Example is
for Object_Dir use “obj”;
end Example;

Before running gnatcov, let’s make sure that this project builds fine:

$gprbuild -Pexample -p$ obj/example
[…]

Great. So now, let’s instrument this program to compute its code coverage:

$gnatcov instrument -Pexample --level=stmt --dump-trigger=atexit As its name suggests, the “gnatcov instrument” command instruments the source code of the given project. The -Pexample and --level=stmt options should be familiar to current GNATcoverage users: the former requests the use of the “example.gpr” project, to compute the code coverage of all of its units, and --level=stmt tells gnatcov to analyze statement coverage. The --dump-trigger=atexit option is interesting. As discussed earlier, instrumented programs need to dump their coverage state into a file (the trace file), that “gnatcov coverage” reads in order to produce a coverage report. But when should that dump happen? Since one generally wants reports to show all discharged obligations (fancy words meaning: executed statements, decision outcomes exercized, …), the goal is to create the trace file after all code has executed, right before the program exits. However some programs are designed to never stop, running an endless loop (Ravenscar profile), so this trace file creation moment needs to be configurable. --dump-trigger=atexit tells the instrumenter to use the libc’s atexit routine to trigger file creation when the process is about to exit. It’s suitable for most programs running on native platforms, and makes trace file creation automatic, which is very convenient. Now is the time to build the instrumented program: $ gprbuild -Pexample -p --src-subdirs=gnatcov-instr --implicit-with=gnatcov_rts_full

Even seasoned GPRbuild users will wonder about the two last options.

--src-subdirs=gnatcov-instr asks GPRbuild to consider, in addition to the regular source directories, all “gnatcov-instr” folders in object directories. Here that means that GPRbuild will first look for sources in “obj/gnatcov-instr” (as “obj” is example.gpr’s object directory), then for sources in “.” (example.gpr’s regular source directory).

But what is “obj/gnatcov-instr” anyway? When it instruments a project, gnatcov must not modify the original sources, so instead it stores instrumented sources in a new directory. The general rule of thumb for programs that deal with project files is to use projects’ object directory (Object_Dir attribute) to store artifacts; “gnatcov instrument” thus creates a “gnatcov-instr” subdirectory there and puts instrumented sources in it. Afterwards, passing --src-subdirs to GPRbuild is the way to tell it to build instrumented sources instead of the original ones.

The job of --implicit-with=gnatcov_rts_full is simple: make GPRbuild consider that all projects use the gnatcov_rts_full.gpr project, even though they don’t contain a “with “gnatcov_rts_full”;” clause. This allows instrumented sources (in obj/gnatcov-instr) to use features in the gnatcov_rts_full project even though “example.gpr” does not request it.

In other words, both --src-subdirs and --implicit-with options allow GPRbuild to build instrumented sources with their extra requirements without having to modify the project file of the project to test/cover (example.gpr).

We are getting closer to the coverage report. All we have to do it to finally run the instrumented program, to create a source trace file:

$obj/example Fact (1) = 1$ ls *.srctrace
example.srctrace

So far, so good. By default, the instrumented program creates in the current directory a trace file called “XXX.srctrace” where XXX is the basename of the executed binary, but one can choose a different filename by setting the GNATCOV_TRACE_FILE environment variable to the name of the trace file to create. Now that we have a trace file, the rest will be familiar to GNATcoverage users:

$gnatcov coverage -Pexample --level=stmt --annotate=xcov example.srctrace$ cat obj/example.adb.xcov
75% of 4 lines covered
Coverage level: stmt
2 .:
3 .: procedure Example is
4 .:    function Fact (N : Natural) return Natural is
5 .:    begin
6 +:       if N <= 1 then
7 +:          return 1;
8 .:       else
9 -:          return N * Fact (N - 1);
10 .:       end if;
11 .:    end Fact;
12 .: begin
13 +:    Put_Line ("Fact (1) =" & Fact (1)'Image);
14 .: end Example;

And voilà! That’s all for today. The next post will demonstrate how to handle programs running on embedded targets.

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Introducing source code instrumentation in GNATcoverage https://blog.adacore.com/source-code-instrumentation-in-gnatcoverage Tue, 08 Sep 2020 08:41:37 -0400 Pierre-Marie de Rodat https://blog.adacore.com/source-code-instrumentation-in-gnatcoverage

This is the first post of a series about GNATcoverage and source code instrumentation.

In order to make GNATcoverage viable in more contexts, we planned several years ago to add instrumentation support in GNATcoverage for Ada sources. This feature reached maturation recently and is available in the last Continuous Release, so it is a good time to present it with a blog series!

GNATcoverage background

GNATcoverage is the tool developed by AdaCore to compute the code coverage of Ada/C programs, available on several platforms, both native and embedded. It is able to assess DO-178 criteria, up to MC/DC, on C and Ada programs, from Ada 95 to Ada 2012.

Its coverage analysis capabilities are versatile. Several output formats are available: “xcov”, which look like text coverage reports from GCC’s gcov tool, a set of static HTML pages, a single modern dynamic HTML page, a XML report for machine processing or a custom text report suitable for certification contexts. In addition, the tool features powerful consolidation capabilities, which allow combining the result of multiple executions into a single report in several fashions and let users specify which packages/sources are of actual relevance to an analysis.

The way it works so far is atypical for a code coverage tool working with programs compiled to machine code:

1. The source code is compiled unchanged to machine code (processor instructions), with a few special compiler options to ease the mapping of machine code back to source constructs and to generate a list of source constructs to cover: the SCOs (Source Coverage Obligations).

2. An “instrumented” execution environment runs the program and generates a “trace file”, which roughly contains the set of machine instructions executed. For native platforms, this execution environment could be Valgrind or DynamoRIO, while embedded programs would execute either in GNATemulator or on a physical board with a hardware probe attached.

3. GNATcoverage computes the coverage report from trace files, compiled programs, source files and the SCOs.

Unlike traditional code coverage tools, which generally inject code in the compiled program so that the program itself computes its coverage state, GNATcoverage works on unmodified programs. Instead, the execution environment builds the coverage state on behalf of the executed program.This has several advantages:

• The main one is that executable code used for coverage analysis can also be used in production, or at least will be very close since the program itself is not modified to embed coverage measurement code and data structures.

• It allows object coverage analysis (coverage of individual processor-level instructions), which is useful in source/object traceability studies.

So why do we need instrumentation?

Our original approach also comes with drawbacks. For instance, the execution environment is an emulator on native platforms (Valgrind, DynamoRIO) which incurs a non trivial performance penalty. For embedded targets, there is sometimes no possibility to use GNATemulator and no hardware probe available to create execution traces, or setting up such a probe can prove tricky enough to turn out impractical. Hence, depending on specific situations, instrumenting programs for code coverage can be a better fit than using unmodified programs.

One of the design goals for this instrumentation scheme was to be as close as possible to the original one. This facilitates transitions from one mode to the other, and makes most existing features, in particular in coverage analysis capabilities, applicable to both in a consistent manner.

The next post will present how the instrumentation scheme works in GNATcoverage with a simple example program.

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The FACE™ open systems strategy gaining traction in the avionics industry https://blog.adacore.com/face-open-systems-strategy-for-avionics Thu, 03 Sep 2020 09:10:00 -0400 Jessie Glockner https://blog.adacore.com/face-open-systems-strategy-for-avionics

The Future Airborne Capability Environment™ (FACE) approach is a US government-industry initiative for reducing avionics system life cycle costs through software component reuse. The technical foundation is based on a portability-focused language-agnostic architecture and data model, common interfaces, and commercially available standards (IDL, POSIX®, and ARINC-653). The architecture comprises a number of segments -- Portable Components Segment, Transport Services Segment, Operating System Segment, Platform-Specific Services Segment, and I/O Services Segment -- that reflect a separation of concerns between portable and platform-dependent functionality.

Cost savings through reusable components is hardly a new idea, and it may look simple on the surface but has proved difficult to achieve in practice. Some hurdles are technical; designing a component for reusability requires experience and a broad set of skills, and programming languages differ in how effectively and efficiently they support reuse. And on the non-technical side, contracts for computer-based systems have historically offered little incentive, for either the agency procuring the software or the company developing it, to focus on reusability. Designing for reuse would add cost to the initial effort but realize savings only on later projects.

The FACE approach resolves this dilemma in a number of ways:

First, from the outset in 2010, it has involved all of the stakeholders: government agencies procuring avionics software, defense contractors developing the software, RTOS vendors and other suppliers of platform infrastructure, and software development and verification tool providers. The FACE Technical Standard has evolved as a consensus from these stakeholders in a “bottom-up” fashion, reflecting the state of the practice in industry.

Second, the FACE Technical Standard has taken a language agnostic approach, with support for the major languages used in airborne systems: C, C++, Ada, and (mostly for user interface components) Java.

Third, it does not attempt to address design goals other than portability. Although the FACE Technical Standard lays out operating system profiles and language subsets (“capability sets”) reflecting safety- and/or security-based restrictions, the FACE conformance tests do not check functionality, safety, or security properties. (Verification of these properties is obviously important but needs to be done separately, using appropriate analysis tools and the relevant certification standard. The SPARK Pro formal methods-based toolset is especially useful in this context, for example allowing proof that the code does not raise any exceptions.) The FACE conformance tests only check the software component and its associated data model -- officially known as a “Unit of Conformance” (UoC) -- to ensure that all interfaces are used correctly. A UoC that provides a FACE service, for example an RTOS in the Operating System Segment, has to implement all of the interfaces that the UoC’s profile requires. A UoC that functions as an application, for example flight management code in the Portable Components Segment, must only use interfaces that the UoC’s profile allows.

And fourth, reflecting the maturity of the FACE Technical Standard and the growing number of certified UoCs, DoD agencies are increasingly adding FACE conformance to the requirements for new avionics system procurements.  Indeed, the FACE approach not only benefits the avionics industry here in the US, but could have a larger impact on other industries in the US and abroad as companies recognize the benefits of reusing portable technologies to provide faster, more cost-effective deployment of new systems and software platforms.

AdaCore is committed to the success of the FACE approach

Both the Ada programming language and the company’s product offerings directly support the FACE initiative’s objectives while helping developers meet the high assurance requirements that are typical in airborne software.

AdaCore has been an active contributor to both the technical and business sides of the FACE community since 2012 and is a Principal Member of the FACE Consortium. Company representatives have served as key members of the Conformance and Operating System Segment (OSS) Subcommittees, where they have reviewed the various versions of the FACE Technical Standard, helped formulate effective policies and procedures for conformance, and worked to incorporate support for Ada 2012 Safety capability sets in the FACE Technical Standard so that developers can take advantage of contract-based programming and other modern features. And just recently, Dr. Benjamin Brosgol, a member of AdaCore’s senior technical staff, was elected Vice Chair of the Technical Working Group by The Open Group FACE™ Consortium Steering Committee.

On September 21, 2020, The Open Group will host its annual FACE™ and SOSA™ (Sensor Open Systems Architecture) Consortia Technical Interchange Meeting. AdaCore is the Premier Sponsor of this year’s free, virtual event offering paper presentations by leading experts from government, industry, and academia on the use of the FACE Technical Standard and related business practices.

On March 23, 2021, The Open Group will host a live FACE™ and SOSA™ Consortia Exposition & Technical Interchange Meeting at the Holiday Inn Solomons Conference Center & Marina in Solomons, MD. This Technical Interchange Meeting will consist of keynote speakers, a panel discussion, and FACE 101 and SOSA 101 sessions. AdaCore will also be the Premier Sponsor of this event, and will be exhibiting alongside customers and partners such as Wind River, Lynx Software Technologies, Rapita, and Verocel.

]]>

Story

Genesis

Last MakeWithAda I worked on getting BLE going with an STM32L476. This project is another communication protocol: LoRa. This came about as my wife and I were musing about how to detect and deter unwanted garden visitors (cats that come into the garden to use it as a toilet and leave, squirrels that climb the walls and scratch at the soffits). Step one for dealing with this was to ID the interlopers using sensors, Step two would be some form of countermeasure. I am sort of at step 0.5 The basic sensor comm is up but the detection and countermeasures have not been decided yet.

LoRa

Early on, I realized that LoRa might make a good choice for long distance sensor use. I did not need a high bandwidth here, just, creature detected, issue countermeasure. On twitter one day, Ronoth was retweeted wrt a Crowdsupply board. So, I waited sometime for that to arrive. It was my first use of Ada and LoRa. The history of how these boards came to be is outlined below.

First Ada + LoRa port: Ronoth S76S

This was a Crowdsupply board and is by Ronoth. The S76S is based on the Ascip S76S module.

The STM32L552 challenge

Whilst the other STM32 ports are not new ground for me wrt getting Ada going on ST platforms, the STM32L552 is exceptionally challenging. I have some experience with embedded platforms and the Cortex-M33 in its various forms from Nordic, NXP and now ST are truly some of the most complex controller designs I have ever worked with. The bugs you can get are really epic. I could fill pages here recounting all the really messy mind bending bugs. The bullet points below outline the work, but at each stage learning and bugs were involved.

1) It uses the ARMv8M architecture. No support in OpenOCD for this targetNo direct Ada support as the lIb is ARMv7M.

2) It is based on a Cortex-M33 with TrustZone.

3) Ada_Drivers_Library is designed for a single view of the peripheral space not a peripheral space partitioned into Secure and Non-secure areas.

4) Debugging also needs to consider ARMv8M+TrustZone and how that effects register/memory/flash reading and writing.

5) To that end openocd was ported to this platform:

https://github.com/morbos/openocd-0.10.1

From there you can attach to the board:

openocd -f board/st_nucleo_l552.cfg

6) The methodology for the port was to bolt two separate Ada ELF32 binaries into the final image. One image is a Secure boot and API handler, then a very small C glue layer that handles gcc ARMv8M stuff since gnat2019's gcc libs cannot be linked with ARMv8M code yet. Also some future pragma's would be needed in the Ada world to accommodate the special S <=> NS bindings + access to NS functions. (the BLXNS and SG instructions and cleaning up the unbanked register state before BLXNS to avoid S state leaks to the NS side). Finally the NS image is the last piece. For S I am using the Ravenscar sfp and for NS, Ravenscar full. Before NS can toggle an LED or touch any peripheral S has to be told what is allowed to be used by the NS side, extra boilerplate unneeded in a non-secure environment.

7) The basic structure of the flash is the boot area (formally 0x08000000, is now S at 0x0c000000 for a secure boot Ada Elf32 binary) it curr occupies 40-60k of flash. A watermark is created in the ST flash option regs to divide the flash into 2 regions, I chose S from 0x0c000000 - 0x0c01ffff) NS from 0x0802000 - end.

The secure boot area also has the veneer code to allow NS to call back into S. You need a magic SG instruction anywhere you want NS's PC to touch down, any other opcode is an abort. Also the region for the NS PC touch down must be marked Non Secure Callable (NSC), or, another cryptic abort. The S & NS Ada programs are Cortex-M4F compiler builds, the veneer code is in C, and it is compiled as Cortex-M33

Above we see the S_To_NS Ada call going to a C s_to_ns.

A magic function ptr with special attribute ensures that the function pointer produces the blxns shown plus a boatload of assembly to wipe out the CPU regs so that no leaks are present to the NS side. Note that the NS side is a full Cortex-M executable with vector table.

Then we have the NS side calling the LoRa radio IP's API. Let's look at Recv_SX1276 as an example API call

The call goes to the C wrapper recv_sx1276_from_ns. Notice that due to it's attribute, its veneer entry point starts with the ARMv8M instruction: sg an sg is the only instruction ns can execute upon arrival in s. Any other instruction and you will be in the debugger, debugging a crash.

After the sg, a veneer branch is made to a special veneer that finally calls our Ada code via an export of its Recv_SX1276 as C callable.

Upon arrival, we see pointer types. How do we on the secure side know that those pointers are safe (i.e. NS). We need to use another ARMv8M instruction to validate the pointer(s). The tt instruction as shown:

If any of those pointers are S, we return. This is defensive coding in a S/NS environment.

A really challenging port. For example 2 Ravenscar runtimes on one SoC, when an exception comes to/from secure/non-secure, how is context switch decided? This was very challenging as process starvation was a very real bug.

Detail:

There was no SW initially for this board, ordered in Nov, which came to my house early Dec. So I had 2 months to get the whole show going.

Secure/Non-secure peripheral base handling, this is interesting. How Ada sets up the peripheral bases for the driver code that shared between S/NS.

We see the magic in the declaration, the S base is always 16#1000_000# shifted from the NS base. In this way, legacy Ada_Drivers_Library code that just refers to GPIO_A for example will 'do the right thing' based on the stance that the library was built with (the Secure_Code constant).

Fifth port the STM32WB55 + LoRa. A LoRa server to BLE bridge

An Ada programmed STM32WB55 rounds out the show. Here is the server having its LoRa module traffic being analyzed by a Saleae.

This was also a big effort done after the last Make with Ada contest. Last contest was a SensorTile with Ada running the BLE stack. This WB55 is a collapsed SensorTile, the BLE radio is not an SPI peripheral here but has been absorbed into an SoC with HW inter-process communication being used instead of SPI. Lots of issues getting a port to this platform. (No SVD file initially(!)). Had a good bug with LoRa SPI. The SPI flags were at the wrong bit offset so no rx/tx notifications were received. This turned out to be an SVD file issue. Took a day to debug that one as an SVD file has a lot of leverage in a port, it will be the last place you look.

For the STM32WB55 nucleo board. There are 2 in the blister pack a large board and a small dongle. I am using the large board and a Blkbox 915MHz LoRa SX1276 module.

Detail:

In the US LoRa is restricted to 915Mhz. I thought this Ada code really elegantly solves the freq to 24bit coding the SX1276 uses:

And its usage:

The Saleae could capture BlkBox LoRa SPI radio traffic between the server on the STM32WB55 and the STM32L552 secure client. Very helpful. At that moment it was 2 Ada programs with a new protocol and issues with retries and list corruption. One bug was quite interesting, the server is in receive 99% of the time, you might think that setting its fifo to 0 and reading the received packet would be solid. No, if the SX1276 state machine never leaves receive, it keeps an internal fifo pointer. Since its internal, it keeps monotonically increasing every packet, even with an overt reset of rx&tx pointers. Thus at each packet notification the code sees stale data at 0 as the internal pointer has moved. The 'fix' is to leave receive just long enough for that pointer to reload. Another day of debug was that issue. There were many of these types of issues.

Putting it all together

OK, so we have 6 nodes of varying pedigree and 1 gateway/server. The design of LoRa packet data is a to/from pair of bytes (ala RadioHead FW) but then we deviate, namely some fields for commands and values for message retries and sequence numbers.

Message Protocol

There are 4 messages:

Client data structures are as so:

Node discovery

Every 5 seconds the server sends a broadcast ping, ping replies come in from:

2) Those nodes whose non-retried reply can get back to the server. A broadcast ping creates a lot of RF hash wrt the replies.

3) Once the server sees the reply, the node is added to a table:

This is how the network gets populated and grows as new nodes are discovered or old ones drop off.

Note that 0 and 255 are not valid client node IDs.

Once a node has been recorded by the server as active, it can begin to send and receive notifications. Today I only support 8bit notifications. 8 is more than enough for my target application. Ideally, I don't want long messages being sent. This 8 bit notify has been set so that bit0 is the LED on the board and bit1 is a user button. Only node#1 the STM32L552 has a usable user button. The plumbing of that interrupt was interesting btw, as its GPIO needs to be allowed to be read by the NS side, further, it can have its interrupt routed to NS. Again, all new ground and requiring deep study of the ref manual and board experimentation (no reference code, everything was the try, see, iterate method). Fortunately, I was using RAM for those tests or I think the flash would be worn out by now.

New Ada_Drivers_Library support for s/ns external interrupt routing:

The server's BLE stack is activated when the BLE phone application shows an LED icon press. Its a BLE notification. We then locally change the server's LED state to the requested value and then set a suspension object to let the LoRa task know that the LoRa network needs to be woken up with a notify8 message.

The server's notification task then walks the actives list and sends notify8's to all connected nodes. Upon receipt of a reply, the original message is removed from the queue. If however after a timeout, a node has failed to respond, then the original message with same seq# is resent. Finally, after 15 retries with no reply, the message is retired.

When the user button is pressed, the reverse happens. The button notification task on the client is activated via a suspension object, this then prepares a LoRa packet with a notify8 with bit2 set. Assuming it got over the air to the server after its own client retry count, then the server processes the notify8 by setting another suspension object that wakes up the same BLE task that handles the local button press. This then is transmitted to the BLE phone application where the event with timestamp is shown.

Demo

A demo has been coded up. The idea is that via ping packet replies, the server can build an array up of who is connected. Once that map is created the server knows what nodes can are alive on the LoRa network. From there via the Android app, we make a connection to LoRaDa (the name I gave the Ada BLE server on the STM32WB55). From there you can turn on a light and see any alert's. At this time, I only have a user button on the STM32L552 Nucleo board, so it is the only board that sends the alert.

The nodes in the LoRa network are as so:

2: Heltec LoRa Node 151 #1 (Ada coded LoRa client)

3: Heltec LoRa Node 151 #2 (Ada coded LoRa client)

4: Ronoth LoDev S76S #1 (Ada coded LoRa client)

5: Ronoth LoDev S76S #2 (Ada coded LoRa client)

6: Bluepill STM32F103

So when nodes arrive into the network, when the light button is toggled on the Android app, it cycles through all the connected nodes and sets each ones light to the state requested. All during this, node #1 can signal an alert that shows up as a red bell icon with a timestamp in the app.

Range

The range is outstanding, I walked prob 500m from our house and still had a signal! This tech is more than adequate for low rate sensor data.

I doubt I could have hacked this without using Ada. Once I had the client code up on the STM32L552 and was migrating the changes to the Heltec, Ronoth and bluepill boards, I don't think I spent more than 30mins getting the client changes up and working smoothly on those targets. I had already done a bare bones LoRa port to each but the client task version was not. So Ada made the job a pleasure and at least in the doldrums of code bugs I could know and rely on the fact the compiler was 100% there and solid despite my code problems.

Conclusion

One item for Ada users that can make LoRa quite inexpensive is a Bluepill (< $2) and a BlkBox LoRa module ($6-$7) so for <$10, you can have an Ada controlled LoRa module. A good value I feel.

Ultimately, ST will release a 48pin 7x7 STM32L5 processor, that will be pin compatible with the Bluepill boards. That will be Cortex-M33 based and might make for another interesting secure IoT solution. I plan to do a chip swap when its available as I did for the STM32L443 before.

About 2 weeks ago ST announced the STM32WL, that is a single chip LoRa controller. So, all the solutions I showed are dual chip, a controller + radio. Through some deal with Semtech, they will have a single chip with the radio IP absorbed. Of course, the Ada code I worked on will need to run on this device when its available.

Finally, whilst an Ada LoRa network and LoRa<->BLE gateway is novel, my feeling is the most interesting part of the work is the progress on the previously Ada untouched Cortex-M33. Ada is well known as a safe&secure language, the Cortex-M33 is a secure processor. So the marriage is a good one. Lets see if the Ada community can make some progress with this CPU, I already prepared a path and have shown its quite possible to get good results from it.

• Access the project code here.
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Relaxing the Data Initialization Policy of SPARK https://blog.adacore.com/relaxing-the-data-initialization-policy-of-spark Tue, 28 Jul 2020 08:36:53 -0400 Claire Dross https://blog.adacore.com/relaxing-the-data-initialization-policy-of-spark

SPARK always being under development, new language features make it in every release of the tool, be they previously unsupported Ada features (like access types) or SPARK specific developments. However, new features generally take a while to make it into actual user code. The feature I am going to present here is in my experience an exception, as it was used both internally and by external users before it made it into any actual release. It was designed to enhance the verification of data initialization, whose limitations have been a long standing issue in SPARK.

In the assurance ladder, data initialization is associated with the bronze level, that is, the easiest to reach through SPARK. Indeed, most of the time, the verification of correct data initialization is achieved automatically without much need for user annotations or code modifications. However, once in a while, users encounter cases where the tool cannot verify the correct initialization of some data in their program, even though it is correct. Until recently, there were no good solutions for this problem. No additional annotation efforts could help, and users had to either accept the check messages and verify proper initialization by other means, or perform unnecessary initialization to please the tool. This has changed in the most recent releases of SPARK (SPARK community 2020 and recent previews of SPARK Pro 21). In this post, I describe a new feature, called Relaxed_Initialization, designed to help in this situation.

First, let's get some insight on the problem. SPARK performs several analyses. Among them, flow analysis is used to follow the flow of information through variables in the program. It is fast and scales well, but it is not sensitive to values in the program. Said otherwise, it follows variable names through the control flow, but does not try to track their values.
The other main analysis is formal proof. It translates the program into logical formulas that are then verified by an automated solver. It is precise, as it models values of variables at every program point, but it is potentially slow and requires user annotations to summarize the effect of subprograms in contracts. Verifications done by flow analysis are in general easier to complete, and so are associated with the bronze level in the assurance ladder, whereas verifications done by proof require more user inputs and are associated with levels silver or higher.

In SPARK, data initialization is in general handled by flow analysis. Indeed, most of the time, it is enough to look at the control flow graph to decide whether something has been initialized or not. However, using flow analysis for verifying data initialization induces some limitations. Most notably:

• Arrays are handled as a whole, because flow analysis would need to track values to know which indexes have been written by a component assignment. As a result, SPARK is sometimes unable to verify code which initializes an array by part (using a loop for example, as opposed to a single assignment through an aggregate).
•  As it does not require user annotations for checking data initialization, SPARK enforces a strict data initialization policy at subprogram boundary. In a nutshell, all inputs should be entirely initialized on subprogram entry, and all outputs should be entirely initialized on subprogram return.

In recent releases of SPARK, it is possible to use proof instead of flow analysis to verify the correct initialization of data. This has the effect of increasing the precision of the analysis, at the cost of a slower verification process and an increased annotation effort. Since this is a trade-off, SPARK allows users to choose if they want to use flow analysis or proof in a fine grained manner on a per variable basis. By default, the lighter approach is preferred, and initialization checks are handled by flow analysis. To use proof instead, users should annotate their variables with the Relaxed_Initialization aspect.

To demonstrate how this can be used to lift previous limitations, let us look at an example. As stated above, arrays are treated as a whole by flow analysis. Since initializing an array using a loop is a regular occurrence, flow analysis has some heuristics to recognize the most common cases. However, this falls short as soon as the loop does not cover the whole range of the array, elements are initialized more than one at a time, or the array is read during the initialization. In particular, this last case occurs if we try to describe the behavior of a loop using a loop invariant. As an example, Add computes the element-wise addition of two arrays of natural numbers:

type Nat_Array is array (Positive range 1 .. 100) of Natural;

function Add (A, B : Nat_Array) return Nat_Array with
Pre => (for all E of A => E < 10000)
and then (for all E of B => E < 10000),
Post => (for all K in A'Range => Add'Result (K) = A (K) + B (K))
is
Res : Nat_Array;
begin
for I in A'Range loop
Res (I) := A (I) + B (I);
pragma Loop_Invariant
(for all K in 1 .. I => Res (K) = A (K) + B (K));
end loop;
return Res;
end Add;

The correct initialization of Res cannot be verified by flow analysis, because it cannot make sure that the invariant only reads initialized values. If we remove the invariant, then the initialization is verified, but of course the postcondition is not... Until now, the only solution to work around this problem was to add a (useless) initial value to Res using an aggregate. This was less than satisfactory... In recent versions of SPARK, I can instead specify that I want the initialization of Res to be verified by proof using the Relaxed_Initialization aspect:

Res : Nat_Array with Relaxed_Initialization;

With this additional invariant, my program is entirely verified. Note that, when Relaxed_Initialization is used, the bronze level of the assurance ladder is no longer enough to ensure the correct initialization of data. We now need to reach the silver level, which may require adding more contracts and doing more code refactoring.

Let's now consider the second major limitation of the classical handling of initialization in SPARK: the data initialization policy. As I have mentioned earlier, it requires that inputs and outputs of subprograms are entirely initialized at subprogram boundaries. As an example, I can consider the following piece of code which tries to read several natural numbers from a string using a Read_Natural procedure. It as an Error output which is used to signal errors occurring during the read:

type Error_Kind is (Empty_Input, Cannot_Read, No_Errors);

subtype Size_Range is Natural range 0 .. 100;

(Input    : String;
Result   : out Natural;
with Post => Num_Read <= Input'Length;

(Input  : String;
Buffer : out Nat_Array;
Size   : out Size_Range;
Error  : out Error_Kind)
is
Start    : Positive range Input'Range;

begin
--  If Input is empty, set the error code appropriately and return

if Input'Length = 0 then
Size := 0;
Error := Empty_Input;
return;
end if;

--  or we have reached the end of Buffer.

Start := Input'First;

for I in Buffer'Range loop

--  If nothing can be read from Input, set the error mode and return

Size := 0;
return;
end if;

--  We have reached the end of Input

if Start > Input'Last - Num_Read then
Size := I;
Error := No_Errors;
return;
end if;

end loop;

--  We have completely filled Buffer

Size := 100;
Error := No_Errors;
end Read;

This example is not following the data initialization policy of SPARK, as I don't initialize Buffer when returning with an error. In addition, if Input contains less than 100 numbers, Buffer will only be initialized up to Size. If I launch SPARK on this example, flow analysis complains, stating that it cannot ensure that Buffer is initialized at the end of Read. To silence it, I can add a dummy initialization for Buffer at the beginning, for example setting every element to 0. However this is not what I want. Indeed, not only might this initialization be costly, but callers of Read may forget to check the error status and read Buffer, and SPARK won't detect it. Instead, I want SPARK to know which parts of Buffer are meaningful after the call, and to check that those only are accessed by callers.

Here again, I can use the Relaxed_Initialization aspect to exempt Buffer from the data initialization policy of SPARK. To annotate a formal parameter, I need to supply the aspect on the subprogram and mention the formal as a parameter:

procedure Read
(Input  : String;
Buffer : out Nat_Array;
Size   : out Natural;
Error  : out Error_Kind)
with Relaxed_Initialization => Buffer;

Now my procedure is successfully verified by SPARK. Note that I have initialized Size even when the call completes with errors. Indeed, Ada says that copying an uninitialized scalar, for example when giving it as an actual parameter to a subprogram call, is a bounded error. So the Relaxed_Initialization aspect wouldn't help here, as I would still need to initialize Size on all paths before returning from Read.

Let's write some user code to see if everything works as expected. Use_Read reads up to 100 numbers from a string and prints them to the standard output:

procedure Use_Read (S : String) is
Buffer : Nat_Array;
Error  : Error_Kind;
Size   : Natural;
begin
for N of Buffer loop
end loop;
end Use_Read;

Here SPARK complains that Buffer might not be initialized on the call to Read. Indeed, as the local Buffer variable does not have the Relaxed_Initialization aspect set to True, SPARK attempts to verify that it is entirely initialized by the call. This is not what I want, so I annotate Buffer with Relaxed_Initialization:

Buffer : Nat_Array with Relaxed_Initialization;

Now, if I run SPARK again on my example, I have another failed initialization check, this time on the call to Put_Line inside my loop. This one is expected, as I do not check the error status after my call to read. So I now fix my code so that it only accesses indices of Buffer which have been initialized by my read:

procedure Use_Read (S : String) is
Buffer : Nat_Array with Relaxed_Initialization;
Error  : Error_Kind;
Size   : Natural;
begin
if Error = No_Errors then
for N of Buffer (1 .. Size) loop
end loop;
end if;
end Use_Read;

Unfortunately, it does not help, and the failed initialization check on the call to Put_Line remains the same. This is because I have not given any information about the initialization of Buffer in the contract of Read. With the usual data initialization policy of SPARK, nothing is needed, because SPARK enforces that all outputs are initialized after the call. However, since I have opted out of this policy for Buffer, I now need to use a postcondition to describe its initialization status after the call. This can be done easily using the 'Initialized attribute:

procedure Read
(Input  : String;
Buffer : out Nat_Array;
Size   : out Size_Range;
Error  : out Error_Kind)
with Relaxed_Initialization => Buffer,
Post => (if Error = No_Errors then Buffer (1 .. Size)'Initialized
else Size = 0);

The postcondition states that if no errors occurred, then Buffer has been initialized up to Size. If I want my code to prove, I also need to supply an invariant at the end of the loop inside read:

pragma Loop_Invariant (Buffer (1 .. I)'Initialized);

Now both Read and Use_Read are entirely proved, and if I tweak Use_Read to access a part of Buffer with no meaningful values, SPARK will produce a failed initialization check.

The Relaxed_Initialization aspect provides a way to opt out of the strict data initialization policy of SPARK and work around the inherent imprecision of flow analysis on value sensitive checks. It enables the verification of valid programs which used to be out of the scope of the proof technology offered by SPARK. You can find more information in the user guide. Don't hesitate to try it in your project, and tell us if you think it is useful and how we can improve it!

]]>

Story

Introduction

Miniature Circuit Breaker (MCB) interrupts mains power when a short circuit or over-current occurs. The purpose is safety of electrical system from fire hazard.

A smart circuit breaker will not only function as a regular MCB but also isolates incoming AC mains supply during disaster by sensing earthquake, fire/smoke, gas leakage or flood water. By disconnecting incoming power lines to equipment and power outlets inside house/office/industry during any disaster, it can reduce the chance of electrical hazard and ensuring safety of peoples life and assets.

This system is programmed with Ada, where safety and security is critical.

Hardware and Theory of Operation

Hardware modules

Following parts are connected together to assemble the hardware according to the schematic below :-

• Microbit: Runs safe firmware written in Ada for the system
• MAA8693 Accelerometer: Earthquake sensing, onboard I2C sensor
• 10 RGY LED Module: Fault Status indication, CC connection
• Buzzer: Fault Alarm beeping and tone generation
• TL431 External Reference: 2.5V reference for ADC measurement
• Laser & Photo Transistor: Smoke Sensing with light interruption
• MQ-5: Natural Gas (CnH2n+2) Leakage Sensor
• Flood Sensor: Electrode to detects presence of flood water
• Infrared Flame Sensor: Detects fire break out nearby
• TP4056LiPo Charger Module: Charges up backup battery
• Boost Module: Convert3.0-4.2 V from LiPo to 5.0V DC
• Protoboards: Substrate and interconnection between modules
• Power Supplies: LiPo Battery (backup) and 5V adapter (primary)
• MCB / Relay Module***: Connect/Disconnect mains
• Servo Motor: Trips MCB when smoke/fire/gas/vibration/water sensed
• B & A button: Acknowledge fault and resume normal operation

*** Note: Relay not used but can be used instead of MCB

Hardware Pin Map

All the GPIO, ADC, I2C pins are utilized as follows:-

Schematic

Here is the schematic for Smart Circuit Breaker hardware prototype :-

Device Operation

Device operates according to following flowchart :-

• In Ada code, all the I/Os associated with sensors, modules and indication LEDs are initialized first.
• Next, Smoke, Flame, Natural Gas, Earthquake, Flood sensing happens sequentially until a fault condition is detected.
• Immediately after any fault detection, MCB will be tripped by the servo motor.
• Then, LEDs associated with that of fault keeps blinking and buzzer keeps alarming.
• User needs to press button B to acknowledge fault after taking care of the situation/disaster that triggered the fault in the first place.
• Finally, user will flip the MCB manually to 'On' position and then press A to resume sensing again.
• If a short circuit or over current occurs, MCB will just trip like a regular MCB.

Build

On a piece of protoboard the battery and charger modules are connected and secured with double sided tape and hot glue. This is the bottom layer circuit for powering the rest of the components. Header pins are soldered to carry power to the next layer.

On the second layer (i.e the top one), rest of the sensors, modules and microbit is connected according to the schematic.

Servo motor is tied with regular circuit breaker with cable tie and connected to the top layer board to get power from the battery & control signal from Micro:bit.

Install all of these in the same directory/folder.

Where to start: GNAT Programming Studio

C:\GNAT\2019\share\gps\templates\microbit_example

Open one of the examples (i.e. digital_out.gpr) for Microbit according to following steps and edit example project code as needed.

• Step 1: Run GPS (GNAT Programming Studio)
• Step 2: Select digital_out.gpr example project

Step 3: Copy this project's code attached below and replace the example code in main.adb file

Following files are the most important files when working with GNAT Studio :-

.gpr file is the GNAT project file for a project

.adb file is the file where Ada code resides (src folder)

.ads files is where the definitions and declarations goes

Code snippets below are taken from the attached code of this project to briefly explain essential Ada programming styles :-

Comments/ non executable lines in Ada starts with " -- " like this :-

----------------- edge connector pin mapping ----------------------
-------------------------------------------------------------------
-- See here : https://makecode.microbit.org/device/pins -----------
--  pin(code)   pin (edge connector pads)    hardware connected
--   0         --  large pad 0       -- servo motor control  pin
--   1         --  large pad 1       -- Flame Sense IR module

Anything after -- in a single line is a comment, whereas regular syntax ends with semicolon (;)

Syntax with "with" keyword is used to add package support to a program. When 'use' keyword is used for that package, it becomes visible/usable in the code

with MicroBit.IOs; use MicroBit.IOs;     -- includes microbit GPIO package
with MicroBit.Time;                      -- includes microbit time package
with MicroBit.Buttons; use MicroBit.Buttons; -- includes button package
with MMA8653;   use MMA8653;          -- includes hal for accelerometer
with MicroBit.Accelerometer;          -- includes acceleratometer package

For example: "with MicroBit.IOs" includes microbit GPIO control support to main.adb code. But including "use MicroBit.IOs" will enable to use variable types from MicroBit.IOs package (see below: Variables in Ada for detailed explanation)

Similarly MMA8653 and MicroBit.Accelerometer enables support for onboard accelerometer chip microbit

Variables are declared in Ada in following format

• Variable_Name : Type := Initial_Value;
• Variable_Name : Type;

Connected is a variable name, which is Boolean Type, it's initial Value is True.

Fault_Flag is a variable name, which is Integer Type and it's initial Value is 0

Connected   : Boolean := True;           -- boolean type variable
Fault_Flag  : Integer := 0;              -- integer type variable
RedLED1_Smoke    : constant MicroBit.IOs.Pin_Id := 13;
RedLED2_Flame    : constant MicroBit.IOs.Pin_Id := 8;

Variable types are 'strict' in Ada.

For example: ADCVal is not 'Integer' type but 'MicroBit.IOs.Analog_Value' type, although it will hold integer numbers between 0 to 1023

Similarly, RedLED1_Smoke has a constant value of 13, but it is not 'Integer' type constant, it is actually 'MicroBit.IOs.Pin_Id' type constant.

To use these odd types of variable like : MicroBit.IOs.Analog_Value and MicroBit.IOs.Pin_Id. coder must include the 'use MicroBit.IOs; ' line of code before variable declaration.

'use' keyword allows programmer to use package specific types of variable.

Main procedure in Ada is the main function (equivalent of void main in c), which starts with 'procedure Main is' syntax, then comes the variable declaration. After that the 'begin' keyword begins the main procedure. Below 'begin' is the code which is usually initialization or single run code. Next starts the infinite 'loop' (equivalent of while(1) in c). Finally the 'end loop;' encloses the infinite loop and 'end Main;' ends the main procedure.

Here is the ada code skeleton with comments showing what goes where :-

-- package inclusion goes here

procedure Main is

-- variable declaration goes here

begin

-- initialization or one time executable code goes here

loop

-- body of recurring or looping code goes here

end loop;
end Main;

; (semicolon) is the end of a loop or procedure. There are no use for curly-braces {}

In Ada, if-else starts with 'if' keyword followed by logical condition and 'then' keyword, next is the code which will execute if the condition is true, otherwise the code below 'else' will execute. The 'if' statement ends with 'end if;' keyword

if condition_is_true then
-- do this
else
--    do that
end if;

Example :-

if ADCVal >= ADCtemp then
MicroBit.IOs.Set (RedLED1_Smoke, True); -- Write High to Disble LED
else
Fault := True; Fault_Flag := 1;
Connected := False;
end if;

for tempval in 0 .. 9 loop
MicroBit.IOs.Set(Servo_Pin,True);
MicroBit.Time.Delay_Ms(1);
MicroBit.IOs.Set(Servo_Pin,False);
MicroBit.Time.Delay_Ms(19);
end loop;

Case in Ada starts with 'case' keyword, followed by a variable which will be checked and 'is' keywords. Then it checks matching with 'when' keyword followed by different possible values of the variable and ends with '=>' operator. Next is the code which executes when variable check match with a possible value. 'when other' keyword is for no match condition. Case ends with 'end case;' keyword

'null;' is for doing nothing when no match is found, which needs to be explicitly mentioned.

Nothing is left for guess work in Ada !!!

case variable_name is
when 1 =>
-- do this
when 2 =>
-- do that
when others
null; -- do nothing
end case;

Example :-

case Fault_Flag is
when 1 =>
MicroBit.IOs.Set (RedLED1_Smoke, False);
MicroBit.Time.Delay_Ms (100);
MicroBit.IOs.Set (RedLED1_Smoke, True);
MicroBit.IOs.Set(Buzzer_pin,False);
MicroBit.Time.Delay_Ms (100);
when 2 =>
MicroBit.IOs.Set (RedLED2_Flame, False);
MicroBit.Time.Delay_Ms (100);
MicroBit.IOs.Set (RedLED2_Flame, True);
MicroBit.IOs.Set(Buzzer_pin,False);
MicroBit.Time.Delay_Ms (100);
when 3 =>
MicroBit.IOs.Set (RedLED3_NGas, False);
MicroBit.Time.Delay_Ms (100);
MicroBit.IOs.Set (RedLED3_NGas, True);
MicroBit.IOs.Set(Buzzer_pin,False);
MicroBit.Time.Delay_Ms (100);
when 4 =>
MicroBit.IOs.Set (YellowLED1_Quake, False);
MicroBit.Time.Delay_Ms (100);
MicroBit.IOs.Set (YellowLED1_Quake, True);
MicroBit.IOs.Set(Buzzer_pin,False);
MicroBit.Time.Delay_Ms (100);
when 5 =>
MicroBit.IOs.Set (YellowLED2_Flood, False);
MicroBit.Time.Delay_Ms (100);
MicroBit.IOs.Set (YellowLED2_Flood, True);
MicroBit.IOs.Set(Buzzer_pin,False);
MicroBit.Time.Delay_Ms (100);
when others =>
-- do nothing
null;
end case;

Microbit specific APIs

• MicroBit.Time.Delay_Ms (integer) -- delays operation for certain mili seconds
• MicroBit.IOs.Set(Pin_Number, boolean) -- Output drive a GPIO pin
• MicroBit.IOs.Analog(Pin_Number) -- returns a ADC value from an Analog pin
• MicroBit.Buttons.State (Button_Name) = Pressed -- reads A/B buttons

To use these Microbit specific APIs, following packages must be included first:

with MicroBit.IOs;     use MicroBit.IOs;
with MicroBit.Time;
with MicroBit.Buttons; use MicroBit.Buttons;
• use MicroBit.IOs enables the use of MicroBit.IOs.Analog_Value type
• use MicroBit.Buttons enables the use of Pressed type

Examples :-

with MicroBit.IOs;     use MicroBit.IOs;     -- includes microbit GPIO   lib
with MicroBit.Time;                          -- includes microbit timer  lib
with MicroBit.Buttons; use MicroBit.Buttons; -- includes ubit button A/B lib

MicroBit.Time.Delay_Ms(500); -- 500 mS delay

MicroBit.IOs.Set(2, True) -- sets pin 2 Logic-High
MicroBit.IOs.Set(1, False) -- sets pin 1 Logic-Low

ADCVal : MicroBit.IOs.Analog_Value; -- analog_value type variable,not an int
ADCVal:= MicroBit.IOs.Analog(0) -- returns a value between 0 to 1023

MicroBit.Buttons.State (Button_A) = Pressed -- returns True is A is pressed

Once the editing of the code is done, connect Microbit to computer with USB cable ( Windows will make ding-dong sound ).

Then click : Build > Bareboard > Flash to Board > main.adbto flash code to Microbit. The Message window below will show code size and upload percent.

If upload problem occurs, check USB cable or reinstall pyOCD.

Ada isn't just another programming language. It shines where Safety, Security and Reliability matters. In systems where a hidden firmware/software bug could be fatal, life-threatening or damage of equipment might cause huge economic loss, those are the kind of systems where Ada can make a huge difference.

For example, embedded system used in:-

• Pacemaker & ICU Medical Equipment
• Self Driving Vehicles
• Explosive Igniter
• Missile Guidance & Para-suite Launcher
• Spaceship Life Support System
• Lift Control
• Fire Alarm & Safety
• Automated Security
• Enterprise Server Power Monitoring
• Fail Safe Mechanism Monitoring
• Power Plant Steam Generation
• Radioactivity Monitoring
• Chemical Process Control
• Safety Critical Consumer Electronics (e.g. Induction Cooker)

How Ada makes system safe and secure ?

Ada compiler is very strict, it will keep bashing the coder/programmer with errors, warning, suggestions until a clear, well thought code is produced.

Well, compilers in other programming languages do that, too ! But the difference is, things that are not even an error in other programming language is an error in Ada. Someone coming from C or Arduino Land will feel the punch. For example - when trying to add float with integer. In Ada, Apple does not add up with Banana.

“think first, code later” - is the principle which Ada promotes !

Programmer must think clearly about the impact of each type/variable and code in a proper manner. There are other differences like writing style, operators.

Practical Design Considerations

This prototype is designed in a way, so that all the functions can be demonstrated easily. But for practical use, following actions are recommended:

• Both Smoke and Flame sensor are sensitive to strong light, therefore proper shielding from direct light is recommended
• Flood sensor should be placed near floor, where it can easily detect indoor flood water
• Earthquake sensor is susceptible to vibrations, that is why Smart MCB should be mounted on rigid structure
• Gas sensor requires 24 hrs break in period for proper operation
• Proper PCB and enclosure is necessary for hardware reliability

Conclusion

As I have already said, this is just a prototype hardware which I made with my limited resources. But Smart MCB is exactly the kind of application (safety critical) for which the spirit of Ada programming is intended.

MCB was invented and patented by Hugo Stotz almost 100 years ago. I wish someone out there can turn this project into a real product and upgrade the century old MCB technology into Smart MCB for improved safety of next generation electrical distribution systems.

• Access the project code here.
]]>

Story

The project sources

As junior DevOps/SRE, we're quite interested in cryptography and its usage. Therefore, when we were asked to participate to this contest for a uni project, we decided to go for some cryptography-related subject.

For the quite small knowledge we had on the subject, we thought that having a software on the embedded device that could generate RSA keys would be a good start.

The first steps - Bignums

In order to generate RSA private keys, we need a way to obtain big, very big prime numbers. Those numbers we need are way too big for classical number representation in ADA. We need a way to manipulate Bignums, aka number that can have hundreds of number of digits. Typical RSA keys as in 2020 contains 2048 or 4096 bits, so if we want to create such keys, we need bignums that can handle at least 4096 digits.

For the project, we first looked at a library to handle without dependencies bignum, but this library was too limited and too slow. Therefore, we created a bignum library, fitting our needs (allocation on the stack, fast operations, base 256 for better performances, handling negative numbers, ...). We spent quite some time optimizing it, since numbers computations will be most of the CPU time used to generate a key.

Pseudo-prime numbers

The algorithm to find a prime number is quite easy to understand, and leave very few place for optimization. To assert that n is prime, we look at (x * x) mod n, with x going from 2 to sqrt(n).
This operation is slow, and not suitable for 2048 bits numbers. To have an efficient way to assert that a number is prime, we switch to checking it to be pseudo-prime, with a sufficient probability.
A pseudo-prime number means that this number has a probability n of being prime, and functions generating pseudo-prime number must allow us to choose this n value.

We tweaked a bit, and found the best way to generate prime number with a probability > 99.999 % in a reasonable time. The algorithm is the following:
- Check for it being prime with all the prime numbers < 100 using a ~sieve of Eratosthenes
- Check for it being pseudo-prime using Fermat's primality test with {2, 3, 5, 7}
- Run a Miller-Rabin test with 4 iterations using randomly generated witnesses
- Run a Miller-Rabin test with a number of iterations depending on the number being checked with randomly generated witnesses
When a number passes all of this tests, it is safe (up to a defined probability) to consider it being prime.

Random number generator

Random numbers are required to run such algorithms. But random must respect some conditions. We're not only looking for a Pseudo Random Number Generator (PRNG), but for a PRNG that either mix entropy or is Cryptographically Secure (CSPRNG). As we're on an embedded device, we have GPIO pins and sensors, so we chose the first option as entropy could be generated easily.

We implemented a PRNG fed with entropy, inspired by the Linux PRNG (/dev/random). We have an entropy pool of 2048 bits, that is constantly fed by the noise generated by the 3-axis accelerometer using a mixing function. We also have an extraction function, that consume some of the entropy available in the pool (accounted by a estimation function that maintains a gauge every time an operation is performed on the entropy pool) to give a random number.

In our application, entropy is collected periodically in the background, since we created an ADA Task to perform this operation.

The implementation uses chacha20 as a hash function for the extraction process, which can be resumed as:
- Hash the whole entropy pool
- Fold the hash in a 16 bytes hash
- Mix the hash back in the pool to mix-it and re-credit some entropy
- Re-extract 16 bytes of entropy from the pool
- Xor and fold the initial hash and the entropy newly extracted to create a 64 bits output value

The entropy pool can also be fed by any source of entropy, and we also looked on nuclear decay. Unfortunately, our Geiger-Müller counter was shipped with a lot of delay, and we lack the time to interface it.

RSA

With prime numbers and a random number generator, it is finally possible to create RSA keys. The project offers a graphical interface on the touch screen of the board to choose the RSA key size between 256, 512 and 768, which are the biggest key that can be generated in a reasonable time on such low-specs board. The touch screen also permits the user to generate a key, and to print it on the USART connection. The project aims to be a good start for either an embedded crypto library, or for a device like a smartcard (Yubikey).

A RSA key is dumped on the USART using the ASN.1 format, but as a conf file. Here's an example of a very small RSA key generated by the board:

asn1=SEQUENCE:rsa_key

[rsa_key]
version=INTEGER:0
modulus=INTEGER:187031899687190461
pubExp=INTEGER:65537
privExp=INTEGER:113679732755818241
p=INTEGER:191569309
q=INTEGER:976314529
e1=INTEGER:45208217
e2=INTEGER:554591105
coeff=INTEGER:98718018

With openssl, the configuration above can easily be converted to the DER/PEM format, making the key suitable for most applications.

• Access the project code here.
]]>

Story

My wife suffers from a rare condition that causes her to be nauseous for many hours quite often. Through trial and error, she has learned that one thing that seems to alleviate the symptoms is sitting upright and perfectly still. I noticed that often, early in the morning, she would slip out of bed and move to the couch where she can sit upright while I slept in bed.

So, about a year ago, after too many nights of being uncomfortable with our bed, my wife and I decided to finally go bed shopping for a good mattress. In doing so we realized how common it is now that people purchase adjustable bases with their mattresses. It dawned on us that having an adjustable base would be excellent for her condition as she could sit up in bed whenever she needed and she wouldn’t have to go to the couch to do it. After trying one on a showroom floor we were excited to find that we loved it and we went out and bought one with our new mattress.

At first, it was amazing; being able to prop yourself up in bed. But then there were problems. When my wife wanted to raise the bed, if I wasn’t there she’d have to fumble around for the remote (which always seemed to be hard to find). Often, if she wasn’t feeling well, she’d give up and acquiesce to laying supine. Another problem we had was that we tended to fall asleep with the bed in the upright position. This gets quite uncomfortable for the entire night and the effort needed to find the remote and adjust the bed would wake us from our sleep, making it hard to get back to sleep.

One night, when faced with this problem it dawned on me: we already use our Alexa to control our lights and other devices, why not our bed? A quick Google search revealed that nothing like this had ever been done. To make it worse, adjustable beds are primitive affairs --- they don’t have APIs and they aren’t programmable. They rely on power control circuits to drive them.

Project Goals and Overview

The goals of this project were to create an IoT device that was:

• Able to control the movement of a wired adjustable bed.
• Able to be easily reconfigured to work with other wired adjustable beds.
• Able to be controlled by both an Amazon Alexa device as well as the original remote.
• Able to sense occupancy in the room and able to produce under bed safety lighting when walking around at night.
• Nicely designed in terms of fit and finish (since I planned on actually using it!). That means that the entire system should use its own custom PCBs and be hosted inside of a custom-designed case that would look great under my bed.

To explain the product, I’ve prepared two videos that demonstrate the main features of the SmartBase as well as showcase its physical construction.

For the interested reader, you can read the remainder of this document to learn all about the different phases of development for this project. Roughly, the phases of development this project followed were:

1. Reverse engineering, in which I figure out how to control the bed.

2. Prototyping, in which I made some rough perf-board prototypes.

3. PCB design, because no one wants a rats nest of wires under their bed, I designed a series of PCBs to support the function of the SmartBase.

4. Enclosure Design, in which I designed a case for my project.

5. Software Design, which essentially happened at all phases, but I’ve given it its own section here.

6. Verifying Things, in which I do a little bit of verification.

7. Moving to STM32, in which I describe porting the entire project over to a STM32F429ZIT6 and a ESP32 for the WIFI functionality.

8. Bitbanging a WS2812B on a STM32F4 in Ada. Neopixels are only easy to use if you are using an Arduino. Apparently no one had done this one before, so, armed with an oscilloscope and a ARM manual, I figured out how to make that happen.

9. Wrapup, What did I think about using Ada to build the SmartBase?

For the complete code of the SmartBase running on the RPIZeroW, please visit: https://github.com/jsinglet/sm...
For the complete code of the STM32 version of the SmartBase please see this repository: https://github.com/jsinglet/smartbase-stm -- This repo is essentially a fork of the above repository which targets the STM32. I've published them separately so it's easy to understand which repository represents which product.

However, before we jump into all that, here are some pictures of the finished project in action. I hope you enjoy reading about it as much as I enjoyed making it!

Phase 1: Reverse Engineering

The first problem I had to solve was: how do I control the adjustable base? As I’ve mentioned, I’m a software guy, but I’m dangerous with a multimeter. To figure this out, I simply started by sliding under my bed and looking at how it currently functioned.

The configuration of my bed was that of a wired remote, connected with a DIN5-style (big round connector commonly found in MIDI applications). Again, being a software guy, I assumed that perhaps the way the people who built the bed constructed it was to perhaps to implement some sort of protocol over serial. I connected it to my computer, fired up WireShark, but alas, nothing sensible came out of the remote.

So of course, my next step was to disassemble the remote. The wired remote, horrifically disassembled, is pictured below.

After some experimentation with a multimeter I was able to determine that the circuit was quite simple. Here’s a picture of a whiteboard from around the time I was figuring it all out:

The logic for the circuit is actually quite simple. One of the 5 pins functions as the power, which happens to be +30V. The other pins are connected to the power via a switch and fed back into the bed. I tested the idea by taking hookup wire and touching it across the various pins in the configuration I had determined through my schematic analysis and sure enough it worked. With this information, I was ready to construct a rough physical prototype.

Phase 2: The Super Rough Prototype

So after learning the basics of how the control circuit worked, I wanted to test if this could be controlled via relays so after a few trips to my local electronics hacking store SkyCraft (a really one of a kind place!) and some Amazon purchases I managed to put together a rough perfboard prototype, which sadly I didn’t take a picture of when it was all together – however, I saved the board, which is pictured connected to a off the shelf relay module, below:

Phase 3: Moving from Perfboard to PCB

One of my goals for this project was to have a tidy little box I could stick under my bed that my wife would tolerate. The early perfboard prototype sat on top of some roughly cut tile with a lot of duct tape. I don't have a picture, but it was a great sight. Horrifying, but a good sight.

If I was going to have a permanent fixture, I had to fix the ratsnest problem, starting with the circuit. For this, I decided to start by designing my own custom PCB to control the bed. The topic of how many wrong turns and iterations of the PCB I did could frankly fill many pages. To shorten that, I'll just present my finished schematics, below. Before that, however, let me just say that I used JLCPCB to do all my fabrication, including SMT part placement. I can't say enough good things about this company. They are fast, cost effective, and the boards are flawless. For my PCB layout I used Autodesk Eagle, which was great as well.

To gain a better understanding of the hardware required to run the SmartBase, I've prepared the following Block Definition Diagram which details the main systems the SmartBase relies on to provide its functionality.

At the top level we have the Relay and Power Control Module and the LED Mainboard Module. These two components roughly outline the two boards that were created for this project. In the following paragraphs we discuss each board separately.

First, let's look at the board on the bottom, which was responsible for power and hosting the relays on the board that control the bed.

Power is routed in through this board on the bottom through a barrel connector to all of the components within the SmartBase. Other components hosted on this board are:

In order to control the relays (safely) with GPIO signals I used a series of FQP30N06L "logic level" MOSFET transistors connected as shown in the above schematic.

One cool thing I did was make the configuration of the relays jumpered, so if I ever encountered a different bed with a different control configuration they relays can be easily rerouted. The basic idea is that the jumper positions control which pin gets the power signal when the relay is activated.

Next, this is the top board, which hosts the WS2812B LED array as well as the CPU, which is a Raspberry Pi Zero W. The schematic and the layout file for the board is pictured, below:

This board hosts two 40 connectors which host both the GPIO cable connector from the bottom board and the header for the Raspberry Pi itself. If you can believe it, I got this board right on the first try just by following manufacturer spec sheets.

Phase 4: Enclosure Design

To build the enclosure I opted for a very simple design, which features several components:

• Two DIN-5 ports on the back. One connects to the bed and the other allows you to reconnect the remote control to the bed if you wish to ever manually control the bed. In reality I never used this feature because the voice control was so good.
• One power connector for 5VDC power.
• The case actually splits into 3 different bodies. The bottom body, the top body, and the "lens" which I printed out of clear PETG. The rest of the case was printed in black PLA.

One challenging aspect of this case design was actually coming up with appropriate screw sizes. I ended up using a combination of M2 and M3 screws. I found a bunch of assorted hex head sizes on amazon and was able to design the case dimensions around those.

Phase 5: Software Design

In this section we will discuss the design of the Ada software that runs the SmartBase. The software described in this section is the software that drives the version of the SmartBase shown in the demo. Later in the section on STM32 we discuss how these components are handled on a STM32F4-family processor in Ada. However, before we start, a few notes:

• SmartBase makes use of tasking. It is in fact mainly composed of 5 core tasks that handle relay control, command interfacing, MQTT command processing, and LED status control.
• Because I wanted to be able to verify things (and the run it on metal later), I enabled the Ravenscar profile.
• Portions of the application are written in Spark. Notably the components that control the bed are written in Spark with some lightweight specifications around the control sequences.

Concept of Operation

The SmartBase gets inputs from PIR sensors which will trigger fade on / fade off events. These events are able to be processed along with MQTT events which arrive via a AWS Lambda function connected to the Alexa voice service. These MQTT events then turn into motion of the bed via the relay control subsystem. The following diagram provides a high level summary of how the SmartBase performs its main operations.

The typical way one builds microcontroller applications is via a state machine pattern encoded into the main loop of the program running on the microcontroller.

The SmartBase uses 5 tasks for performing its core operations:

2. The CLI Task, which provides a debugging command-line based interface to the SmartBase.

3. The MQTT Task, which listens for protocol events from Amazon IoT (from spoken voice commands) and talks to the bed task to execute protocol events.

4. The LED Task, which is responsible for providing a structured interface to controlling the LED ring. The LED ring defines states for connecting, connected, fading on, and fading off.

5. The Motion Detector Task, which is really comprised of several tasks and is easily the most complex of the 5. I describe the Motion Detector tasks more detail later in this section.

The following diagram details the relationship of these five tasks in more detail. Note that in the diagram I use the method loop to indicate the main loop of the task. One stipulation of Ravenscar is that tasks do not exit and notation calls that restriction out.

The design of the system is that all interaction with the LED and Bed components happens strictly through the Commands interface with the exception of the Motion Detector tasks. This interface itself in turn only acts with protected objects. For example, if a command arrives via the command line and the MQTT task at the same time (assuming we have more than one CPU) they will both attempt to process the command through the Commands interface which in turn will ensure the access to the resources is serialized.

One interesting task is the LED_Status_Task, which is responsible for processing changes to the LED status ring. There are two problems solved in this component: 1) How to provided serialized access to the underling LED hardware and 2) How to ensure that transitions to different LED states are valid? The first problem is solved through protected objects. The second part of the program is covered in more detail in the next section on verification.

Lastly, the most complicated use of tasking is easily in the way the SmartBase does motion detection. As can be seen in the above figure, the MotionDetector package is composed of two Tasks and two protected objects. They function in the following fashion:

• Interrupts arrive on the protected object Detector, which very quickly sets an interrupt flag.
• The Motion_Detector_Trigger_Task monitors this flag by waiting on the entry to Triggered_Entry. Once the barrier releases, the Motion_Detector_Trigger_Task engages the Timer_Task via the MotionDetector_Control object's Start method.
• The Timer_Task is then responsible for changing the status of the LED ring to OFF once the detection is finished.
• Re-triggering is handled at the level of the Motion_Detector_Trigger_Task. If the LED hasn't already faded off, more time is simply added to the timeout. That way, if people keep moving the lights remain on.

Other Software Details

Some other items worth mentioning pertaining to the Pi Zero version of this implementation are the following:

• For MQTT I wrote C bindings to the Eclipse Paho library. On the STM32 platform, I use the MQTT AT interface directly built into the ESP32 and program it over UART.
• For LED Control on the Pi I used a neopixel library and bound to it in C. On STM32 there is no such type of library for Ada (or anything else, really) so I wrote my own hand-coded bit-banged version of a WS2812B control library. What is nice about my implementation is that it is optimized for my particular use case and uses constant RAM (whereas other implementations use RAM on the order of the number of pixels in the array). You can read about it in the section on STM32.
• I did all my work in GPS, but I would really like to get my hands on the Eclipse version of the GNAT tool set and would happily accept any complimentary licenses!

Phase 6: Verifying Things

One of my goals was to check out the specification-related features of Ada with this project. To that end, I came up with two small verification tasks for my project.

1. First, rather than use a timer, part of the way that the bed is controlled is through the use of tasks and protected bodies. These tasks use a barrier to control when a task should start waiting to see if it should stop moving the bed. One aspect of this protocol is that since other commands may be received while the bed is moving (thus, nullifying the current action), the tasks controlling the timeout have to know that they have been cancelled without having race conditions around the starting and stopping of the bed. I will show a little bit of what I did in the protocol with relation to specification.

2. Second, a critical element of this application is the LED status ring. In this application the LED status ring is used for indicating when the SmartBase is connecting to the internet, disconnected, and when motion is detected. Designing a system that can process all of these states at any time is trickier than it sounds and I discuss the model I used for managing the LED status ring.

Specifications on the Bed Controller

The first thing I wanted to write some specifications for were the behaviors surrounding the behavior of the stopping and starting of the bed. As I described earlier. To do this I wrote the following specifications, which I will explain after the listing.

procedure Do_Stop_At (Pin : in out Pin_Type;
Expiration_Slot : Time;
Actual_Expiration_Slot : Time)
with
Contract_Cases => (
-- the slots match. This means
-- we will be performing the action
-- on the pin we expected.
(Actual_Expiration_Slot = Expiration_Slot) => Pin'Old = Pin,
-- the slots DON'T match, which means we missed our window
(Actual_Expiration_Slot /= Expiration_Slot) => Pin = Pin_None
);

procedure Stop_At (Pin : in out Pin_Type; Expiration_Slot : Time);

procedure Stop
with
Global => (Input => (Device.GPIO),
Output => (Bed_State_Ghost.Moving)),
Pre  => True,
Post => Bed_State_Ghost.Moving = False;

In the first specification, there are two cases. The first case is when the time slot that the timer task used to cancel the task is the one currently executing. In this case, we expect that the pre-state value of the Pin matches the post-state value of the pin, that is, we actually perform the stop on the pin we expected to. In the second case we require that if they are not the same that no pin is used to do anything. This is represented by the expression Pin = Pin_None.

The second set of specs are on the Stop and Start methods. These specs simply require that, in the case of Stop, we actually stop the bed, and in the case of Start, we assign a pin when the move was successful.

In the above listing, you might note that I am using an explicit Ghost package to hold ghost state. Why am I doing this when Ada/Spark 2012 has this feature built in? This is to get around an incompatibility I was having getting this to run on the Raspberry Pi, which only had an older version of GNAT available (and didn't support ghost fields). I was able to replicate this behavior by just encoding the ghost state into an actual package. Since ghosts are really just syntactic sugar for that it works quite nicely.

Design of LED Status Control Ring State Machine

Of course, verifying things doesn't always mean you write specifications. There are lots of ways to add more assurance around your design. The next thing I decided to analyze was the LED status ring, the states for which can be seen in the following diagram.

In the above diagram we have all of the states that the LED ring can be in. One thing that was important to me was ensuring that, the following properties would hold:

• When the LED faded ON, the system would not attempt to fade it back on. It sounds like a simple thing, but I didn't want to create a disco effect.
• I wanted to ensure that no matter what state the system was in, if the connection was every lost the system would, as soon as it could, notify the user. This point is subtle. I didn't want the system to interrupt the visual effect of a fade, however I did want the connection sequence to begin as soon as possible. Ensuring this sort of separation was critical to my design.
• Once the system was trying to connect, the visual feedback that the system was connecting should not be interrupted, even if motion detection events are happening. In my design, one doesn't have to disable motion detection -- the state machine of the LED ring simply subsumes this logic and makes such interruptions impossible.

What's nice about having a model like the one pictured above is that we can simply look at it and see if the desired properties hold. So let's do that one by one.

The first property obviously holds. Specifically, the only edges going into Fade On come from None, which is not reachable for any state after Fade on. So far so good.

The second property holds because all states in the state machine may make a transition to Connecting after completing their state transition.

The third property holds because for all nodes in the graph there are no transitions from the Connecting state except Connected.

Phase 7: Moving to the STM32

As I mentioned in the introduction, after completing this project on the Raspberry Pi Zero W platform I immediately began work on a version that could run on bare metal on a STM32 processor.

For my development boards I selected:

I'm not totally done getting the SmartBase running on the STM32 platform, but here's a quick rundown of what is and isn't done so far:

• The Bed Control is done and works perfectly with the CLI interface over a serial console. You can see that demonstrated in the video, above.
• The Motion Detection is done and perfectly works with the LED array.
• The LED Controls are done, thanks to a driver for the WS2812B I wrote which I describe in detail in Section 8.
• The MQTT work is almost done. To do this I had to build a custom ESP32 firmware to enable the MQTT AT command set on my board. This works perfectly through the serial console and I'm able to get it to pull down MQTT messages. I'm currently writing the driver that will send the AT commands over UART to the ESP32.
• PCB design hasn't been redone for the ESP32 + STM32 combo. However, I'm really excited about getting these guys onto my PCBs and I've already looked over the reference designs.

Phase 8: Bit Banging a WS2812B on a STM32F4 in Ada

One of the most interesting parts of this project was when I had to get the LED array working on the STM platform in Ada. The timings on the WS2812B are relatively tight and unlike platforms like Arduino and others where there is a wealth of information available, on the STM32 basically nothing exists. What little does exist could never be used in an Ada program because it is hopelessly coupled to the hardware abstraction layers and drivers provided by ST Microelectronics. Therefore, to do this I had to start from from principles and work my way forward.

From the manufacturer specification sheets, the WS2812B implements the following protocol:

To trick to producing colors with the WS2812B is to set the rightmost 24 bits of a number to the GRB value you want to set. For example, to set a LED to green, you would send a waveform corresponding to the following number:

0000 0000 1111 11111 0000 0000 0000 0000
(white)      (green)          (red)          (blue)

To control multiple pixels you just repeat this for as many pixels as you have, making sure to say within the Treset window, which practically just means you do it as fast as possible.

To get reliable timings the codes that drive the WS2812B must be written in assembly with interrupts disabled. As a proof of concept I decided to build up a basic loop to set a single LED to green. To do this I cross referenced the cycle cost for each of the instructions, which can be found on the ARM website

As for how many instructions are required, the calculations were performed as follows:

The STM32F429ZIT6 operates at 180Mhz, i.e. 180000000 CPS, which means each cycle costs: 0.00555555556 microseconds (5.555555559999999 ns). To achieve 800ns we need 800/5.555555559999999 = 143.9 instructions (i.e., 144) of delay time added to the pipeline. This works out to 144 * 5.55ns = 799.2 ns, which is well within the +/- 150ns needed. Similarly, to achieve 450 ns, we need 450/5.555555559999999 = 81 instructions of delay added. That gives us 81 * 5.55ns = 449.55, which is well within +/- 150ns needed.

The trick to implementing this is to work in the time to load and set the bits high (or low) and combine that with a loop that keeps that timing in check.

From reading other codes (such as the official neopixel code), one way people have done this is by adding explicitly chains of NOP instructions to the pipeline. This works well on a 16Mhz processor like the Arduino's AVR. It doesn't work as well on a 180Mhz processor. Because I didn't want pages and pages of NOPs pasted into my program, I decided instead to work out the code using a loop structure.

The following code works essentially according to the following procedure:

1. Load up a counter that dictates the number of pixels we want to do this for.

2. Load up a number that indicates which GRB value to load.

3. Loop over each bit of the color and, following the convention above, send the appropriate bit by introducing enough cycle delay to get the required timeouts.

Initially to create this program I worked it out using the math shown above for a single color value. This worked well for a small proof of concept code, but it was difficult to scale to the entire GRB value (let alone multiple ones). To help me expand the program I used an ARM simulator that could count cycles to ensure the loops I had written were correct. I used the VisUAL simulator for this purpose.

Finally, however, I wanted to be absolutely sure that my timings were correct. I actually didn't own an oscilloscope so I finally broke down and ordered one off of Amazon. I ended up getting a Hantek DSO5072P, which is a reasonably good scope for my purposes, if not a fancy one. The image below shows an example cursor session, measuring the delay of one bit.

The following code shows the general principle of operation -- if you'd like to see how I implemented it in the actual Ada code, please see the end of this post for the full code. One interesting thing to note here is that most neopixel implementations consume ram that grows with the number of pixels. Because I don't ever want any pixel to be a different color than another one I was able to make a simplification that allows my program to always require constant memory, regardless of the number of pixels being driven.

For now I'm just using this library in my own project, however, if others find it useful I'd consider contributing a library that makes STM32 control of neopixels accessible to all.

cpsie	i
init
mov		r3, #2              ; the number of pixels to do this for
ldr		r5, =0x40021418     ; set r5 to the GPIOF register + 18 offset for BSRR
ldr		r6, =0x2000         ; pin 13 HIGH mask
ldr		r9, =0x20000000     ; pin 13 LOW mask
mov		r8, #1              ; we use this to test bits
send_pixel
cmp		r3, #0        ; test if we are done
beq		done          ; if we are out of pixels, finish up
mov		r4, #23       ; we are going to send 24 bits, prime it here.
sub		r3, r3, #1    ; decrement this pixel
send_bit
lsl		r2, r8, r4    ; build the mask by shifting over the number of bits we have
bne		send_one      ; send a one
b		send_zero     ; otherwise, send a zero
;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;
send_one
str		r6, [r5]            ; set pin 13 HIGH
;;		delay for ~ 800ns
mov		r0, #36
delay_T1H
subs		r0, r0, #1
bne		delay_T1H
;;		end delay
str		r9, [r5]            ; set pin 13 LOW
;;		delay for ~ 450ns
mov		r0, #20
delay_T1L
subs		r0, r0, #1
bne		delay_T1L
;;		end delay
b		bit_sent
;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;
send_zero
str		r6, [r5]            ; set pin 13 HIGH
;;		delay for ~ 400ns
mov		r0, #17
delay_T0H
subs		r0, r0, #1
bne		delay_T0H
;;		end delay
str		r9, [r5]            ; set pin 13 LOW
;;		delay for ~ 850ns
mov		r0, #38
delay_T0L
subs		r0, r0, #1
bne		delay_T0L
;;		end delay
b		bit_sent
;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;;
bit_sent
cmp		r4, #0       ; was that the last bit?
sub		r4, r4, #1   ; otherwise, decrement our counter
beq		send_pixel   ; if so, go to the next pixel
b		send_bit     ; and send the next bit
done
cpie	i

Again, if you want to see what this all looks like when it makes it back to Ada, check out the Code Section of this project.

Wrapup

In this article I detailed the creation of a novel IoT device, the SmartBase. I described how it works, and described in detail each phase of its development.

So after all of this, what is my opinion of developing this product in Ada?

Frankly, my reaction is that I cannot even imagine doing it in another language. Even though I am quite familiar with so-called "exotic" languages like Haskell, the blunt, unapologetic efficiency of Ada, the pickiness of the compiler, the the robust, built-in support for tasking and specification were a pleasure to work with and I'm confident these features helped me to find many errors that would have otherwise potentially caused problems in my system. Some other items, in no particular order are:

• Scenarios are brilliant. Using scenarios I was able to make multiple versions of the code for SmartBase for different boards and tie it all together quite easily.
• Compared to Java and other languages, the package system of Ada, in general, makes structuring a complex application much better. I love how I can have package initialization, nest objects, and in general encapsulate functionality. I think that many people learn the OO facilities in a language like Java and walk away thinking "this is what objects are all about." The package system reminded me of the one in OCaml, which I also quite like.
• I love love love the support for separate complication. When you combine that with the package system, you have a powerful mechanism for build management.
• Having a SPARK > Prove All menu in my IDE is a beautiful thing to see.

In the future I plan to use Ada on a few more projects, for example, I wanted to make some small, battery powered WIFI LED strips. I can't think of a better language for the job.

• Access the project code here.

]]>
CuBit: A General-Purpose Operating System in SPARK/Ada https://blog.adacore.com/cubit-a-general-purpose-operating-system-in-spark-ada Wed, 10 Jun 2020 06:10:00 -0400 Jon Andrew https://blog.adacore.com/cubit-a-general-purpose-operating-system-in-spark-ada pragma Suppress (Index_Check); pragma Suppress (Range_Check); pragma Suppress (Overflow_Check); ... pragma Restrictions (No_Floating_Point);
package Compiler is
(
...
"-mno-sse"
"-mno-sse2"
...
);
end Compiler;
---------------------------------------------------------------------------
-- Read from a model-specific register (MSR)
---------------------------------------------------------------------------
function rdmsr(msraddr : in MSR) return Unsigned_64
is
low : Unsigned_32;
high : Unsigned_32;
begin
Asm("rdmsr",
Outputs => (Unsigned_32'Asm_Output("=a", low),
Unsigned_32'Asm_Output("=d", high)),
Volatile => True);
return (Shift_Left(Unsigned_64(high), 32) or Unsigned_64(low));
end rdmsr;
/* AP starting point */
AP_START        = 0x7000;

KERNEL_PHYS     = 0x00100000;
KERNEL_BASE     = 0xFFFFFFFF80000000;

...
SECTIONS
{
. = AP_START;

.text_ap : AT(AP_START) {
stext_ap = .;
*(.text_ap_entry)
etext_ap = .;
}

. = KERNEL_PHYS + KERNEL_BASE;

KERNEL_START_VIRT = .;

.text : AT(KERNEL_PHYS)
{
stext = .;
build/boot.o (.text .text.*)    /* need this at the front */
*( EXCLUDE_FILE(build/init.o) .text .text.*)
}

. = ALIGN(4K);
etext = .;

.rodata :
{
srodata = .;
*(.rodata .rodata.*)
}
...
.text_ap : AT(AP_START) {
stext_ap = .;
*(.text_ap_entry)
etext_ap = .;
}
---------------------------------------------------------------------------
-- Symbol is a useless type, used to prevent us from forgetting to use
-- 'Address when referring to one.
---------------------------------------------------------------------------
type Symbol is (USELESS) with Size => System.Word_Size;
...

BITS 16

; we'll link this section down low, since it has to be in first
; 65535 bytes for real mode.
section .text_ap_entry
...
> readelf -a build/boot_ap.o
...
Size              EntSize          Flags  Link  Info  Align
[ 0]                   NULL             0000000000000000  00000000
0000000000000000  0000000000000000           0     0     0
[ 1] .shstrtab         STRTAB           0000000000000000  00000320
000000000000009e  0000000000000000           0     0     0
[ 2] .strtab           STRTAB           0000000000000000  000003c0
000000000000009d  0000000000000000           0     0     0
[ 3] .symtab           SYMTAB           0000000000000000  00000460
0000000000000198  0000000000000018           2    15     8
[ 4] .text             PROGBITS         0000000000000000  00000040
0000000000000000  0000000000000000  AX       0     0     16
[ 5] .text_ap_entry    PROGBITS         0000000000000000  00000040
000000000000008e  0000000000000000   A       0     0     16
...
...
subtype kernelTextPages is Virtmem.PFN range

subtype kernelROPages is Virtmem.PFN range

subtype kernelRWPages is Virtmem.PFN range
...
procedure determineFlagsAndMapFrame(frame : in Virtmem.PFN) is
...
begin

if frame in kernelTextPages then
mapPage(fromPhys,
toVirtLinear,
Virtmem.PG_KERNELDATARO,
kernelP4,
ok);

if not ok then raise RemapException; end if;

mapPage(fromPhys,
toVirtKernel,
Virtmem.PG_KERNELCODE,
kernelP4,
ok);
...
procedure unlink(ord : in Order; addr : in Virtmem.PhysAddress) with
SPARK_Mode => On,
Pre  => freeLists(ord).numFreeBlocks > 0,
Post => freeLists(ord).numFreeBlocks =
freeLists(ord).numFreeBlocks'Old - 1
is
block : aliased FreeBlock with

begin

declare
prevBlock : aliased FreeBlock with

nextBlock : aliased FreeBlock with
begin

-- decrement the free list count when we unlink somebody
freeLists(ord).numFreeBlocks :=
freeLists(ord).numFreeBlocks - 1;
end unlink;
type FreeNode is
record
end record with Size => 16 * 8;

record
next at 0 range 0..63;
prev at 8 range 0..63;
end record;

type Slab is limited record
freeList    : FreeNode;

numFree     : Integer := 0;
capacity    : Integer := 0;

blockOrder  : BuddyAllocator.Order;

mutex       : aliased Spinlock.Spinlock;

alignment   : System.Storage_Elements.Storage_Count;

initialized : Boolean := False;
end record;

-- GNAT-specific pragma
pragma Simple_Storage_Pool_Type(Slab);
...
objSlab : SlabAllocator.Slab;
type myObjPtr is access myObject;
for myObjPtr'Simple_Storage_Pool use objSlab;

procedure free is new Ada.Unchecked_Deallocation(myObject, myObjPtr);
obj : myObjPtr;

begin

SlabAllocator.setup(objSlab, myObject'Size);
obj := new Object;
...
---------------------------------------------------------------------------
-- FADT - Fixed ACPI Description Table.
---------------------------------------------------------------------------
record
firmwareControl     : Unsigned_32;  -- ignored if exFirmwareControl present
dsdt                : Unsigned_32;  -- ignored if exDsdt present
reserved1           : Unsigned_8;
powerMgmtProfile    : PowerManagementProfile;
sciInterrupt        : Unsigned_16;
smiCommand          : Unsigned_32;
acpiEnable          : Unsigned_8;
acpiDisable         : Unsigned_8;
S4BIOSReq           : Unsigned_8;
pStateControl       : Unsigned_8;
PM1AEventBlock      : Unsigned_32;
PM1BEventBlock      : Unsigned_32;
PM1AControlBlock    : Unsigned_32;
PM1BControlBlock    : Unsigned_32;
PM2ControlBlock     : Unsigned_32;
PMTimerBlock        : Unsigned_32;
GPE0Block           : Unsigned_32;
GPE1Block           : Unsigned_32;
PM1EventLength      : Unsigned_8;
PM1ControlLength    : Unsigned_8;
PM2ControlLength    : Unsigned_8;
PMTimerLength       : Unsigned_8;
GPE0BlockLength     : Unsigned_8;
GPE1BlockLength     : Unsigned_8;
GPE1Base            : Unsigned_8;
cStateControl       : Unsigned_8;
pLevel2Latency      : Unsigned_16;
pLevel3Latency      : Unsigned_16;
flushSize           : Unsigned_16;
flushStride         : Unsigned_16;
dutyOffset          : Unsigned_8;
dutyWidth           : Unsigned_8;
dayAlarm            : Unsigned_8;
monthAlarm          : Unsigned_8;
century             : Unsigned_8;   -- RTC index into RTC RAM if not 0
intelBootArch       : Unsigned_16;  -- IA-PC boot architecture flags
reserved2           : Unsigned_8;   -- always 0
flags               : Unsigned_32;  -- fixed feature flags
resetValue          : Unsigned_8;
armBootArch         : Unsigned_16;
exFirmwareControl   : Unsigned_64;
exDsdt              : Unsigned_64;

-- ACPI 6 fields (not supported yet)
--hypervisorVendor    : Unsigned_64;
end record with Size => 244*8;

record
firmwareControl     at 36  range 0..31;
dsdt                at 40  range 0..31;
reserved1           at 44  range 0..7;
powerMgmtProfile    at 45  range 0..7;
sciInterrupt        at 46  range 0..15;
smiCommand          at 48  range 0..31;
acpiEnable          at 52  range 0..7;
acpiDisable         at 53  range 0..7;
S4BIOSReq           at 54  range 0..7;
pStateControl       at 55  range 0..7;
PM1AEventBlock      at 56  range 0..31;
PM1BEventBlock      at 60  range 0..31;
PM1AControlBlock    at 64  range 0..31;
PM1BControlBlock    at 68  range 0..31;
PM2ControlBlock     at 72  range 0..31;
PMTimerBlock        at 76  range 0..31;
GPE0Block           at 80  range 0..31;
GPE1Block           at 84  range 0..31;
PM1EventLength      at 88  range 0..7;
PM1ControlLength    at 89  range 0..7;
PM2ControlLength    at 90  range 0..7;
PMTimerLength       at 91  range 0..7;
GPE0BlockLength     at 92  range 0..7;
GPE1BlockLength     at 93  range 0..7;
GPE1Base            at 94  range 0..7;
cStateControl       at 95  range 0..7;
pLevel2Latency      at 96  range 0..15;
pLevel3Latency      at 98  range 0..15;
flushSize           at 100 range 0..15;
flushStride         at 102 range 0..15;
dutyOffset          at 104 range 0..7;
dutyWidth           at 105 range 0..7;
dayAlarm            at 106 range 0..7;
monthAlarm          at 107 range 0..7;
century             at 108 range 0..7;
intelBootArch       at 109 range 0..15;
reserved2           at 111 range 0..7;
flags               at 112 range 0..31;
resetRegister       at 116 range 0..95;
resetValue          at 128 range 0..7;
armBootArch         at 129 range 0..15;
exFirmwareControl   at 132 range 0..63;
exDsdt              at 140 range 0..63;
exPM1AEventBlock    at 148 range 0..95;
exPM1BEventBlock    at 160 range 0..95;
exPM1AControlBlock  at 172 range 0..95;
exPM1BControlBlock  at 184 range 0..95;
exPM2ControlBlock   at 196 range 0..95;
exPMTimerBlock      at 208 range 0..95;
exGPE0Block         at 220 range 0..95;
exGPE1Block         at 232 range 0..95;

-- ACPI 6 fields
--sleepControlReg     at 244 range 0..95;
--sleepStatusReg      at 256 range 0..95;
--hypervisorVendor    at 268 range 0..63;
end record;
...
setup_bsp:
; Setup our kernel stack.
mov rsp, qword (STACK_TOP)

; Add a stack canary to bottom of primary stack for CPU #0
mov rax, qword (STACK_TOP - PER_CPU_STACK_SIZE + SEC_STACK_SIZE)
mov [rax], rbx

; Save rdi, rsi so adainit doesn't clobber them
push rdi
push rsi

call rax

; Restore arguments to kmain
pop rsi
pop rdi

mov rax, qword kmain
call rax
...
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GNAT Community 2020 is here! https://blog.adacore.com/gnat-community-2020-is-here Tue, 26 May 2020 09:59:00 -0400 Nicolas Setton https://blog.adacore.com/gnat-community-2020-is-here

We are happy to announce that the GNAT Community 2020 release is now available via https://www.adacore.com/download. Here are some release highlights:

GNAT compiler toolchain

The 2020 compiler includes tightening and enforcing of Ada rules, performance enhancements, and support for some Ada 202x features - watch this space for further news on this.

The compiler back-end has been upgraded to GCC 9 on all platforms except Mac OS - see below for further information about this exception.

ASIS is no longer supported, and we encourage you to switch to Libadalang for all your code intelligence needs. GNAT Community 2019 remains available for legacy support of ASIS.

RISC-V 64-bits

This year we have added a toolchain for RISC-V 64-bits hosted on Linux - you can try it out on boards like the HiFive Unleashed - and we include the emulator for this platform as well.

IDE

This release includes GNAT Studio, the evolution of GPS, our multi-language IDE for Ada, SPARK, C, C++ and Python. Notable features are:

• A Completely new engine for Ada/SPARK navigation, implemented via a language server based on Libadalang. This means, in particular, that navigation works without requiring you to compile the codebase first.
• Improved overall performance in the editors, the omnisearch, and the debugger.

• Several UI enhancements, especially the contextual menus which have been reorganized.

Please also note that we no longer support GNAT Studio on Mac OS.

Libadalang, a library for parsing and semantic analysis of Ada code, has made a lot of progress in the past year. In this GNAT Community release, you'll find:

• The Python-facing API is now compatible with Python 3..

• Support of Aggregate Projects has been added.

SPARK

For those looking to take their Ada programs to the next level, GNAT Community includes a complete SPARK toolchain now including the Lemma library (doc).

Toolchain and development environment enhancements are:

• New SPARK submenus and key shortcuts in GNAT Studio.

• Parallel analysis of subprograms.

• Automatic target configuration for GNAT runtimes.

Proving engine enhancements are:

• Support for infinite precision arithmetic in Ada.Numerics.Big_Numbers.Big_Integers/Big_Reals (doc).

• Support for partially initialized data in proof (doc).

• Detection of memory leaks by proof.

• Dead code detected by proof warnings.

• Improved floating-point support in Alt-Ergo prover.

SPARK language enhancements are:

• Support for local borrowers as part of pointer support through ownership.

• Many fixes in the new pointer support based on ownership.

• Detection of wrap-around on modular arithmetic with annotation No_Wrap_Around (doc).

• Support for forward goto.

• Support for raise expressions (doc).

• Detection of unsafe use of Unchecked_Conversion (doc).

• New annotation Might_Not_Return on procedures (doc).

• Volatility refinement aspects supported for types (doc).

• Allow SPARK_Mode Off inside subprograms.

• Support for volatile variables to prevent compiler optimizations (doc).

Support for Visual Studio Code

If you are using Visual Studio Code, we have written a prototype extension for Ada and SPARK as part of our work on the Ada Language Server: you can find it on the Visual Studio Marketplace.

Notes on Mac OS

Mac OS is becoming harder to maintain, especially the latest versions, which require code signing for binaries. For the GNAT Community 2020 release we decided not to codesign and notarize the binaries, so you'll have to circumvent the protections: see the README for the specific instructions. We have also removed support for the ARM cross compiler hosted on this platform, as well as GNAT Studio.

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1. Introduction

To learn a new programming language, an effective approach is to implement data structures common to computer programming. This is an effective strategy because the problem to be solved is well understood and several different forms of a given data structure are possible: bounded versus unbounded, sequential versus thread-safe, and so on. A clear understanding of the problem allows one to focus on the language details, and the multiple forms likely require a wide range of language features.

Fortunately, when learning SPARK, Ada programmers need not start from scratch. We can begin with an existing, production-ready Ada implementation for a common data structure and make the changes necessary to conform to SPARK. This approach is possible because the fundamental design, based on the principles of software engineering, is the same in both languages. We would have a package exporting a private type, with primitive operations manipulating that type; in other words, an abstract data type (ADT). The type might be limited, and might be tagged, using the same criteria in both languages to decide. Those primitive operations that change state would be procedures, with functions designed to be "pure" and side effects avoided. As a result, the changes need not be fundamental or extensive, although they are important and in some cases subtle.

The chosen Ada component is one that I have had for decades and have used in real-world applications. Specifically, this component defines a sequential, bounded stack ADT. The enclosing package is a generic so that the type of data contained in the stack objects need not be hard-coded. By "sequential" I mean that the code is not thread-safe. By "bounded" I mean that it is backed by an array, which as usual entails a discriminant on the private type to set the upper bound of the internal array component. Client misuse of the Push and Pop routines, e.g., pushing onto a full stack, raises exceptions. As Ada has evolved I have applied new features to make the code more robust, for example the Push and Pop routines use preconditions to prevent callers from misusing the abstraction, raising exceptions from within the preconditions instead of the procedure bodies.

This blog entry describes the transformation of that Ada stack ADT into a completely proven SPARK implementation that relies on static verification instead of run-time enforcement of the abstraction’s semantics. We will prove that there are no reads of unassigned variables, no array indexing errors, no range errors, no numeric overflow errors, no attempts to push onto a full stack, no attempts to pop from an empty stack, that subprogram bodies implement their functional requirements, and so on. As a result, we get a maximally robust implementation of a reusable stack abstraction providing all the facilities required for production use.

The transformation will occur in phases, following the adoption levels described in section 2. Each adoption level introduces more rigor and thus defines a simple, incremental transition approach.

Note that I assume familiarity with Ada, including preconditions and postconditions. Language details can be obtained from the online learning facilities available at https://learn.adacore.com/, an interactive site allowing one to enter, compile, and execute Ada programs in a web browser. We also assume a degree of familiarity with SPARK. That same web site provides a similar interactive environment and materials for learning SPARK, including formal proof.

In 2016, AdaCore collaborated with Thales in a series of experiments on the application of SPARK to existing software projects written in Ada. The resulting document presents a set of guidelines for adopting formal verification in existing projects. These guidelines are arranged in terms of five levels of software assurance, in increasing order of benefits and costs. The levels are named Stone, Bronze, Silver, Gold and Platinum. Successfully reaching a given level requires successfully achieving the goals of the previous levels as well.

The guidelines were developed jointly by AdaCore and Thales for the adoption of the SPARK language technology at Thales but are applicable across a wide range of application domains. The document is available online: http://www.adacore.com/knowled...

2.1 Stone Level

The goal at the Stone level is to identify as much code as possible that belongs to the SPARK subset. That subset provides a strong semantic coding standard that enforces safer use of Ada language features and forbids those features precluding analysis (e.g., exception handlers). The result is potentially more understandable, maintainable code.

2.2 Bronze Level

The goal at the Bronze level is to verify initialization and correct data flow, as indicated by the absence of GNATprove messages during SPARK flow analysis. Flow analysis detects programming errors such as reading uninitialized data, problematic aliasing between formal parameters, and data races between concurrent tasks. In addition, GNATprove checks unit specifications for the actual data read or written, and the flow of information from inputs to outputs. As one can see, this level provides significant benefits, and can be reached with comparatively low cost. There are no proofs attempted at this level, only data and flow analyses.

2.3 Silver Level

The goal at the Silver level is to statically prove absence of run-time errors (AoRTE), i.e., that there are no exceptions raised. Proof at this level detects programming errors such as divide by zero, array indexes that are out of bounds, and numeric overflow (integer, fixed-point and floating-point), among others. These errors are detected via the implicit language-defined checks that raise language-defined exceptions. The checks themselves preclude a number of significant situations, including, for example, buffer overflow, which is often exploited to inject malicious executable code.

Preconditions, among other additions, may be required to prove these checks. To illustrate the benefit and part of the cost of achieving the Silver level, consider the way the Ada version of the stack ADT uses preconditions for this purpose. (The complete Ada implementation is explored in section 4.1.) First, here is the full declaration for type Stack in the Ada package private part:

type Content is array (Positive range <>) of Element;

type Stack (Capacity : Positive) is record
Values : Content (1 .. Capacity);
Top    : Natural := 0;
end record;

The type Element represents the kind of individual values contained by stack objects. Top is used as the index into the array Values and can be zero. The Values array uses 1 for the lower index bound so when Top is zero the enclosing stack object is logically empty. The following function checks for that condition:

function Empty (This : Stack) return Boolean is
(This.Top = 0);

Consider, then, a function using Empty as a precondition. The function takes a stack parameter as input and returns the Element value at the logical top of the stack:

19    function Top_Element (This : Stack) return Element with
20      Pre => not Empty (This);

Given the precondition on line 20, within the function completion we know that Top has a value that is a potentially valid array index. (We'll also have to be more precise about Top's upper bound, as explained later in section 4.4.) There is no need for defensive code so the body is simply as follows:

57    function Top_Element (This : Stack) return Element is
58      (This.Values (This.Top));

If we did not have the precondition specified, GNATprove would issue a message:

58:24: medium: array index check might fail, (e.g. when This = (…, Top => 0) and …)


The message shows an example situation in which the check could fail: Top is zero, i.e., the stack is empty. (We have elided some of the message content to highlight the part mentioning Top.)

GNATprove will attempt to prove, statically, that the preconditions hold at every call site, flagging those calls, if any, in which the preconditions might not hold. Those failures must be addressed at the Silver level because the preconditions are necessary to the proof of absence of run-time errors.

As you can see, the Silver level provides highly significant benefits, but does require more contracts and potentially complex changes to the code. The effort required to achieve this level can be high. Arguably, however, this level should be the minimum target level, especially if the application executable is to be deployed with run-time checks disabled.

2.4 Gold Level

The goal at the Gold level is proof of key integrity properties. These properties are typically derived from software requirements but also include maintaining critical data invariants throughout execution.

Working at this level assumes prior completion at the Silver level to ensure program integrity, such that control flow cannot be circumvented through run-time errors and data cannot be corrupted. Verification at this level is also expected to pass without any violations.

Key integrity properties are expressed as additional preconditions and postconditions beyond those used for defensive purposes.  In addition, the application may explicitly raise application-defined exceptions to signal violations of integrity properties. GNATprove will attempt to prove that the code raising an exception is never reached, and thus, that the property violation never occurs. This approach may also require further proof-oriented code.

The Gold level provides extremely significant benefits. In particular, it can be less expensive to prove at this level than to test to the same degree of confidence. However, the analysis may take a long time, may require adding more precise types (ranges), and may require adding more preconditions and postconditions. Even if a property is provable, automatic provers may fail to prove it due to limitations of the provers, requiring either manual proof or, alternatively, testing.

2.5 Platinum Level

The goal at the Platinum level is nothing less than full functional proof of the requirements, including the functional unit level requirements, but also any abstract requirements such as, for example, safety and security.

As with the Gold level, the application code must pass SPARK analysis without any violations. Furthermore, at the Platinum level GNATprove must verify complete user specifications for type invariants, preconditions, postconditions, type predicates, loop variants, and loop termination.

The effort to achieve Platinum level is high, so high that this level is not recommended during initial adoption of SPARK.

3. Development Environment and Configuration

When we say we use SPARK, we mean that we develop the sources in the SPARK language, but also that we use the SPARK analysis tool to examine and verify those sources. We developed our sources in GNAT Studio (formerly GPS), a multi-lingual IDE supporting both Ada and SPARK, among others. The SPARK analysis tool is named GNATprove, a command-line tool integrated with GNAT Studio. GNAT Studio facilitates invocation of GNATprove with control over switches and source files, providing traversable results and even, if need be, interactive proof.

3.1 The Provers

A critical concept for using GNATprove is that it transparently invokes third-party “provers” to analyze the given source files. These provers are somewhat specialized in their ability to analyze specific semantics expressed by the source code. As a result, invocation of a series of provers may be required before some source code is successfully proven. In addition, we may need to ask the provers to “try harder” when attempting to analyze difficult situations. GNATprove can do both for us via the “level=n” switch, where “n” is a number from 0 to 4 indicating increasing strength of analysis and additional provers invoked. In proving our stack implementation we use level 4.

3.2 Language-Defined Run-time Checks

GNATprove is also integrated with the GNAT Ada compiler, including the analysis of language-defined run-time checks produced by the compiler. GNATprove attempts to verify that no exceptions are raised due to these checks. It will do so even if we suppress the checks with compiler switches or pragma Suppress, so we can interpret lack of corresponding messages as successful verification of those checks.

Integer overflow checks are a special case, and as a result have a dedicated GNAT switch that affects whether that specific check is generated by the compiler. They are a special case because, in addition to the functional code, they may appear in the logical assertions about the functional code, including subprogram preconditions and postconditions. In these contexts, we might expect them to behave mathematically, without implementation bounds. For example, consider the following declaration for a procedure that enters a log entry into a file:

5    Entry_Num : Natural := 0;
6
7    procedure Log (This : String) with
8      Pre    => Entry_Num + 1 <= Integer'Last,
9      Global => (In_Out => Entry_Num);

The procedure body increments Entry_Num by one and then prepends the result to the string passed as the log entry. This addition in the body might overflow, but the issue under consideration is the addition in the precondition on line 8. If Entry_Num is Integer’Last at the point of the call, the addition on line 8 will overflow, as GNATprove indicates:

8:26: medium: overflow check might fail (e.g. when Entry_Num = Natural'Last)


We could revise the code so that the expression cannot overflow:

Pre => Entry_Num <= Integer'Last - 1,

although that is slightly less readable. Other alternatives within the code are possible as well. However, with regard to switches pertinent for check generation, GNAT provides the “-gnato” switch that allows us to control how integer overflow is treated. (There is a pragma as well, with the same effects.) We can use that switch to have the compiler implement integer arithmetic mathematically, without bounds, the way we might conceptually expect it to work within logical, non-functional assertions. As a result, there will be no integer overflow checks generated. The default effect for the switch, and the default if the switch is not present, is to enable overflow checks in both functional and assertion code so we just need to be aware of non-default usage when we want to determine whether integer overflow checks have been verified. (See the SPARK User Guide, section 5.7 “Overflow Modes” for the switch parameters.) In our GNAT project file, the switch is explicitly set to enable overflow checks in both the functional code and the assertion code.

3.3 Source Code File Organization

The main program declares objects of a type Stack able to contain character values. That Stack type is provided by the package Character_Stacks, which is an instantiation of a generic package defining a stack abstract data type. The instantiation is specified such that objects of the resulting Stack type can contain character values.

Logically, there are four source files in the application: two (declaration and body) for the generic package, one for the instantiation of that generic package, and one containing the demonstration main subprogram.  Operationally, however, there are multiple source files for the generic package. Rather than have one implementation that we alter as we progress through the SPARK adoption levels, we have chosen to have a distinct generic package for each level. Each generic package implements a common stack ADT in a manner consistent with an adoption level. The differences among them reflect the changes required for the different levels. This approach makes it easier to keep the differences straight when examining the code. Furthermore, we can apply the proof analyses to a conceptually common abstraction at arbitrary adoption levels without having to alter the code.

In addition to the content differences required by the adoption levels, each generic package name reflects the corresponding level. We have generic package Bounded_Stacks_Stone for the Stone level, Bounded_Stacks_Gold for the Gold level, and so on. Therefore, although the instantiation is always named Character_Stacks, we have multiple generic packages available to declare the one instantiation used.

There are also multiple files for the instantiations. Each instantiation is located within a dedicated source file corresponding to a given adoption level (lines 2 and 3 below). For example, here is the content of the file providing the instance for the Stone level:

1 pragma Spark_Mode (On);
2 with Bounded_Stacks_Stone;
3 package Character_Stacks is new Bounded_Stacks_Stone
4   (…);

The file names for these instances must be unique but are otherwise arbitrary. For the above, the file name is “character_stacks-stone.ads” because it is the instance of the Stone level generic.

Only one of these instances can be used when GNATprove analyzes the code (or when building the executable). To select among them we use a “scenario variable” defined in the GNAT project file that has scenario values matching the adoption level names. In the IDE this scenario variable is presented with a pull-down menu so all we must do to work at a given level is select the adoption level name in the pull-down list. The project file then selects the instantiation file corresponding to the level, e.g., “character_stacks-silver.ads” when the Silver level is selected.

There are also multiple source files for the main program. Rather than have one file that must be edited as we prove the higher levels, we have two: one for all levels up to and including the Silver level, and one for all levels above that. The scenario variable also determines which of these two source files is active.

3.4 Verifying Generic Units

One of the current limitations of GNATprove is that it cannot verify generic units on their own. GNATprove must instead be provided an instantiation to verify. Therefore, whenever we say that we are verifying the generic package defining the stack ADT, we mean we are invoking GNATprove on an instantiation of that generic. As noted earlier in section 3.3, there are multiple source files containing these instantiations so we must select the file corresponding to the desired level when we want to verify the generic package alone.

However, because there are only four total files required at any one time, we usually invoke the IDE action that has GNATprove analyze all the files in the closure of the application. The instantiation file corresponding to the scenario variable’s current selection will be analyzed; other instantiation files are ignored. This approach also verifies the main program’s calls to the stack routines, which is vital to the higher adoption levels.

Our first main procedure, used for all adoption levels up through Silver, declares two stack objects (line 6 below) and manipulates them via the abstraction’s interface:

1 with Ada.Text_IO;       use Ada.Text_IO;
2 with Character_Stacks;  use Character_Stacks;
3
4 procedure Demo_AoRTE with SPARK_Mode is
5
6    S1, S2 : Stack (Capacity => 10);  -- arbitrary
7
8    X, Y : Character;
9
10 begin
11    pragma Assert (Empty (S1) and Empty (S2));
12    pragma Assert (S1 = S2);
13    Push (S1, 'a');
14    Push (S1, 'b');
15    Put_Line ("Top of S1 is '" & Top_Element (S1) & "'");
16
17    Pop (S1, X);
18    Put_Line ("Top of S1 is '" & Top_Element (S1) & "'");
19    Pop (S1, Y);
20    pragma Assert (Empty (S1) and Empty (S2));
21    Put_Line (X & Y);
22
23    Reset (S1);
24    Put_Line ("Extent of S1 is" & Extent (S1)'Image);
25
26    Put_Line ("Done");
27 end Demo_AoRTE;

This is the “demo_aorte.adb” file. The purpose of the code is to illustrate issues found at the initial levels, including proof in a caller context. It has no other functional purpose whatsoever. As we progress through the levels, we will add more assertions to highlight more issues, as will be seen in the other main procedure in the  “demo_gold.adb” file.

The initial version defines a canonical representation of a sequential, bounded stack. As an abstract data type, the Stack type is declared as a private type with routines manipulating objects of the type. The type is declared within a generic package that has one generic formal parameter, a type representing the kind of elements contained by Stack objects. This approach is used in all the implementations.

Some routines have “defensive” preconditions to ensure correct functionality. They raise exceptions, declared within the package, when the preconditions do not hold.

The generic package in Ada is declared as follows:

1 generic
2    type Element is private;
3 package Bounded_Stacks_Magma is
4
5    type Stack (Capacity : Positive) is private;
6
7    procedure Push (This : in out Stack; Item : in Element) with
8      Pre => not Full (This) or else raise Overflow;
9
10    procedure Pop (This : in out Stack; Item : out Element) with
11      Pre => not Empty (This) or else raise Underflow;
12
13    function Top_Element (This : Stack) return Element with
14      Pre => not Empty (This) or else raise Underflow;
15    --  Returns the value of the Element at the "top" of This
16    --  stack, i.e., the most recent Element pushed. Does not
17    --  remove that Element or alter the state of This stack
18    --  in any way.
19
20    overriding function "=" (Left, Right : Stack) return Boolean;
21
22    procedure Copy (Destination : out Stack; Source : Stack) with
23      Pre => Destination.Capacity >= Extent (Source)
24               or else raise Overflow;
25    --  An alternative to predefined assignment that does not
26    --  copy all the values unless necessary. It only copies
27    --  the part "logically" contained, so is more efficient
28    --  when Source is not full.
29
30    function Extent (This : Stack) return Natural;
31    --  Returns the number of Element values currently
32    --  contained within This stack.
33
34    function Empty (This : Stack) return Boolean;
35
36    function Full (This : Stack) return Boolean;
37
38    procedure Reset (This : out Stack);
39
40    Overflow  : exception;
41    Underflow : exception;
42
43 private
44
45    type Content is array (Positive range <>) of Element;
46
47    type Stack (Capacity : Positive) is record
48       Values : Content (1 .. Capacity);
49       Top    : Natural := 0;
50    end record;
51
52 end Bounded_Stacks_Magma;

This version is below the Stone level because it is not within the SPARK subset, due to the raise expressions on lines 8, 11, 14, and 24. We will address those constructs in the Stone version.

The generic package body is shown below.

1 package body Bounded_Stacks_Magma is
2
3    procedure Reset (This : out Stack) is
4    begin
5       This.Top := 0;
6    end Reset;
7
8    function Extent (This : Stack) return Natural is
9       (This.Top);
10
11    function Empty (This : Stack) return Boolean is
12      (This.Top = 0);
13
14    function Full (This : Stack) return Boolean is
15      (This.Top = This.Capacity);
16
17    procedure Push (This : in out Stack; Item : in Element) is
18    begin
19       This.Top := This.Top + 1;
20       This.Values (This.Top) := Item;
21    end Push;
22
23    procedure Pop (This : in out Stack; Item : out Element) is
24    begin
25       Item := This.Values (This.Top);
26       This.Top := This.Top - 1;
27    end Pop;
28
29    function Top_Element (This : Stack) return Element is
30      (This.Values (This.Top));
31
32    function "=" (Left, Right : Stack) return Boolean is
33      (Left.Top = Right.Top and then
34       Left.Values (1 .. Left.Top) = Right.Values (1 .. Right.Top));
35
36    procedure Copy (Destination : out Stack; Source : Stack) is
37       subtype Contained is Integer range 1 .. Source.Top;
38    begin
39       Destination.Top := Source.Top;
40       Destination.Values (Contained) := Source.Values (Contained);
41    end Copy;
42
43 end Bounded_Stacks_Magma;

Note that both procedure Copy and function “=” are defined for the sake of increased efficiency when the objects in question are not full. The procedure only copies the slice of Source.Values that represents the Element values logically contained at the time of the call. The language-defined assignment operation, in contrast, would copy the entire contents. Similarly, the overridden equality operator only compares the array slices, rather than the entire arrays, after first ensuring the stacks are the same logical size.

However, in addition to efficiency, the "=" function is also required for proper semantics. The comparison should not compare array elements that are not, and perhaps never have been, currently contained in the stack objects. The predefined equality would do so and must, therefore, be replaced.

The changes to the body made for the sake of SPARK will amount to moving certain bodies to the package declaration so we will not show the package body again. The full Platinum implementation, both declaration and body, is provided in section 6.

4.2 Stone Implementation

The Stone level version of the package cannot have the "raise expressions" in the preconditions because they are not in the SPARK subset. The rest of the preconditions are unchanged. Here are the updated declarations for Push and Pop, for example:

procedure Push (This : in out Stack; Item : in Element) with
Pre => not Full (This);

procedure Pop (This : in out Stack; Item : out Element) with
Pre => not Empty (This);

When we get to the adoption levels involving proof, GNATprove will attempt to verify statically that the preconditions will hold at each call site. Either that verification will succeed, or we will know that we must change the calling code accordingly. Therefore, the prohibited “raise expressions” are not needed.

The exception declarations, although within the subset, are also removed because they are no longer needed.

The remaining code is wholly within the SPARK subset so we have reached the Stone level.

4.3 Bronze Implementation

The Bronze level is about initialization and data flow. When we apply GNATprove to the Stone version in flow analysis mode, GNATprove issues messages on the declarations of procedures Copy and Reset in the generic package declaration:

medium: "Destination.Values" might not be initialized in "Copy"
high: "This.Values" is not initialized in "Reset"


The procedure declarations are repeated below for reference:

procedure Copy (Destination : out Stack; Source : Stack) with
Pre => Destination.Capacity >= Extent (Source);

procedure Reset (This : out Stack);

Both messages result from the fact that the updated formal stack parameters have mode “out” specified. That mode, in SPARK, means more than it does in Ada. It indicates that the actual parameters are fully assigned by the procedures, but these two procedure bodies do not do so. Procedure Reset simply sets the Top to zero because that is all that a stack requires, at run-time, to be fully reset. It does nothing at all to the Values array component. Likewise, procedure Copy may only assign part of the array, i.e., just those array components that are logically part of the Source object. (Of course, if Source is full, the entire array is copied.) In both subprograms our notion of being fully assigned is less than SPARK requires. Therefore, we have two choices. Either we assign values to all components of the record, or we change the modes to “in out.” These two procedures exist for the sake of efficiency, i.e., not writing any more data than logically necessary. Having Reset assign anything to the array component would defeat the purpose. For the same reason, having Copy assign more than the partial slice (when the stack is not full) is clearly inappropriate. Therefore, we change the mode to “in out” for these two subprograms. In other cases we might change the implementations to fully assign the objects.

The other change required for initialization concerns the type Stack itself. In the main subprogram, GNATprove complains that the two objects of type Stack have not been initialized:

warning: "S1" may be referenced before it has a value
high: private part of "S1" is not initialized
warning: "S2" may be referenced before it has a value
high: private part of "S2" is not initialized
high: private part of "S1" is not initialized


Our full definition of the Stack type in the private part is such that default initialization (i.e., elaboration of object declarations without an explicit initial value) will assign the record components so that a stack will behave as if initially empty. Specifically, default initialization assigns zero to Top (line 5 below), and since function Empty examines only the Top component, such objects are empty.

1 type Content is array (Positive range <>) of Element;
2
3 type Stack (Capacity : Positive) is record
4    Values : Content (1 .. Capacity);
5    Top    : Natural := 0;
6 end record;

Proper run-time functionality of the Stack ADT does not require the Values array component to be assigned by default initialization. But just as with Reset and Copy, although this approach is sufficient at run-time, the resulting objects will not be fully initialized in SPARK, which analyzes the code prior to run-time. As a result, we need to assign an array aggregate to the Values component as well. Expressing the array aggregate is problematic because the array component type is the generic formal private type Element, with a private view within the package. Inside the generic package we don’t know how to construct a value of type Element so we cannot construct an aggregate containing such values. Therefore, we add the Default_Value generic formal object parameter and use it to initialize the array components.

This new generic formal parameter, shown below on line 5, is added from the Bronze version onward:

1 generic
2    type Element is private;
3    --  The type of values contained by objects of type Stack
4
5    Default_Value : Element;
6    --  The default value used for stack contents. Never
7    --  acquired as a value from the API, but required for
8    --  initialization in SPARK.
9 package Bounded_Stacks_Bronze is

The full definition for type Stack then uses that parameter to initialize Values (line 2):

1 type Stack (Capacity : Positive) is record
2    Values : Content (1 .. Capacity) := (others => Default_Value);
3    Top    : Natural := 0;
4 end record;

With those changes in place flow analysis completes without further complaint. The implementation has reached the Bronze level.

The need for that additional generic formal parameter is unfortunate because it becomes part of the user’s interface without any functional use. None of the API routines ever return it as such, and the actual value chosen is immaterial.

Note that SPARK will not allow the aggregate to contain default components (line 2):

1 type Stack (Capacity : Positive) is record
2    Values : Content (1 .. Capacity) := (others => <>);
3    Top    : Natural := 0;
4 end record;

as per SPARK RM 4.3(1).

Alternatively, we could omit this generic formal object parameter if we use an aspect to promise that the objects are initially empty, and then manually justify any resulting messages. We will in fact add that aspect for other reasons, but we prefer to have proof as automated as possible, for convenience and to avoid human error.

Finally, although the data dependency contracts, i.e., the “Global” aspects, would be generated automatically, we add them explicitly, indicating that there are no intended accesses to any global objects. For example, on line 3 in the following:

1 procedure Push (This : in out Stack;  Item : Element) with
2   Pre    => not Full (This),
3   Global => null;

We do so because mismatches between reality and the generated contracts are not reported by GNATprove, but we prefer positive confirmation for our understanding of the dependencies.

The flow dependency contracts (the “Depends” aspects) also can be generated automatically. Unlike the data dependency contracts, however, usually these can be omitted from the code even though mismatches with the corresponding bodies are not reported. That lack of notification is not a problem because the generated contracts are safe: they express at least the dependencies that the code actually exhibits. Therefore, all actual dependencies are covered. For example, a generated flow dependency will state that all outputs depend on all inputs, which is possible but not necessarily the case.

However, overly conservative contracts can lead to otherwise-avoidable issues with proof, leading the developer to add precise contracts explicitly when necessary. The other reason to express them explicitly is when we want to prove data flow dependencies as part of the abstract properties, for example data flowing only between units at appropriate security levels. We are not doing so in this case.

4.4 Silver Implementation

If we try to prove the Bronze level version of the generic package, GNATprove will complain about various run-time checks that cannot be proved in the generic package body. The Silver level requires these checks to be proven not to fail, i.e., not to raise exceptions.

The check messages are as follows, preceded by the code fragments they reference, with some message content elided in order to emphasize parts that lead us to the solution:

37    procedure Push (This : in out Stack; Item : in Element) is
38    begin
39       This.Top := This.Top + 1;
40       This.Values (This.Top) := Item;
41    end Push;
bounded_stacks_silver.adb:39:28: medium: overflow check might fail, … (e.g. when This = (…, Top => Natural'Last) …

bounded_stacks_silver.adb:40:24: medium: array index check might fail, … (e.g. when This = (…, Top => 2) and This.Values'First = 1 and This.Values'Last = 1)
47    procedure Pop (This : in out Stack; Item : out Element) is
48    begin
49       Item := This.Values (This.Top);
50       This.Top := This.Top - 1;
51    end Pop;
bounded_stacks_silver.adb:49:32: medium: array index check might fail, … (e.g. when This = (…, Top => 2) and This.Values'First = 1 and This.Values'Last = 1)
57    function Top_Element (This : Stack) return Element is
58      (This.Values (This.Top));
bounded_stacks_silver.adb:58:24: medium: array index check might fail, … (e.g. when This = (…, Top => 2) and This.Values'First = 1 and This.Values'Last = 1)
64    function "=" (Left, Right : Stack) return Boolean is
65       (Left.Top = Right.Top and then
66        Left.Values (1 .. Left.Top) = Right.Values (1 .. Right.Top));
bounded_stacks_silver.adb:66:12: medium: range check might fail, … (e.g. when Left = (Capacity => 1, …, Top => 2) …

bounded_stacks_silver.adb:66:43: medium: range check might fail, … (e.g. when Right = (Capacity => 1, …, Top => 2) …
72    procedure Copy (Destination : in out Stack; Source : Stack) is
73       subtype Contained is Integer range 1 .. Source.Top;
74    begin
75       Destination.Top := Source.Top;
76       Destination.Values (Contained) := Source.Values (Contained);
77    end Copy;
bounded_stacks_silver.adb:76:47: medium: range check might fail, … (e.g. when Destination = (Capacity => 1, …) and Source = (Capacity => 1, …), Top => 2)


All of these messages indicate that the provers do not know that the Top component is always in the range 0 .. Capacity. The code has not said so, and indeed, there is no way to use a discriminant in a scalar record component declaration to constrain the component’s range.  This is what we would write for the record type implementing type Stack in the full view, if we could (line 3):

1 type Stack (Capacity : Positive) is record
2    Values : Content (1 .. Capacity) := (others => Default_Value);
3    Top    : Natural range 0 .. Capacity := 0;
4 end record;

but that range constraint on Top is not legal. The reason it is illegal is that the application can change the value of a discriminant at run-time, under controlled circumstances, but there is no way at run-time to change the range checks in the object code generated by the compiler. However, with Ada and SPARK there is now a way to express the constraint on Top, and the provers will recognize the meaning during analysis. Specifically, we apply a “subtype predicate” to the record type declaration (line 5):

1 type Stack (Capacity : Positive) is record
2    Values : Content (1 .. Capacity) := (others => Default_Value);
3    Top    : Natural := 0;
4 end record with
5   Predicate => Top in 0 .. Capacity;

This aspect informs the provers that the Top component for any object of type Stack is always in the range 0 .. Capacity. That addition successfully addresses all the messages about the generic package body. Note that the provers will verify the predicate too.

However, GNATprove also complains about the main program. Consider that the first two assertions in the main procedure are not verified:

10   begin
11      pragma Assert (Empty (S1) and Empty (S2));
12      pragma Assert (S1 = S2);

GNATprove emits:

11:19: medium: assertion might fail, cannot prove Empty (S1)
12:19: medium: assertion might fail, cannot prove S1 = S2

We can address the issue for function Empty, partly, by adding another aspect to the declaration of type Stack, this time to the visible declaration:

type Stack (Capacity : Positive) is private
with Default_Initial_Condition => Empty (Stack);

The new aspect indicates that default initialization results in stack objects that are empty, making explicit, and especially, verifiable, the intended initial object state. We will be notified if GNATprove determines that the aspect does not hold.

That new aspect will handle the first assertion in the main program on line 11 but GNATprove complains throughout the main procedure that the preconditions involving Empty and Full cannot be proven. For example:

13    Push (S1, 'a');
14    Push (S1, 'b');
15    Put_Line ("Top of S1 is '" & Top_Element (S1) & "'");

GNATprove emits:

13:06: medium: precondition might fail, cannot prove not Full (This)

14:06: medium: precondition might fail, cannot prove not Full (This) [possible explanation: call at line 13 should mention This (for argument S1) in a postcondition]

15:35: medium: precondition might fail, cannot prove not Empty (This) [possible explanation: call at line 14 should mention This (for argument S1) in a postcondition]

Note the “possible explanations” that GNATprove gives us. These are clear indications that we are not specifying sufficient postconditions. Remember that when analyzing code that includes a call to some procedure, the provers’ knowledge of the call’s effect is provided entirely by the procedure’s postcondition. That postcondition might be insufficient, especially if it is absent!

Therefore, we must tell the provers about the effects of calling Push and Pop, as well as the other routines that change state. We add a new postcondition on Push (line 3):

1 procedure Push (This : in out Stack;  Item : Element) with
2   Pre    => not Full (This),
3   Post   => Extent (This) = Extent (This)'Old + 1,
4   Global => null;

The new postcondition expresses the fact that the Stack contains one more Element value after the call. This is sufficient because the provers know that function Extent is simply the value of Top:

function Extent (This : Stack) return Natural is
(This.Top);

Hence the provers know that Top is incremented by Push.

The same approach addresses the messages for Pop (line 3):

1 procedure Pop (This : in out Stack; Item : out Element) with
2   Pre    => not Empty (This),
3   Post   => Extent (This) = Extent (This)'Old - 1,
4   Global => null;

In the above we say that the provers know what the function Extent means. For that to be the case when verifying client calls, we must move the function completion from the generic package body to the generic package declaration. In addition, the function must be implemented as an “expression function,” which Extent already is (see above). As expression functions in the package spec, the provers will know the semantics of those functions automatically, as if each is given a postcondition restating the corresponding expression explicitly. We also need functions Full and Empty to be known in this manner. Therefore, we move the Extent, Empty, and Full function completions, already expression functions, from the generic package body to the package declaration. We put them in the private part because these implementation details should not be exported to clients.

However, we have a potential overflow in the postcondition for Push, i.e., the increment of the number of elements contained after Push returns (line 3 below). The postcondition for procedure Pop, of course, does not have that problem.

1 procedure Push (This : in out Stack;  Item : Element) with
2   Pre    => not Full (This),
3   Post   => Extent (This) = Extent (This)'Old + 1,
4   Global => null;

The increment might overflow because Extent returns a value of subtype Natural, which could be the value Integer'Last. Hence the increment could raise Constraint_Error and the check cannot be verified. We must either apply the “-gnato” switch so that assertions can never overflow, or alternatively, declare a safe subrange so that the result of the addition cannot be greater than Integer'Last.

Our choice is to declare a safe subrange because the effects are explicit in the code, as opposed to an external switch. Here are the added subtype declarations:

subtype Element_Count is
Integer range 0 .. Integer'Last - 1;
--  The number of Element values currently contained
--  within any given stack. The lower bound is zero
--  because a stack can be empty. We limit the upper
--  bound (minimally) to preclude overflow issues.

subtype Physical_Capacity is
Element_Count range 1 .. Element_Count'Last;
--  The range of values that any given stack object can
--  specify (via the discriminant) for the number of
--  Element values the object can physically contain.
--  Must be at least one.

We use the second subtype for the discriminant in the partial view for Stack (line 1):

1 type Stack (Capacity : Physical_Capacity) is private
2    with Default_Initial_Condition => Empty (Stack);

and both subtypes in the full declaration in the private part (lines 1, 3, and 5):

1 type Content is array (Physical_Capacity range <>) of Element;
2
3 type Stack (Capacity : Physical_Capacity) is record
4    Values : Content (1 .. Capacity) := (others => Default_Value);
5    Top    : Element_Count := 0;
6 end record with
7   Predicate => Top in 0 .. Capacity;

The function Extent is changed to return a value of the subtype Element_Count so adding one in the postcondition cannot go past Integer’Last. Overflow is precluded but note that there will now be range checks for GNATprove to verify.

With these changes in place we have achieved the Silver level. There are no run-time check verification failures and the defensive preconditions are proven at their call sites.

4.5 Gold Implementation

We will now address the remaining changes needed to reach the Gold level. The process involves iteratively attempting to prove the main program that calls the stack routines and makes assertions about the conditions that follow. This process will result in changes to the generic package, especially postconditions, so it will require verification along with the main procedure. Those additional postconditions may require additional preconditions as well.

In general, a good way to identify postcondition candidates is to ask ourselves what conditions we, as the developers, know to be true after a call to the routine in question. Then we can add assertions after the calls to see if the provers can verify those conditions. If not, we extend the postcondition on the routine.

For example, we can say that after a call to Push, the corresponding stack cannot be empty. Likewise, after a call to Pop, the stack cannot be full. These additions are not required for the sake of assertions or other preconditions because the Extent function already tells the provers what they need to know in this regard. However, they are good documentation and may be required to prove additional conditions added later. (That is the case, in fact, as will be shown.)

To see what other postconditions are required, we now switch to the other main procedure, in the “demo_gold.adb” file. This version of the demo program includes a number of additional assertions:

1 with Ada.Text_IO;       use Ada.Text_IO;
2 with Character_Stacks;  use Character_Stacks;
3
4 procedure Demo_Gold with SPARK_Mode is
5
6    S1, S2 : Stack (Capacity => 10);  -- arbitrary
7
8    X, Y : Character;
9
10 begin
11    pragma Assert (Empty (S1) and Empty (S2));
12    pragma Assert (S1 = S2);
13    Push (S1, 'a');
14    pragma Assert (not Empty (S1));
15    pragma Assert (Top_Element (S1) = 'a');
16    Push (S1, 'b');
17    pragma Assert (S1 /= S2);
18
19    Put_Line ("Top of S1 is '" & Top_Element (S1) & "'");
20
21    Pop (S1, X);
22    Put_Line ("Top of S1 is '" & Top_Element (S1) & "'");
23    Pop (S1, Y);
24    pragma Assert (X = 'b');
25    pragma Assert (Y = 'a');
26    pragma Assert (S1 = S2);
27    Put_Line (X & Y);
28
29    Push (S1, 'a');
30    Copy (Source => S1, Destination => S2);
31    pragma Assert (S1 = S2);
32    pragma Assert (Top_Element (S1) = Top_Element (S2));
33    pragma Assert (Extent (S1) = Extent (S2));
34
35    Reset (S1);
36    pragma Assert (Empty (S1));
37    pragma Assert (S1 /= S2);
38
39    Put_Line ("Done");
40 end Demo_Gold;

For example, we have added assertions after the calls to Reset and Copy, on lines 31 through 33 and 36 through 37, respectively. GNATprove now emits the following (elided) messages for those assertions:

demo_gold.adb:31:19: medium: assertion might fail, cannot prove S1 = S2 (e.g. when S1 = (…, Top => 0) and S2 = (…, Top => 0)) [possible explanation: call at line 30 should mention Destination (for argument S2) in a postcondition]
demo_gold.adb:36:19: medium: assertion might fail, cannot prove Empty (S1) … [possible explanation: call at line 35 should mention This (for argument S1) in a postcondition]

Note again the “possible explanation” hints. For the first message we need to add a postcondition on Copy specifying that the value of the argument passed to Destination will be equal to that of the Source argument (line 3):

1 procedure Copy (Destination : in out Stack; Source : Stack) with
2   Pre    => Destination.Capacity >= Extent (Source),
3   Post   => Destination = Source,
4   Global => null;

We must move the “=” function implementation to the package spec so that the provers will know the meaning. The function was already completed as an expression function so moving it to the spec is all that is required.

For the second message, regarding the failure to prove that a stack is Empty after Reset, we add a postcondition to that effect (line 2):

1 procedure Reset (This : in out Stack) with
2   Post   => Empty (This),
3   Global => null;

The completion for function Empty was already moved to the package spec, earlier.

The implementations of procedure Copy and function “=” might have required explicit loops, likely requiring loop invariants, but using array slicing we can express the loop implicitly. Here is function “=” again, for example:

1 function "=" (Left, Right : Stack) return Boolean is
2   (Left.Top = Right.Top and then
3    Left.Values (1 .. Left.Top) = Right.Values (1 .. Right.Top));

The slice comparison on line 3 expresses an implicit loop for us, as does the slice assignment in procedure Copy.

The function could have been implemented as follows, with an explicit loop:

1 function "=" (Left, Right : Stack) return Boolean is
2 begin
3    if Left.Top /= Right.Top then
4       --  They hold a different number of element values so
5       --  cannot be equal.
6       return False;
7    end if;
8    --  The two Top values are the same, and the arrays
9    --  are 1-based, so the bounds are the same. Hence the
10    --  choice of Left.Top or Right.Top is arbitrary and
11    --  there is no need for index offsets.
12    for K in 1 .. Left.Top loop
13       if Left.Values (K) /= Right.Values (K) then
14          return False;
15       end if;
16       pragma Loop_Invariant
17                (Left.Values (1 .. K) = Right.Values (1 .. K));
18    end loop;
19    --  We didn't find a difference
20    return True;
21 end "=";

Note the loop invariant on lines 16 and 17. In some circumstances GNATprove will handle the invariants for us but often it cannot. In practice, writing sufficient loop invariants is one of the more difficult facets of SPARK development so the chance to avoid them is welcome.

Continuing, we know that after the body of Push executes, the top element contained in the stack will be the value passed to Push as an argument. But the provers cannot verify an assertion to that effect (line 15 below):

13      Push (S1, 'a');
14      pragma Assert (not Empty (S1));
15      pragma Assert (Top_Element (S1) = 'a');

GNATprove emits this message:

demo_gold.adb:15:19: medium: assertion might fail, cannot prove Top_Element (S1) = 'a'

We must extend the postcondition for Push to state that Top_Element would return the value just pushed, as shown on line 4 below:

1 procedure Push (This : in out Stack;  Item : Element) with
2   Pre    => not Full (This),
3   Post   => not Empty (This)
4             and then Top_Element (This) = Item
5             and then Extent (This) = Extent (This)'Old + 1,
6   Global => null;

Now the assertion on line 15 is verified successfully.

Recall that the precondition for function Top_Element is that the stack is not empty. We already have that assertion in the postcondition (line 3) so the precondition for Top_Element is satisfied. We must use the short circuit form for the conjunction, though, to control the order of evaluation so that “not Empty” is verified before Top_Element.

The short-circuit form on line 4 necessitates the same form on line 5, per Ada rules. That triggers a subtle issue flagged by GNATprove. The short-circuit form, by definition, means that the evaluation of line 5 might not occur. If it is not evaluated, we’ve told the compiler to call Extent and make a copy of the result (via ‘Old, on the right-hand side of “=”) that will not be needed. Moreover, the execution of Extent might raise an exception. Therefore, the language disallows applying ‘Old in any potentially unevaluated expression that might raise exceptions. As a consequence, in line 5 we cannot apply ‘Old to the result of calling Extent. GNATprove issues this error message:

prefix of attribute "Old" that is potentially unevaluated must denote an entity

We could address the error by changing line 5 to use Extent(This'Old) instead, but there is a potential performance difference between Extent(This)'Old and Extent(This'Old). With the former, only the result of the function call is copied, whereas with the latter, the value of the parameter is copied. Copying the parameter could take significant time and space if This is a large object. Of course, if the function returns a large value the copy will be large too, but in this case Extent only returns an integer.

In SPARK, unlike Ada, preconditions, postconditions, and assertions in general are verified statically, prior to execution, so there is no performance issue. Ultimately, though, the application will be executed. Having statically proven the preconditions and postconditions successfully, we can safely deploy the final executable without them enabled, but not all projects follow that approach (at least, not on that basis). Therefore, for the sake of emphasizing the idiom with typically better performance, we prefer applying ‘Old to the function in our implementation.

We can tell GNATprove that this is a benign case, using a pragma in the package spec:

pragma Unevaluated_Use_of_Old (Allow);

GNATprove will then allow use of ‘Old on the call to function Extent and will ensure that no exceptions will be raised by the function.

As with procedure Push, we can also use Top_Element to strengthen the postcondition for procedure Pop (line 4 below):

1 procedure Pop (This : in out Stack;  Item : out Element) with
2   Pre    => not Empty (This),
3   Post   => not Full (This)
4             and Item = Top_Element (This)'Old
5             and Extent (This) = Extent (This)'Old – 1,
6   Global => null;

Line 4 states that the Item returned in the parameter to Pop is the value that would be returned by Top_Element prior to the call to Pop.

One last significant enhancement now remains to be made. Consider the assertions in the main procedure about the effects of Pop on lines 24 and 25, repeated below:

21    Pop (S1, X);
22    Put_Line ("Top of S1 is '" & Top_Element (S1) & "'");
23    Pop (S1, Y);
24    pragma Assert (X = 'b');
25    pragma Assert (Y = 'a');

Previous lines had pushed ‘a’ and then ‘b’ in that order onto S1. GNATprove emits this one message:

25:19: medium: assertion might fail, cannot prove Y = 'a' (e.g. when Y = 'b')

The message is about the assertion on line 25, alone. The assertion on line 24 was verified. Also, the message indicates that Y could be some arbitrary character. We can conclude that the provers do not know enough about the state of the stack after a call to Pop. The postcondition requires strengthening.

The necessary postcondition extension reflects a unit-level functional requirement for both Push and Pop. If one considers that postconditions correspond to the low-level unit functional requirements (if not more), one can see why the postconditions must be complete. Identifying and expressing complete functional requirements is difficult in itself, and indeed the need for this additional postcondition content is not obvious at first.

The unit-level requirement for both operations is that the prior array components within the stack are not altered, other than the one added or removed. We need to state that Push and Pop have not reordered them, for example. Specifically, for Push we need to say that the new stack state has exactly the same prior array slice contents, ignoring the newly pushed value. For Pop, we need to say that the new state has exactly the prior array slice contents without the old value at the top.

A new function can be used to express these requirements for both Push and Pop:

function Unchanged (Invariant_Part, Within : Stack) return Boolean;

The Within parameter is a stack whose internal state will be compared against that of the Invariant_Part parameter. The name “Invariant_Part” is chosen to indicate the stack state that has not changed. The name "Within" is chosen for readability in named parameter associations on the calls. For example:

Unchanged (X, Within => Y)

means that the Element values of X should be equal to precisely the corresponding values within Y.

However, this function is not one that users would call directly. We only need it for proof. Therefore, we mark the Unchanged function as a "ghost" function so that the compiler will neither generate code for it nor allow the application code to call it. The function is declared with that aspect (on line 2) as follows:

1 function Unchanged (Invariant_Part, Within : Stack) return Boolean
2   with Ghost;

Key to the usage is the fact that by passing This'Old and This to the two parameters we can compare the before/after states of a single object. Viewing the function's implementation will help understand its use in the postconditions:

1 function Unchanged (Invariant_Part, Within : Stack) return Boolean is
2   (Invariant_Part.Top <= Within.Top and then
3    (for all K in 1 .. Invariant_Part.Top =>
4        Within.Values (K) = Invariant_Part.Values (K)));

This approach is based directly on a very clever one by Rod Chapman, as seen in some similar code.

The function states that the array components logically contained in Invariant_Part must have the same values as those corresponding array components in Within. Note how we allow Invariant_Part to contain fewer values than the other stack (line 2 above). That is necessary because we use this function in the postconditions for both the Push and Pop operations, in which one more or one less Element value will be present, respectively.

For Push, we add a call to the function in the postcondition as line 6, below:

1 procedure Push (This : in out Stack;  Item : Element) with
2   Pre    => not Full (This),
3   Post   => not Empty (This)
4             and then Top_Element (This) = Item
5             and then Extent (This) = Extent (This)'Old + 1
6             and then Unchanged (This'Old, Within => This),
7   Global => null;

This'Old provides the value of the stack prior to the call of Push, without the new value included, whereas This represents the stack state after Push returns, with the new value in place. Thus, the prior values are compared to the corresponding values in the new state, with the newly included value ignored.

Likewise, we add the function call to the postcondition for Pop, also line 6, below:

1 procedure Pop (This : in out Stack;  Item : out Element) with
2   Pre    => not Empty (This),
3   Post   => not Full (This)
4             and Item = Top_Element (This)'Old
5             and Extent (This) = Extent (This)'Old - 1
6             and Unchanged (This, Within => This'Old),
7   Global => null;

In contrast with procedure Push, on line 6 the values This and This'Old are passed to the opposite parameters. In this case the new state of the stack, with one less array component logically present, is used as the invariant to compare against. Line 6 expresses the requirement that the new state's content is the same as the old state's content except for the one array component no longer present. Because the function only compares the number of array components within the Invariant_Part, the additional top element value within This'Old is ignored.

Note that we must apply ‘Old to This in the calls to Unchanged in both procedures, rather than to some function result. That is unavoidable because we must refer to the prior state of the one stack object being compared.

With those additions to the postconditions we get no further messages from GNATprove from the main procedure, including assertions about the states resulting from a series of calls. We have achieved the Gold level.

Some additional postconditions are possible, however, for completeness. We can also use function Unchanged in a new postcondition for the "=" function:

1 function "=" (Left, Right : Stack) return Boolean with
2    Post => "="'Result = (Extent (Left) = Extent (Right)
3                          and then Unchanged (Left, Right));

This postcondition expresses an implication: whenever the “=” function comparing the two stacks returns True, the Extent (i.e., Top) values will be the same and Unchanged will hold. In other words, they will have the same logical size and content. Whenever “=” returns False, the conjunction will not hold either. Note that on line 3, neither argument to function Unchanged has ‘Old applied because we are comparing two distinct stack objects, rather than different states for one object. The sizes will be the same (from line 2) so Unchanged will compare the entire slices logically contained by Left and Right.

We can use the same implication approach in a new postcondition for function Empty:

function Empty (This : Stack) return Boolean with
Post => Empty'Result = (Extent (This) = 0);

Whenever Empty returns True, Top (i.e., Extent) will be zero, otherwise Top will not be zero.

4.6 Platinum Implementation

Our Gold level implementation also achieved the Platinum level because our postconditions fully covered the functional requirements and there were no abstract properties to be proven. Achieving the Platinum level is rare in itself, all the more so using the Gold level implementation. Doing so is possible in no small part because stacks are simple abstractions.

5. Concluding Remarks

We have shown how to transition an Ada implementation of a sequential, bounded stack abstract data type into a SPARK implementation supporting formal proof of the abstraction’s semantics. The full project, including sources for each level, are available on GitHub.

Overall, the changes were relatively simple and brief. The truly difficult part of the effort, of course, was determining what changes to make in order to satisfy the provers. That difficulty is somewhat understated in the text because we go directly from specific problems to their solutions, without indicating the time and effort required to identify those solutions. Similarly, we elided parts of the GNATprove messages to highlight the parts indicating the actual problem. Knowing how to interpret the messages, the counterexamples, and possible explanations is a skill that comes with experience.

In addition, we must point out that stacks are simple, especially bounded stacks based on arrays. The relative ease in reaching the Gold or Platinum levels would likely not be possible for other data structures. In particular, a “model” of the abstraction’s state will often be required, resulting in complexity well beyond the Unchanged function that was sufficient for bounded stacks. See, for example, the formal containers shipped with GNAT.

Thanks are due to Yannick Moy and the entire SPARK team at AdaCore for their essential help.

6. Gold/Platinum Implementation Listing

The following is the generic package declaration and body for the Platinum level implementation. As described earlier, the Platinum level implementation is the same as the Gold level implementation. We have kept the two versions in separate packages and files.

Rather than using the "_Platinum" suffix in this unit name, we use the name shown below because this is the final, production-ready version and, as such, should include the indicator of whether it is thread-safe (it is not).

The Platinum version, like the Gold version, did not include the Depends contracts. In the source directory we include a version with those contracts, for completeness.

generic
type Element is private;
--  The type of values contained by objects of type Stack

Default_Value : Element;
--  The default value used for stack contents. Never
--  acquired as a value from the API, but required for
--  initialization in SPARK.
package Sequential_Bounded_Stacks is

pragma Unevaluated_Use_of_Old (Allow);

subtype Element_Count is Integer range 0 .. Integer'Last - 1;
--  The number of Element values currently contained
--  within any given stack. The lower bound is zero
--  because a stack can be empty. We limit the upper
--  bound (minimally) to preclude overflow issues.

subtype Physical_Capacity is
Element_Count range 1 .. Element_Count'Last;
--  The range of values that any given stack object can
--  specify (via the discriminant) for the number of
--  Element values the object can physically contain.
--  Must be at least one.

type Stack (Capacity : Physical_Capacity) is private
with Default_Initial_Condition => Empty (Stack);

procedure Push (This : in out Stack;  Item : Element) with
Pre    => not Full (This),
Post   => not Empty (This)
and then Top_Element (This) = Item
and then Extent (This) = Extent (This)'Old + 1
and then Unchanged (This'Old, Within => This),
Global => null;

procedure Pop (This : in out Stack;  Item : out Element) with
Pre    => not Empty (This),
Post   => not Full (This)
and Item = Top_Element (This)'Old
and Extent (This) = Extent (This)'Old - 1
and Unchanged (This, Within => This'Old),
Global => null;

function Top_Element (This : Stack) return Element with
Pre    => not Empty (This),
Global => null;
--  Returns the value of the Element at the "top" of This
--  stack, i.e., the most recent Element pushed. Does not
--  remove that Element or alter the state of This stack
--  in any way.

overriding function "=" (Left, Right : Stack) return Boolean with
Post   => "="'Result = (Extent (Left) = Extent (Right)
and then Unchanged (Left, Right)),
Global => null;

procedure Copy (Destination : in out Stack; Source : Stack) with
Pre    => Destination.Capacity >= Extent (Source),
Post   => Destination = Source,
Global => null;
--  An alternative to predefined assignment that does not
--  copy all the values unless necessary. It only copies
--  the part "logically" contained, so is more efficient
--  when Source is not full.

function Extent (This : Stack) return Element_Count with
Global => null;
--  Returns the number of Element values currently
--  contained within This stack.

function Empty (This : Stack) return Boolean with
Post   => Empty'Result = (Extent (This) = 0),
Global => null;

function Full (This : Stack) return Boolean with
Post   => Full'Result = (Extent (This) = This.Capacity),
Global => null;

procedure Reset (This : in out Stack) with
Post   => Empty (This),
Global => null;

function Unchanged (Invariant_Part, Within : Stack) return Boolean
with Ghost;
--  Returns whether the Element values of Invariant_Part
--  are unchanged in the stack Within, e.g., that inserting
--  or removing an Element value does not change the other
--  Element values held.

private

type Content is array (Physical_Capacity range <>) of Element;

type Stack (Capacity : Physical_Capacity) is record
Values : Content (1 .. Capacity) := (others => Default_Value);
Top    : Element_Count := 0;
end record with
Predicate => Top in 0 .. Capacity;

------------
-- Extent --
------------

function Extent (This : Stack) return Element_Count is
(This.Top);

-----------
-- Empty --
-----------

function Empty (This : Stack) return Boolean is
(This.Top = 0);

----------
-- Full --
----------

function Full (This : Stack) return Boolean is
(This.Top = This.Capacity);

-----------------
-- Top_Element --
-----------------

function Top_Element (This : Stack) return Element is
(This.Values (This.Top));

---------
-- "=" --
---------

function "=" (Left, Right : Stack) return Boolean is
(Left.Top = Right.Top and then
Left.Values (1 .. Left.Top) = Right.Values (1 .. Right.Top));

---------------
-- Unchanged --
---------------

function Unchanged (Invariant_Part, Within : Stack) return Boolean is
(Invariant_Part.Top <= Within.Top and then
(for all K in 1 .. Invariant_Part.Top =>
Within.Values (K) = Invariant_Part.Values (K)));

end Sequential_Bounded_Stacks;

The package body:

package body Sequential_Bounded_Stacks is

-----------
-- Reset --
-----------

procedure Reset (This : in out Stack) is
begin
This.Top := 0;
end Reset;

----------
-- Push --
----------

procedure Push (This : in out Stack; Item : in Element) is
begin
This.Top := This.Top + 1;
This.Values (This.Top) := Item;
end Push;

---------
-- Pop --
---------

procedure Pop (This : in out Stack; Item : out Element) is
begin
Item := This.Values (This.Top);
This.Top := This.Top - 1;
end Pop;

----------
-- Copy --
----------

procedure Copy (Destination : in out Stack; Source : Stack) is
subtype Contained is Element_Count range 1 .. Source.Top;
begin
Destination.Top := Source.Top;
Destination.Values (Contained) := Source.Values (Contained);
end Copy;

end Sequential_Bounded_Stacks;
]]>
An Introduction to Contract-Based Programming in Ada https://blog.adacore.com/the-case-for-contracts Tue, 21 Apr 2020 08:26:00 -0400 Abe Cohen https://blog.adacore.com/the-case-for-contracts

One of the most powerful features of Ada 2012* is the ability to specify contracts on your code. Contracts describe conditions that must be satisfied upon entry (preconditions) and upon exit (postconditions) of your subprogram. Preconditions describe the context in which the subprogram must be called, and postconditions describe conditions that will be adhered to by the subprogram’s implementation. If you think about it, contracts are a natural evolution of Ada’s core design principle. To encourage developers to be as explicit as possible with their expressions, putting both the compiler/toolchain and other developers in the best position to help them develop better code.

The addition of contracts into a standard Ada application accomplishes several elusive objectives; specifically, they act as a static method of handling potential errors, as documentation that gets updated and checked for consistency by the compiler alongside your code, and provide static analysis tools like SPARK and CodePeer with more application-specific detail they can use to produce higher-quality results. So let’s get started.

package Graph is

type Graph_Record (Nodes : Positive) is record
Node_List : Node_List_Type (1 .. Nodes);
end record;
procedure Set_Source (Graph : in out Graph_Record; ID : Positive);
end Graph;

package body Graph is
procedure Set_Source (Graph : in out Graph_Record; ID : Positive) is
begin
Graph.Node_List (ID).dist := 0;
end Set_Source;
end Graph;

Here is a package with a simple subprogram that sets a property of a graph. One thing to notice about the graph from its definition is that its nodes are labelled with IDs from 1 to the number of nodes. In order to make sure that our subprogram doesn’t index into the graph’s list of nodes out of bounds, we might do a number of things. We can change Set_Source to a function that returns a boolean - True if the operation was successful, False if the supplied ID is out of range. Another option is to do nothing and make use of the default compiler-inserted array access check (I'll get into the drawbacks of this later), or we can even insert an explicit defensive check of our own if we want to raise a specific exception with a specific message.

However, all of these approaches come with two fundamental issues: they require additional documentation to be effective, and they rely on checks and/or exception handlers at run-time to prevent errors which can hurt performance. By adding a simple precondition, we can mitigate both of these problems at the same time.

procedure Set_Source (Graph : in out Graph_Record; ID : Positive)
with Pre => (ID <= Graph.Nodes);

The documentation issue is more obvious, so I’ll address that one first. Anyone using this API, even someone without access to the implementation, now knows that this subprogram expects to be called with the ID parameter in a specific range, yet no additional documentation is needed to express this. If we were using conventional methods, we would need another way to tell API users how to correctly use this subprogram. However, using contracts in this manner integrates the task of writing and updating documentation with the subprogram’s design process. On top of that, if the subprogram were to be redesigned, say if the Graph record type was broadened to accept characters as indices for Node_List, those new requirements would be reflected in the new preconditions, with no additional information needed.

In addition to helping other developers use your subprograms properly, contracts introduce a static methodology for dealing with errors. Conventionally, errors are dealt with via defensive checks and exception handlers at run-time. Particularly in an embedded context, where the final executable size in memory and computational demands need to be optimized, the reduction of run-time code is essential to dealing with hardware constraints. Accordingly, many programs have no choice but to trust that their testing infrastructure was sufficient and ship code with most run-time checks turned off. However, it’s not revolutionary to say that all programs wish their applications ran safely with less overhead. Using contracts provides an elegant way for developers higher in the call chain to take appropriate action to avoid violating known conditions that will cause program failure without adding run-time code at every level, as would happen with either explicit or compiler-inserted defensive checks or propagating exception handlers.

Sometimes though, as in the case of input validation, there’s no way to get around defensive code at run-time. Contracts provide the flexibility to add these checks both broadly and on a granular level. If you pass the ‘-gnata’ switch to the compiler, it will insert additional checks assuring your contracts are not violated alongside the standard Ada run-time checks, like range checks on types. However, if you just want to enable a single defensive check, you can do something like this:

pragma Assertion_Policy (Pre => Check, Post => Ignore);
procedure Set_Source (Graph : out Graph_Record; ID : Positive) is
begin
Graph.Node_List (ID).dist := 0;
end;

The use of contracts can also increase organizational confidence that testing was in fact sufficient, and accounted for all the potential ways in which the application could fail. If you’re not at the level of statically verifying contracts to be unbreakable within the context of your application with SPARK, other static analysis tools, like CodePeer, can benefit from the extra information contracts provide about the intended use of your code. This is because in this context, contracts are language-level proxies of your application’s requirements, and CodePeer, like many other tools, only works on language-level constructs.

When CodePeer analyzes a subprogram, it generates implicit pre- and postconditions as part of the analysis. If one of those implicit contracts might be violated, you might get a message like this:

medium: precondition (array index check) might fail on call to graph.set_source: requires ID <= Graph.Nodes

However, when you supply CodePeer with your own contracts to compare against,  it can output situations in which user-supplied contracts contradict some of its own, leading to more specific, more actionable findings, and fewer false positives. To learn more about contracts, check out this chapter from learn.adacore.com, or this section of the SPARK documentation.

*Contracts can also be used via pragma Precondition and pragma Postcondition with older versions of GNAT, or approximated with pragma Assert as defined in Ada 05. Learn more about that here.

]]>
Ada on the ESP8266 https://blog.adacore.com/ada-on-the-esp8266 Thu, 09 Apr 2020 07:31:00 -0400 Johannes Kliemann https://blog.adacore.com/ada-on-the-esp8266 $llvm-gnatmake -c unit.adb -cargs --target=xtensa -mcpu=esp8266 $ cd /path/to/Arduino/libraries
$git clone --recursive https://github.com/jklmnn/esp8266-ada-example.git$ cd esp8266-ada-example
$make $ screen /dev/ttyUSB0 115200
Make with Ada!
]]>
A Trivial File Transfer Protocol Server written in Ada https://blog.adacore.com/the-elegance-of-open-source-collaboration Tue, 07 Apr 2020 07:50:00 -0400 Martyn Pike https://blog.adacore.com/the-elegance-of-open-source-collaboration

For an upcoming project, I needed a simple way of transferring binary files over an Ethernet connection with minimal (if any at all) user interaction.

A protocol that's particularly appropriate for this kind of usage is the Trivial File Transfer Protocol (TFTP).  You can find a high level description on Wikipedia and a more detailed breakdown of the protocol here.

My previous experience with this protocol has mostly been within test rig environments,  where a target computer accesses its operational software payload from a TFTP server at boot-time.  The beauty of this is approach is that it allows for different payloads to be used between reboots by switching the files being served.  That is exactly how this server will be used in my forthcoming project, also to be documented on blog.adacore.com.

The Ada TFTP server will be hosted on Ubuntu Linux and support a subset of the transactions provided by the protocol.  For example, the ability for the client to write a file to the server will not be supported.

There are a number of TFTP servers available for Ubuntu Linux. However I wanted to implement my own in Ada, mainly to prove it could be done, but also to test a couple of different options for handling UDP/IP transactions from within Ada applications.

To do this, I needed something to mimic the capabilities of GNAT.Sockets.

To start with, I reviewed the current catalogue of available Ada software on Github and the repository from my good friends over at CodeLabs, who happen to be the developers of the Muen separation kernel which will also figure in my forthcoming project.

The CodeLabs team had exactly what I was looking for: Anet.  I highly recommend that you review its code on the CodeLabs repository.   That repository is now my go-to for (publicly available and open source) high quality examples of Ada and SPARK software development.  Many of my future blog posts about applying the AdaCore tools and techniques will be oriented around examples based on CodeLabs source code.

Back to my TFTP server.  Since I had two options for the UDP/IP transaction functionality, I decided to set about creating one version of my server using GNAT.Sockets and another using Anet from CodeLabs.

If you want to get ahead of the game, the code is available on my GitHub repository for all three parts of this blog series.

The first step towards my objective was to obtain Anet, by cloning the repository, building the library and installing it in a location where the GNAT build tools can locate it.

All the code for this blog post can be built with GNAT Community 2019 as well as GNAT Pro.

The TFTP server code has been reviewed by CodePeer for the detection of run-time vulnerabilities that may lead to unexpected code execution paths and by GNATcheck against a suitable coding standard.

Both CodePeer and GNATcheck are available from AdaCore as professionally assured products and can be qualified as TQL-5 review tools for use by DO-178B/C projects.

The following command sequence installs the Anet libraries into ~/sw/adalibs and with a suitable GNAT compiler in my PATH,  I would execute the following commands:

git clone https://git.codelabs.ch/anet.git
cd anet
git checkout master
make all
make PREFIX=~/sw/adalibs install

After doing this, there will be a shared library called libanet.so stored in the ~/sw/adalibs/lib directory.

Before writing code that will use this library,  the GPR_PROJECT_PATH environment variable needs to identify the ~/sw/adalibs/lib/gnat directory.

This can be done by using the following

export GPR_PROJECT_PATH=~/sw/adalibs/lib/gnat

I encountered a slight learning curve with the Anet API because it's architecture differs from that of GNAT.Sockets.  However, the code documentation is very good and before I knew it I had a proof of concept working.

To build the code from Github (with the same GNAT compiler in the path which was used to build Anet), you can use the following:

export GPR_PROJECT_PATH=~/sw/adalibs/lib/gnat
make all

The makefile and GNAT project file are as follows:

# Assumes GPR_PROJECT_PATH includes Anet installation
# Try to use the same GNAT to build adatftpd that was used to
# build Anet.
all:

check:
gnatcheck -P adatftpd.gpr --show-rule -rules -from=gnatcheck.rules

review:
codepeer -P adatftpd.gpr -level 2 -output-msg

clean:
gprclean -q -P adatftpd.gpr
with "anet.gpr";

for Source_Dirs use ("src/**");
for Object_Dir use "obj";
for Exec_Dir use "test";

package Builder is
end Builder;

package Compiler is
end Compiler;

end Adatftpd;

Assuming no errors occurred during this sequence of commands, the 'test' sub-directory will contain the 'adatftpf-anet' executable.

You can also checkout the verification program I wrote for my TFTP server,  which is also available on my Github repository.

I'd welcome feedback and collaboration on either of these TFTP related projects.

]]>
Proving properties of constant-time crypto code in SPARKNaCl https://blog.adacore.com/proving-constant-time-crypto-code-in-sparknacl Thu, 02 Apr 2020 08:15:00 -0400 Roderick Chapman https://blog.adacore.com/proving-constant-time-crypto-code-in-sparknacl #define FOR(i,n) for (i = 0; i < n; ++i) #define sv static void typedef unsigned char u8; typedef long long i64; typedef i64 gf[16];
sv pack25519(u8 *o, const gf n);
subtype I32 is Integer_32;
subtype N32 is I32 range 0 .. I32'Last;
subtype I64 is Integer_64;

subtype Index_32 is I32 range 0 .. 31;

type Byte_Seq is array (N32 range <>) of Byte;
subtype Bytes_32 is Byte_Seq (Index_32);

--  "LM"   = "Limb Modulus"
--  "LMM1" = "Limb Modulus Minus 1"
LM   : constant := 65536;
LMM1 : constant := 65535;
--  "R2256" = "Remainder of 2**256 (modulo 2**255-19)"
R2256 : constant := 38;

--  "Maximum GF Limb Coefficient"
MGFLC : constant := (R2256 * 15) + 1;

--  "Maximum GF Limb Product"
MGFLP : constant := LMM1 * LMM1;

subtype GF_Any_Limb is I64 range -LM .. (MGFLC * MGFLP);

type GF is array (Index_16) of GF_Any_Limb;

subtype GF_Normal_Limb is I64 range 0 .. LMM1;

subtype Normal_GF is GF
with Dynamic_Predicate =>
(for all I in Index_16 => Normal_GF (I) in GF_Normal_Limb);
--  Reduces N modulo (2**255 - 19) then packs the
--  value into 32 bytes little-endian.
function Pack_25519 (N : in Normal_GF) return Bytes_32
with Global => null;
sv pack25519 (u8 *o, const gf n)
{
int i, j, b;
gf m, t;
FOR(i,16) t[i]=n[i];
car25519(t);
car25519(t);
car25519(t);
FOR(j, 2) {
m[0]=t[0]-0xffed;
for(i=1;i<15;i++) {
m[i]=t[i]-0xffff-((m[i-1]>>16)&1);
m[i-1]&=0xffff;
}
m[15]=t[15]-0x7fff-((m[14]>>16)&1);
b=(m[15]>>16)&1;
m[14]&=0xffff;
sel25519 (t, m, 1-b);
}
FOR(i, 16) {
o[2*i]=t[i]&0xff;
o[2*i+1]=t[i]>>8;
}
}
sv sel25519 (gf p, gf q, int b);
--  Constant time conditional swap of P and Q.
procedure CSwap (P    : in out GF;
Q    : in out GF;
Swap : in     Boolean)
with Global => null,
Contract_Cases =>
(Swap     => (P = Q'Old and Q = P'Old)
not Swap => (P = P'Old and Q = Q'Old));
if Swap then
Temp := P;
P := Q;
Q := Temp;
end if;
sv sel25519 (gf p, gf q, int b)
{
i64 t, i, c = ~(b-1);
FOR(i, 16) {
t= c&(p[i]^q[i]);
p[i]^=t;
q[i]^=t;
}
}
type Bit_To_Swapmask_Table is array (Boolean) of U64;
(False => 16#0000_0000_0000_0000#,
True  => 16#FFFF_FFFF_FFFF_FFFF#);
pragma Assume
(for all K in I64 => To_I64 (To_U64 (K)) = K);
procedure CSwap (P    : in out GF;
Q    : in out GF;
Swap : in     Boolean)
is
T : U64;
C : U64 := Bit_To_Swapmask (Swap);
begin
for I in Index_16 loop
T := C and (To_U64 (P (I)) xor To_U64 (Q (I)));
P (I) := To_I64 (To_U64 (P (I)) xor T);
Q (I) := To_I64 (To_U64 (Q (I)) xor T);

pragma Loop_Invariant
(if Swap then
(for all J in Index_16 range 0 .. I =>
(P (J) = Q'Loop_Entry (J) and
Q (J) = P'Loop_Entry (J)))
else
(for all J in Index_16 range 0 .. I =>
(P (J) = P'Loop_Entry (J) and
Q (J) = Q'Loop_Entry (J)))
);
end loop;
end CSwap;
--  Subtracting P twice from a Normal_GF might result
--  in a GF where limb 15 can be negative with lower bound -65536
subtype Temp_GF_MSL is I64 range -LM .. LMM1;
subtype Temp_GF is GF
with Dynamic_Predicate =>
(Temp_GF (15) in Temp_GF_MSL and
(for all K in Index_16 range 0 .. 14 =>
Temp_GF (K) in GF_Normal_Limb));

procedure Subtract_P (T         : in     Temp_GF;
Result    :    out Temp_GF;
Underflow :    out Boolean)
with Global => null,
Pre    => T (15) >= -16#8000#,
Post   => (Result (15) >= T (15) - 16#8000#);
subtype I64_Bit is I64 range 0 .. 1;

procedure Subtract_P (T         : in     Temp_GF;
Result    :    out Temp_GF;
Underflow :    out Boolean)
is
Carry : I64_Bit;
R     : GF;
begin
R := (others => 0);

--  Limb 0 - subtract LSL of P, which is 16#FFED#
R (0) := T (0) - 16#FFED#;

--  Limbs 1 .. 14 - subtract FFFF with carry
for I in Index_16 range 1 .. 14 loop
Carry     := ASR_16 (R (I - 1)) mod 2;
R (I)     := T (I) - 16#FFFF# - Carry;
R (I - 1) := R (I - 1) mod LM;

pragma Loop_Invariant
(for all J in Index_16 range 0 .. I - 1 =>
R (J) in GF_Normal_Limb);
pragma Loop_Invariant (T in Temp_GF);
end loop;

--  Limb 15 - Subtract MSL (Most Significant Limb)
--  of P (16#7FFF#) with carry.
--  Note that Limb 15 might become negative on underflow
Carry  := ASR_16 (R (14)) mod 2;
R (15) := (T (15) - 16#7FFF#) - Carry;
R (14) := R (14) mod LM;

--  Note that R (15) is not normalized here, so that the
--  result of the first subtraction is numerically correct
--  as the input to the second.
Underflow := R (15) < 0;
Result    := R;
end Subtract_P;
function Pack_25519 (N : in Normal_GF) return Bytes_32
is
L      : GF;
R1, R2 : Temp_GF;
First_Underflow  : Boolean;
Second_Underflow : Boolean;
begin
L := N;
Subtract_P (L,  R1, First_Underflow);
Subtract_P (R1, R2, Second_Underflow);
CSwap (R1, R2, Second_Underflow);
CSwap (L,  R2, First_Underflow);
end Pack_25519;
sparknacl-utils.adb:197:27: medium: predicate check might fail
--  Result := T - P;
--  if     Underflow, then Result is not a Normal_GF
--  if not Underflow, then Result is     a Normal_GF
procedure Subtract_P (T         : in     Temp_GF;
Result    :    out Temp_GF;
Underflow :    out Boolean)
with Global => null,
Pre    => T (15) >= -16#8000#,
Post   => (Result (15) >= T (15) - 16#8000#) and then
(Underflow /= (Result in Normal_GF));
R (14) := R (14) mod LM;
R (15) := R (15) mod LM;
sparknacl-utils.adb:139:23: medium: predicate check might fail
]]>
Time travel debugging in GNAT Studio with GDB and RR https://blog.adacore.com/time-travel-debugging-in-gnat-studio Tue, 17 Mar 2020 09:18:07 -0400 Ghjuvan Lacambre https://blog.adacore.com/time-travel-debugging-in-gnat-studio
with Ada.Numerics.Discrete_Random;

procedure Main is

Generator : Rand_Positive.Generator;

Error : exception;

Bug : Boolean := False;

procedure Make_Bug is
begin
Bug := True;
end Make_Bug;

procedure Do_Bug is
begin
Bug := True;
end Do_Bug;

begin
Rand_Positive.Reset(Generator);

for I in 1..10 loop
if Rand_Positive.Random(Generator) < (Positive'Last / 100) then
if Rand_Positive.Random(Generator) < (Positive'Last / 2) then
Make_Bug;
else
Do_Bug;
end if;
end if;
end loop;

if Bug then
raise Error;
end if;

end Main;
]]>

Having previously shown how to create a Web application in Ada, it's not so difficult to create an Android application in Ada. Perhaps the simplest way is to install Android Studio. Then just create a new project and choose "Empty Activity". Open the layout, delete TextView and put WebView instead.

In onCreate function write the initialization code:

WebView webView = (WebView)findViewById(R.id.webView);
WebSettings settings = webView.getSettings();
settings.setJavaScriptEnabled(true);

To make WebView work offline, you need to provide content. One way to do this is just to put content in the asset folder and open it as a URL in WebView. When a user starts the application, WebView will load HTML and corresponding JavaScript. Then JavaScript loads WebAssembly and so, actually, launches Ada code. But it can't use a file:/// schema to load JavaScript and WebAssembly files because of the default security settings. So we trick WebView by intercepting requests and also provide correct MIME types for them. We do this using the shouldInterceptRequest method of WebViewClient class to intercept any request to HTML/WASM/JS/JPEG resources and load the corresponding file from the asset folder:

public WebResourceResponse shouldInterceptRequest(WebView view,
WebResourceRequest request) {
String path = request.getUrl().getLastPathSegment();

try {
String mime;
AssetManager assetManager = getAssets();

if (path.endsWith(".html")) mime = "text/html";
else if (path.endsWith(".wasm")) mime = "application/wasm";
else if (path.endsWith(".mjs")) mime = "text/javascript";
else if (path.endsWith(".jpg")) mime = "image/jpeg";
else
return super.shouldInterceptRequest(view, request);

InputStream input = assetManager.open("www/" + path);

return new WebResourceResponse(mime, "utf-8", input);
} catch (IOException e) {
e.printStackTrace();
ByteArrayInputStream result = new ByteArrayInputStream
(("X:" + path + " E:" + e.toString()).getBytes());
return new WebResourceResponse("text/plain", "utf-8", result);
}
}

Now connect this code to WebView, like this:

webView.setWebViewClient(new WebViewClient() {
@Override
public WebResourceResponse shouldInterceptRequest(WebView view,
....
});

For debug purposes, let's connect the WebView console to Android log. We just add this function below the code for shouldInterceptRequest:

public boolean onConsoleMessage(ConsoleMessage cm) {
Log.d("MyApplication", cm.message() + " -- From line "
+ cm.lineNumber() + " of "
+ cm.sourceId() );
return true;
}

Now we're able to build and run an Android Package. Here is how it looks like on Android Studio emulator (it's been tested on my phone too!):

If you need the complete code, there's a repository on github!

PS: This article doesn't discuss how we produced WebAssembly from Ada code for running with WebGL integration. We will write a follow-up post about that soon!

]]>

As a demonstration for the use of Ada and SPARK in very small embedded targets, I created a remote-controlled (RC) car using Lego NXT Mindstorms motors and sensors but without using the Lego computer or Lego software. I used an ARM Cortex System-on-Chip board for the computer, and all the code -- the control program, the device drivers, everything -- is written in Ada. Over time, I’ve upgraded some of the code to be in SPARK. This blog post describes the hardware, the software, the SPARK upgrades, and the repositories that are used and created for this purpose.

Why use Lego NXT parts? The Lego NXT robotics kit was extremely popular. Many schools and individuals still have kits and third-party components. Even if the latest Lego kit is much more capable, the ubiquity and low cost of the NXT components make them an attractive basis for experiments and demonstrations.

In addition, there are many existing NXT projects upon which to base demonstrations using Ada. For example, the RC car is based on the third-party HiTechnic IR RC Car, following instructions available here: http://www.hitechnic.com/models. The car turns extremely well because it has an Ackerman steering mechanism, so that the inside wheel turns sharper than the outside wheel, and a differential on the drive shaft so that the drive wheels can rotate at different speeds during a turn. The original car uses the HiTechnic IR (infra-red) receiver to communicate with a Lego remote control. This new car uses that same receiver and controller, but also supports another controller communicating over Bluetooth LE.

Replacing the NXT Brick

The NXT embedded computer controlling NXT robots is known as the “brick,” probably because of its appearance. (See Figure 1.) It consists of an older 48 MHz ARMv7, with 256 KB of FLASH and 64 KB of RAM, as well as an AVR co-processor. The brick enclosure provides an LCD screen, a speaker, Bluetooth, and four user-buttons, combined with the electronics required to interface to the external world. A battery pack is on the back.

Our replacement computer is one of the “Discovery Kit” products from STMicroelectronics. The Discovery Kits have ARM Cortex processors and include many on-package devices for interfacing to the external world, including A/D and D/A converters, timers, UARTs, DMA controllers, I2C and SPI communication, and others. Sophisticated external components are also included, depending upon the specific kit.

Specifically, we use the STM32F4 Discovery Kit which has a Cortex M4 MCU running at up to 168 MHz, a floating-point co-processor, a megabyte of FLASH and 192 KB of RAM. It also includes an accelerometer, MEMS microphone, audio codec, a user button, and four user LEDs. (See figure 2.) It is very inexpensive– approximately 15. Details are available here: https://www.st.com/en/evaluation-tools/stm32f4discovery.html I made one change to the Discovery Kit board as received from the factory. Because the on-package devices, such as the serial ports, I2C devices, timers, etc. all share potentially overlapping groups of GPIO pins, and because not all pins are available on the headers, not all the pins required were exclusively available for all the devices needed for the RC car. Ultimately, I found a set of pin allocations that would almost work, but I needed pin PA0 to do it. However, pin PA0 is dedicated to the blue User button by a solder bridge on the underside of the board. I removed that solder bridge to make PA0 available. Of course, doing so disabled the blue User button but I didn’t need it for this project. Replacing the NXT brick also removed the internal interface electronics for the motors and sensors. I used a combination of a third-party board and hand-made circuits to replace them. A brief examination of the motors will serve to explain why the additional board was chosen. The Lego Mindstorms motors are 9-volt DC motors with a precise rotation sensor and significant gear reduction producing high torque. The motors rotate at a rate relative to the power applied and can rotate in either direction. The polarity of the power lines controls the rotation direction: positive rotates one way, negative rotates the other way. Figure 3 illustrates the partial internals of the NXT motor, including the gear train in light blue, and the rotation sensor to the left in dark blue, next to the motor itself in dark orange. (The dark gray part at far left is the connector housing.) I mentioned that the polarity of the applied power determines the rotation direction. That polarity control requires an external circuit, specifically an ‘H-bridge” circuit that allows us to achieve that effect. Figure 4 shows the functional layout of the H-bridge circuit, in particular the arrangement of the four switches S1 through S4 around the motor M. By selectively closing two switches and leaving the other two open we can control the direction of the current flow, and thereby control the direction of the motor rotation. Figure 5 illustrates two of the three useful switch configurations. The red line shows the current flow. Another option is to close two switches on the same side and end, in which case the rotor will “lock” in place. Opening all the switches removes all power and thus does not cause rotation. The fourth possible combination, in which all switches are closed, is not used. Rather than build my own H-bridge circuit I used a low-cost product dedicated to interfacing with NXT motors and sensors. In addition to the H-bridge circuits, they also provide filters for the rotation sensor’s discrete inputs so that noise does not result in too many false rotation counts. There are a number of these products available. One such is the “Arduino NXT Shield Version 2” by TKJ Electronics: http://www.tkjelectronics.dk/ in Denmark. The product is described in their blog, here: http://blog.tkjelectronics.dk/2011/10/nxt-shield-ver2/ and is available for sale here: http://shop.tkjelectronics.dk/product_info.php?products_id=29 for a reasonable price. The “NXT Shield” can control two NXT motors and one sensor requiring 9 volts input, including a Mindstorms NXT Ultrasonic Sensor. Figure 6 shows the NXT Shield with the two standard NXT connectors on the left for the two motors, and the sensor connector on the right. The kit requires assembly but it is just through-board soldering. As long as you get the diodes oriented correctly everything is straightforward. Figure 7 (below) shows our build, already located in an enclosure and connected to the Discovery Kit, power, two NXT motors, and the ultrasonic sensor. The in-coming 9 volts is routed to a DC power jack on the back of the enclosure, visible on the bottom left with red and black wires connecting it to the board. The 5 volts for the on-board electronics comes via the Discovery Kit header and is bundled with the white and green wires coming in through the left side in the figure. The enclosure itself is one of the “Make with Ada” boxes. “Make with Ada” is a competition offering serious prize money for cool projects using embedded targets and Ada. See http://www.makewithada.org/ for more information. The power supply replacing the battery pack on the back of the NXT brick is an external battery intended for charging cell phones and tablets. This battery provides separate connections for +5 and +9 (or +12) volts, which is very convenient: the +5V is provided via USB connector, which is precisely what the STM32F4 card requires, and both the NXT motors and the NXT ultrasonic sensor require +9 volts. The battery isn't light but holds a charge for a very long time, especially with this relatively light load. Note that the battery can also provide +12 volts instead of +9, selected by a physical slider switch on the side of the battery. Using +12 volts will drive the motors considerably faster and is (evidently) tolerated by the NXT Shield sensor circuit and the NXT Ultrasonic Sensor itself. Finally, I required a small circuit supporting the I2C communication with the HiTechnic IR Receiver. The circuit is as simple as one can imagine: power, ground, and a pull-up resistor for each of the two I2C communication lines. These components are housed in the traditional Altoids tin and take power and ground from the Discovery Kit header pins. The communication lines go to specific GPIO header pins. All of these replacements and the overall completed car (known as "Bob"), are shown in the following images: Figure 10 shows the rear enclosure containing the NXT Shield board, labeled “Make With Ada” on the outside, and the Altoids tin on the side containing the small circuit for the IR receiver. Here is the car in action: Replacing the NXT Software The Ada Drivers Library (ADL) provided by AdaCore and the Ada community supplies the device drivers for the timers, I2C, A/D and D/A converters, and other devices required to replace those in the the NXT brick. The ADL supports a variety of development platforms from various vendors, including the STM32 series boards. The ADL is available on GitHub for both non-proprietary and commercial use here: https://github.com/AdaCore/Ada_Drivers_Library. Replacing the brick will also require drivers for the NXT sensors and motors, software that is not included in the ADL. However, we can base them on the ADL drivers for our target board. For example, the motor rotary encoder driver uses the STM32 timer driver internally because those timers directly support quadrature rotation encoders. All these abstractions, including some that are not hardware specific, are in the Robotics with Ada repository: https://github.com/AdaCore/Robotics_with_Ada. This repo supports the NXT motors and all the basic sensors, as well as some third-party sensors. Abstract base types are used for the more complex sensors so that new sensors can be created easily using inheritance. In addition, the repository contains some signal processing and control system software, e.g., a “recursive moving average” (RMA) noise filer type and a closed loop PID controller type. These require further packages, such as a bounded ring buffer abstraction. For example, the analog sensors (e.g., the light and sound sensors), have an abstract base class controlling an ADC, and two abstract subclasses using DMA and polling to transfer the converted data. The concrete light and sound sensor types are derived from the DMA-based parent type (figure 11). The so-called NXT “digital” devices contain an embedded chip. These follow a similar design with an abstract base class and concrete subclass drivers for the more sophisticated, complex sensors. Lego refers to these sensors as “digital” sensors because they do not provide an analog signal to be sampled. Instead, the drivers both command and query the internal chips to operate the sensors. The sensors’ chips use the NXT hardware cable connectors’ two discrete I/O lines to communicate. Therefore, a serial communications protocol based on two wires is applied. This communication protocol is usually, but not always, the “I2C” serial protocol. The Lego Ultrasonic Sonar sensor and the HiTechnic IR Receiver sensor both use I2C for communication. In contrast, version 2 of the Lego Color sensor uses the two discrete lines with an ad-hoc protocol. The HiTechnic IR Receiver driver uses the I2C driver from the ADL for the on-package I2C hardware. That is a simple approach that also offloads the work from the MCU. The NXT Ultrasonic sensor, on the other hand, was a problem. I could send data to the Ultrasonic sensor successfully using the on-package I2C hardware (via the ADL driver) but could not get any data back. As discussed on the Internet, the problem is that the sensor does not follow the standard I2C protocol. It requires an extra communication line state change in the middle of the receiving steps. I could not find a way to make the on-package I2C hardware in the ARM package do this extra line change. The NXT Shield hardware even includes a GPIO “back door” connection to the I2C data line for this purpose, but I could not make that work with the STM32 hardware. Ultimately, I had to use a bit-banged approach in place of the I2C hardware and ADL driver. Fortunately, the vendor of the NXT Shield also provides the source code for an ultrasonic sensor driver in C++ using the Arduino “Wire” interface for I2C so I could see exactly what was required. Bit-banging has system-wide implications. Since the software is doing the low-level communication instead of the on-package I2C hardware, interrupting the software execution in the middle of the protocol could be a problem. That would mean that the priority of the task handing the device must be sufficiently high relative to the other tasks in the system. Bit-banging also means an additional utilization of the MPU that would otherwise be offloaded to a separate I2C hardware device. Our application is rather simple, so processor overload is not a problem. Care with the task priorities was required, though. You cannot hear the ultrasonic sensor pings, as the sensor name indicates. However, I recorded the videos on my cellphone and its microphone detects the pings. They are very directional, necessarily, so they are only heard in the video when the car is pointing at the phone. Here is another short video of the car, stationary, with the camera immediately in front. The pings are quite noticeable: System Architecture The overall architecture of the control software is shown below in figure 12. In the diagram, the parallelograms are periodic tasks (threads), running until power is removed. Each task is located inside a dedicated package. The Remote_Control package and the Vehicle package also provide functions that are callable by clients. Calls are indicated by dotted lines, with the arrowhead indicating the flow of data. For example, the Servo task in the Steering_Control package calls the Remote_Control package’s function to get the currently requested steering angle. The Steering Motor and Propulsion Motor boxes represent the two NXT motors. Each motor has a dedicated rotary encoder inside the motor housing but the diagram depicts them as distinct in order to more clearly show their usage. The PID controller and vehicle noise filter are completely in software. The PID (Proportional Integral Derivative) controller is a closed-loop control mechanism that uses feedback from the system under control to maintain a requested value. These mechanisms are ubiquitous, for example in your house's thermostat maintaining your requested heating and cooling temperatures. In our case, the PID controller maintains the requested steering angle using the steering motor's encoder data as the feedback signal. The noise filter is a “recursive moving average” filter commonly used in digital signal processing to smooth sensor inputs. (Although the third-party interface board removed most encoder noise, some noise remained.) The PID controller did not require an encoder noise filter because the mechanical steering mechanism has enough “play” in it that encoder noise has no observable effect. The vehicle measured speed calculation, however, needed the filter because the values are used only within the software, not in a physical effector. The collision detection logic determines whether a collision is imminent, using the NXT ultrasonic sensor data and the vehicle's current speed as inputs. If a sufficiently close object is detected ahead and the vehicle is moving forward, the engine Controller task stops the car immediately. Otherwise, such objects, if any, are ignored. Application Source Code Example: Steering Servo As the system diagram shows, the application consists of four primary packages, each containing a dedicated task. (There are other packages as well, but they do not contain tasks.) The task in the Steering_Control package is named “Servo” because it is acting as a servomechanism: it has a feedback control loop. In contrast, the task “Controller” in the Engine_Control package is not acting as a servo because it uses “open loop” control without any feedback. It simply sets the motor power to the requested percentage, with the resulting speed depending on the available battery power and the load on the wheels. I could also use a PID controller to maintain a requested speed, varying the power as required, but did not bother to do so in this version of the application. The source code for the “Servo” task in the Steering_Control package is shown below, along with the declarations for two subprograms called by the task. function Current_Motor_Angle (This : Basic_Motor) return Real with Inline; procedure Convert_To_Motor_Values (Signed_Power : Real; Motor_Power : out NXT.Motors.Power_Level; Direction : out NXT.Motors.Directions) with Inline, Pre => Within_Limits (Signed_Power, Power_Level_Limits); task body Servo is Next_Release : Time; Target_Angle : Real; Current_Angle : Real := 0.0; -- zero for call to Steering_Computer.Enable Steering_Power : Real := 0.0; -- zero for call to Steering_Computer.Enable Motor_Power : NXT.Motors.Power_Level; Rotation_Direction : NXT.Motors.Directions; Steering_Offset : Real; Steering_Computer : Closed_Loop.PID_Controller; begin Steering_Computer.Configure (Proportional_Gain => Kp, Integral_Gain => Ki, Derivative_Gain => Kd, Period => System_Configuration.Steering_Control_Period, Output_Limits => Power_Level_Limits, Direction => Closed_Loop.Direct); Initialize_Steering_Mechanism (Steering_Offset); Global_Initialization.Critical_Instant.Wait (Epoch => Next_Release); Steering_Computer.Enable (Current_Angle, Steering_Power); loop pragma Loop_Invariant (Steering_Computer.Current_Output_Limits = Power_Level_Limits); pragma Loop_Invariant (Within_Limits (Steering_Power, Power_Level_Limits)); Current_Angle := Current_Motor_Angle (Steering_Motor) - Steering_Offset; Target_Angle := Real (Remote_Control.Requested_Steering_Angle); Limit (Target_Angle, -Steering_Offset, Steering_Offset); Steering_Computer.Compute_Output (Process_Variable => Current_Angle, Setpoint => Target_Angle, Control_Variable => Steering_Power); Convert_To_Motor_Values (Steering_Power, Motor_Power, Rotation_Direction); Steering_Motor.Engage (Rotation_Direction, Motor_Power); Next_Release := Next_Release + Period; delay until Next_Release; end loop; end Servo; The PID controller object declared on line 19 is of a type declared in package Closed_Loop, an instantiation of a generic package. The package is a generic so that the specific floating-point input/output type is not hard-coded. The task first configures the PID controller object named Steering_Computer to specify the PID gain parameters, the interval at which the output routine is called, and the upper and lower limits for the output value (lines 21 through 27). The task then initializes the mechanical steering mechanism in order to get the steering offset (line 29). This offset is required because the steering angle requests from the user (via the remote control) are based on a frame of reference oriented on the major axis of the vehicle. Because I use the steering motor rotation angle to steer the vehicle, the code must translate the requests from the user's frame of reference (ie, the vehicle's) into the frame of reference of the steering motor. The steering motor's frame of reference is defined by the steering mechanism's physical connection to the car’s frame and is not aligned with the car’s major axis. Therefore, to do the translation the code sets the motor encoder to zero at some known point relative to the vehicle's major axis (line 29) and then handles the difference (line 38) between that motor "zero" and the "zero" corresponding to the vehicle. The code thus orients the steering motor's frame of reference to that of the vehicle, and hence to the user. Having completed these local initialization steps, the Servo task then waits for the “critical instant” in which all the tasks should begin their periodic execution (line 31). The critical instant is time T0 (usually), so the main procedure passes a common absolute time value to each task from the Epoch formal parameter to the Next_Release variable. Each task uses its local Next_Release variable to compute its next iteration release time (lines 52 and 53) using the same initial epoch time. Waiting for this critical instant release also allows each task to wait for any prior processing in the main procedure to occur. The task then enables the PID controller and goes into the loop. In each iteration, the task determines the current steering angle from the steering motor’s rotary encoder and the computed offset (line 38), gets the requested angle from the remote control and ensures it is within the steering mechanism’s physical limits (lines 40 and 41), then feeds the current angle and target angle into the PID controller (lines 43 through 46). The resulting output value is the steering motor power value required to reach the target angle. The signed steering power is then converted into an NXT motor power percentage and rotation direction (line 48). Those values are used to engage the steering motor on line 50. Finally, the task computes the next time it should be released for execution and then suspends itself until that point in time arrives (lines 52 and 53). All the tasks in the system use this same periodic looping idiom, as is expected for time-driven tasks in a Ravenscar tasking profile. (We are actually using the Jorvik tasking profile, based on Ravenscar and defined in Ada 202x. See http://www.ada-auth.org/standa...) The PID controller is based on the Arduino PID library, version 1.1.1. The primary difference between my design and the Arduino design is that this Ada version does not compute the next time to execute. Instead, because Ada has such good real-time support, barring a design error we can be sure that the periodic task will call the PID output calculation routine at a fixed rate. Therefore, the Configure routine specifies this period, which is then used internally in the output computation. In addition, the PID object does not retain pointers to the input, setpoint, and output objects, for the sake of SPARK compatibility. We pass them as parameters instead. For a great explanation of the Arduino PID design and implementation, step-by-step, see this web page: The PID controller abstract data type is declared within a generic package so that the input and output types need not be hard-coded. This specific implementation uses floating-point for the inputs and output, which gives us considerable dynamic range. The ARM MCU includes a floating-point unit so there is no performance penalty. However, if desired, a version using fixed-point types could be defined with the same API, trading some of the problems with floating point computations for problems with fixed-point computations. Neither is perfect. SPARK Upgrade One of my long terms goals for the RC Car was to upgrade to SPARK as much as possible. That effort is currently underway and some of the packages and reusable components are now in SPARK. For example, the Steering_Control package, containing the Servo task and PID controller object, are now at the Silver level of SPARK, meaning that it is proven to have no run-time errors, including no overflows. That is the reason for the loop invariants in the Servo task (lines 35 and 36 above), and the precondition on procedure Convert_To_Motor_Values (line 9 above). In particular, the provers needed to be told that the output value limits for the PID controller remain unchanged in each iteration, and that the value of the PID controller output variable remains within those limits. Other parts of the software are merely in the SPARK subset currently, but some are at the highest level. The recursive moving average (RMA) filter uses a bounded ring buffer type, for example, that is at Gold level, the level of functional proof of unit correctness. I will continue to upgrade the code to the higher levels, at least the Silver level for proving absence of runtime errors. Ultimately, however, this process will require changes to the ADL drivers because they use access discriminants which are not compatible with SPARK. That is the remaining issue preventing clean proof for the Vehicle package and its Controller task, for instance. Source Code Availability The full project for the RC car, including some relevant documents, is here: https://github.com/AdaCore/RC_... ]]> A Further Expedition into Libadalang: Save Time with Libadalang.Helpers.App https://blog.adacore.com/save-time-with-libadalang-helpers-app Thu, 06 Feb 2020 09:24:27 -0500 Pierre-Marie de Rodat https://blog.adacore.com/save-time-with-libadalang-helpers-app Martyn’s recent blog post showed small programs based on Libadalang to find uses of access types in Ada sources. Albeit short, these programs need to take care of all the tedious logistics around processing Ada sources: find the files to work on, create a Libadalang analysis context, use it to read the source files, etc. Besides, they are not very convenient to run:  gprls -s -P test.gpr | ./ptrfinder1 | ./ptrfinder2

The gprls command (shipped with GNAT Pro) is used here in order to get the list of sources that belong to the test.gpr project file. Wouldn’t it be nice if our programs could use the GNATCOLL.Projects API in order to read this project file themselves and get the list of sources to process from there? It’s definitely doable, but also definitely cumbersome: first we need to get the appropriate info from the command line (project file name, potentially target and runtime information, or a *.cgpr configuration file), then call all the various APIs to load the project, and many more operations.

Such operations are so common for tools using Libadalang that we have decided to include helpers to factor this in the library itself, so that programs can focus on their real purpose. The 20.1 Libadalang release provides building blocks to save you this trouble: check the App generic package in Libadalang.Helpers. Note that you can see a tutorial and its API reference for it in our nightly documentation.

This package is intended to be used as a framework: you instantiate it with your settings at the top-level of your program and call its Run procedure. App then takes over the control of the program: it parses command-line options and invokes when appropriate the callbacks you provided it. Let’s update Martyn’s programs to use App. The job of the first program (ptrfinder1) is to go through source files and report access type declarations and object declarations that have access types.

First, we declare some shortcuts for code brevity:

package Helpers renames Libadalang.Helpers;
package Slocs renames Langkit_Support.Slocs;

Next, we can instantiate App:

procedure Process_Unit
(Job_Ctx : Helpers.App_Job_Context; Unit : LAL.Analysis_Unit);
--  Look for the use of access types in Unit

package App is new Helpers.App
(Name         => "ptrfinder1",
Description  => "Look for the use of access types in the input sources",
Process_Unit => Process_Unit);

Naturally, the Process_Unit procedure will be called once for each file to process. The Name and Description formals allow the automatic generation of a “help” message on the command-line (see later). Implementing the Process_Unit procedure is as easy as running minor adjustments on Martyn’s original code:

procedure Report (Node : LAL.Ada_Node'Class);
--  Report the use of an access type at Filename/Line_Number on the standard
--  output.

------------
-- Report --
------------

procedure Report (Node : LAL.Ada_Node'Class) is
Filename : constant String := Node.Unit.Get_Filename;
Line     : constant Slocs.Line_Number :=
Node.Sloc_Range.Start_Line;
begin
Put_Line (Filename & ":"
end Report;

------------------
-- Process_Unit --
------------------

procedure Process_Unit
(Job_Ctx : Helpers.App_Job_Context; Unit : LAL.Analysis_Unit)
is
pragma Unreferenced (Job_Ctx);

function Process_Node (Node : Ada_Node'Class) return Visit_Status;
--  Callback for LAL.Traverse

------------------
-- Process_Node --
------------------

function Process_Node (Node : Ada_Node'Class) return Visit_Status is
begin
case Node.Kind is
if Node.As_Base_Type_Decl.P_Is_Access_Type then
Report (Node);
end if;

if Node.As_Object_Decl.F_Type_Expr
.P_Designated_Type_Decl.P_Is_Access_Type
then
Report (Node);
end if;

when others =>

--  Nothing interesting was found in this Node so continue
--  processing it for other violations.

return Into;
end case;

--  A violation was detected, skip over any further processing of this
--  node.

return Over;
end Process_Node;

begin
if not Unit.Has_Diagnostics then
Unit.Root.Traverse (Process_Node'Access);
end if;
end Process_Unit;

We’re nearly done! All that’s left to do is to make our program only call the Run procedure:

begin
App.Run;
end ptrfinder1;

That’s it. Build and run this program:

$./ptrfinder1 No source file to process$ ./ptrfinder1 basic_pointers.adb
/tmp/access-type-detector/test/basic_pointers.adb:5

So far, so good.

$./ptrfinder1 --help usage: ptrfinder1 [--help|-h] [--charset|-C CHARSET] [--project|-P PROJECT] [--scenario-variable|-X SCENARIO-VARIABLE[SCENARIO-VARIABLE...]] [--target TARGET] [--RTS RTS] [--config CONFIG] [--auto-dir|-A AUTO-DIR[AUTO-DIR...]] [--no-traceback] [--symbolic-traceback] files [files ...] Look for the use of access types in the input sources positional arguments: files Files to analyze optional arguments: --help, -h Show this help message […] Wow, that’s a lot! As you can see, App takes care of parsing command-line arguments and provides a lot of built-in options. Most of them are for the various ways to communicate to the application the set of source files to process: • "ptrfinder1 source1.adb source2.adb …" will process all source files on the command-line, assuming that all source files belong to the current directory; • "ptrfinder1 -P my_project.gpr [-XKEY=VALUE] [--target=…] [--RTS=…] [--config=…]" will process all source files that belong to the my_project.gpr project file. If additional source files appear on the command-line, ptrfinder1 will process only them, but my_project.gpr will still be used to find the other source files. • "ptrfinder1 --auto-dir=src1 --auto-dir=src2" will process all Ada source files that can be found in the src1 and src2 directories. Likewise, additional source files on the command-line will restrict processing to them. These three use cases should cover most needs, the most reliable one being the project file way: calling gprbuild on the project file (with the same arguments) is a cheap way to check using the compiler that the set of sources passed to the application/Libadalang is complete, consistent and valid Ada. As it is a common gotcha, let’s take a moment to note that even though your application may process only one source file, Libadalang may need to get access to other source files. For instance, computing the type of a variable in source1.adb may require to read pkg.ads, which defines the type of this variable. This is why passing a project file or --auto-dir options is useful even when you pass the list of source files to process explicitly on the command-line. Martyn’s second program (ptrfinder2) doesn’t use Libadalang, so rewriting it to use App isn’t very interesting. Instead, let’s extend the previous program to run the text verification on the fly. We are going to add a command-line option to our application to optionally do the verification. Right after the App instantiation, add: package Do_Verify is new GNATCOLL.Opt_Parse.Parse_Flag (App.Args.Parser, Long => "--verify", Help => "Verify detected ""access"" occurences"); App’s command-line parser (App.Args.Parser) uses the GNATCOLL.Opt_Parse library, so adding support for new command-line options is very easy. Here, we add a flag, i.e. a switch with no argument: it’s either present or absent. Just doing this already extends the automatic help message: $ ./ptrfinder1 --help
usage: ptrfinder1 […]
files [files ...] [--verify]

Look for the use of access types in the input sources

positional arguments:
files                 Files to analyze

optional arguments:
[…]
--verify,             Verify detected "access" occurences

Now we can modify the Report procedure to handle this option:

function Verify
(Filename : String; Line : Slocs.Line_Number) return Boolean;
--  Return whether Filename can be read and that its Line'th line contains
--  the " access " substring.

--  Report the use of an access type at Filename/Line_Number on the standard
--  output. If --verify is enabled, check that the first source line
--  corresponding to Node contains the " access " substring.

------------
-- Verify --
------------

function Verify
(Filename : String; Line : Slocs.Line_Number) return Boolean
is
--  Here, we could directly look for an "access" token in the list of
--  tokens corresponding to Line in this unit. However, in the spirit of

Found : Boolean := False;
--  Whether we have found the substring on the expected line

File : File_Type;
begin
Open (File, In_File, Filename);
for I in 1 .. Line loop
declare
use type Slocs.Line_Number;

Line_Content : constant String := Get_Line (File);
begin
if I = Line
and then Ada.Strings.Fixed.Index (Line_Content, " access ") > 0
then
Found := True;
end if;
end;
end loop;
Close (File);
return Found;
exception
when Use_Error | Name_Error | Device_Error =>
Close (File);
return Found;
end Verify;

------------
-- Report --
------------

procedure Report (Node : LAL.Ada_Node'Class) is
Filename   : constant String := Node.Unit.Get_Filename;
Line       : constant Slocs.Line_Number :=
Node.Sloc_Range.Start_Line;
Line_Image : constant String :=
begin
if Do_Verify.Get then
if Verify (Filename, Line) then
Put_Line ("Access Type Verified on line #"
& Line_Image & " of " & Filename);
else
Put_Line ("Suspected Access Type *NOT* Verified on line #"
& Line_Image & " of " & Filename);
end if;

else
Put_Line (Filename & ":" & Line_Image);
end if;
end Report;

And voilà! Let’s check how it works:

./ptrfinder1 basic_pointers.adb --verify Access Type Verified on line #3 of /tmp/access-type-detector/test/basic_pointers.adb Access Type Verified on line #5 of /tmp/access-type-detector/test/basic_pointers.adb When writing Libadalang-based tools, don’t waste time with trivialities such as command-line parsing: use Libadalang.Helpers.App and go directly to the interesting parts! You can find the compilable project for this post on my GitHub fork. Just make sure you get Libadalang 20.1 or the next Continuous Release (coming in February 2020). As usual, please send us suggestions and bug reports on GNATtracker (if you are an AdaCore customer) or on Libadalang’s GitHub project. ]]> Using GNAT-LLVM to target Ada to WebAssembly https://blog.adacore.com/use-of-gnat-llvm-to-translate-ada-applications-to-webassembly Tue, 04 Feb 2020 07:37:00 -0500 Vadim Godunko https://blog.adacore.com/use-of-gnat-llvm-to-translate-ada-applications-to-webassembly The GNAT-LLVM project provides an opportunity to port Ada to new platforms, one of which is WebAssembly. We conducted an experiment to evaluate the porting of Ada and the development of bindings to use Web API provided by the browser directly from Ada applications. Subtotals As a result of the experiment, the standard language library and runtime library were partially ported. Together with a binding for the Web API, this allowed us to write a simple example showing the possibility of using Ada for developing applications compiled into WebAssembly and executed inside the browser. At the same time, there are some limitations both of WebAssembly and of the current GNAT-LLVM implementation: • the inability to use tasks and protected types • support for exceptions limited to local propagation and the last chance handler • the inability to use nested subprograms Example Here is small example of an Ada program that shows/hides the text when pressing the button by manipulating attributes of document nodes. with Web.DOM.Event_Listeners; with Web.DOM.Events; with Web.HTML.Buttons; with Web.HTML.Elements; with Web.Strings; with Web.Window; package body Demo is function "+" (Item : Wide_Wide_String) return Web.Strings.Web_String renames Web.Strings.To_Web_String; type Listener is limited new Web.DOM.Event_Listeners.Event_Listener with null record; overriding procedure Handle_Event (Self : in out Listener; Event : in out Web.DOM.Events.Event'Class); L : aliased Listener; ------------------ -- Handle_Event -- ------------------ overriding procedure Handle_Event (Self : in out Listener; Event : in out Web.DOM.Events.Event'Class) is X : Web.HTML.Elements.HTML_Element := Web.Window.Document.Get_Element_By_Id (+"toggle_label"); begin X.Set_Hidden (not X.Get_Hidden); end Handle_Event; --------------------- -- Initialize_Demo -- --------------------- procedure Initialize_Demo is B : Web.HTML.Buttons.HTML_Button := Web.Window.Document.Get_Element_By_Id (+"toggle_button").As_HTML_Button; begin B.Add_Event_Listener (+"click", L'Access); B.Set_Disabled (False); end Initialize_Demo; begin Initialize_Demo; end Demo; As you can see, it uses elaboration, tagged and interface types, and callbacks. Live demo Setup & Build To compile the examples you need to setup GNAT-LLVM & GNAT WASM RTL following instructions in README.md file. To compile specific example use gprbuild to build application and open index.html in the browser to run it. Next steps The source code is published in a repository on GitHub and we invite everyone to participate in the project. Attachments ]]> AdaCore at FOSDEM 2020 https://blog.adacore.com/adacore-at-fosdem-2020 Thu, 30 Jan 2020 11:03:04 -0500 Fabien Chouteau https://blog.adacore.com/adacore-at-fosdem-2020 Like last year and the year before, AdaCore will participate to the celebration of Open Source software at FOSDEM. It is always a key event for the Ada/SPARK community and we are looking forward to meet Ada enthusiasts. You can check the program of the Ada/SPARK devroom here. AdaCore engineers will give two talks in the Ada devroom: We have a talk in the Hardware Enablement devroom: And there is a related talk in the Security devroom on the use of SPARK for security: Hope to see you at FOSDEM this week-end! ]]> Ada on a Feather https://blog.adacore.com/ada-on-a-feather Thu, 23 Jan 2020 09:35:59 -0500 Fabien Chouteau https://blog.adacore.com/ada-on-a-feather In the last couple of years, the maker community switched from AVR based micro-controllers (popularized by Arduino) to the ARM Cortex-M architecture. AdaFruit was at the forefront of this migration, with boards like the Circuit Playground Express or some of the Feathers. AdaFruit chose to adopt the Atmel (now Microchip) SAMD micro-controller family. Unfortunately for us it is not in the list of platforms with the most Ada support so far (stay tuned, this might change soon ;)). So I was quite happy to see AdaFruit release their first Feather format board including a micro-controller with plenty of Ada support, the STM32F4. I bought a board right away and implemented some support code for it. The support for the Feather STM32F405 is now available in the Ada Drivers Library, along with two examples. The first just blinks the on-board LED and the second displays Make With Ada on a CharlieWing expansion board. Setup To compile the examples, you need to download and install a couple of things: the GNAT arm-elf package from adacore.com/download (I also recommend the native package to get the IDE GNAT Studio) and the Ada Driver Library code from GitHub (AdaCore/Ada_Drivers_Library). You then have to run the script scripts/install_dependencies.py to install the run-time BSPs. Build To build the example, open one of the the project files examples/feather_stm32f405/blinky/blinky.gpr or examples/feather_stm32f405/charlie_wing/charlie_wing.gpr with GNATstudio (Aka GPS), and click on the “build all” icon. Program the board To program the example on the board, I recommend using the Black Magic Probe debugger (also available from AdaFruit). This neat little device provides a GDB remote server interface to the STM32F4, allowing you not only to program the micro-controller but also to debug it. An alternative is to use the DFU mode of the STM32F4. Happy hacking :) ]]> Witnessing the Emergence of a New Ada Era https://blog.adacore.com/the-emergence-of-a-new-ada-era Tue, 21 Jan 2020 09:17:00 -0500 Quentin Ochem https://blog.adacore.com/the-emergence-of-a-new-ada-era For nearly four decades the Ada language (in all versions of the standard) has been helping developers meet the most stringent reliability, safety and security requirements in the embedded market. As such, Ada has become an entrenched player in its historic A&D niche, where its technical advantages are recognized and well understood. Ada has also seen usage in other domains (such as medical and transportation) but its penetration has progressed at a somewhat slower pace. In these other markets Ada stands in particular contrast with the C language, which, although suffering from extremely well known and documented flaws, remains a strong and seldom questioned default choice. Or at least, when it’s not the choice, C is still the starting point (a gateway drug?) for alternatives such as C++ or Java, which in the end still lack the software engineering benefits that Ada embodies.. Throughout AdaCore’s twenty-five year history, we’ve seen underground activities of software engineers willing to question the status quo and embark on new technological grounds. But driving such a change is a tough sell. While the merits of the language are usually relatively easy to establish, overcoming the surrounding inertia often feels like an insurmountable obstacle. Other engineers have to be willing to change old habits. Management has to be willing to invest in new technology. All have to agree on the need for safer, more secure and more reliable software. Even if we’ve been able to report some successes over the years, we were falling short of the critical mass. Or so it seemed. The tide has turned. 2018 and 2019 have been exceptional vintages in terms of Ada and SPARK adoption, all the signs are showing that 2020 will be at least as exciting. What’s more - the new adopters are coming from industries that were never part of the initial Ada and SPARK user base. What used to be inertia is now momentum. Let’s take a look at the information that can be gathered from the web over the past two years to demonstrate the new dynamic of Ada and SPARK usage. The Established User Base Before talking about new adopters, it’s important to step back and re-establish the basis of the Ada and SPARK usage, which is the root of its viability over the very long term. Ada and SPARK are used by a very large user base in the defense and avionics domains. A glance at AdaCore customer list - a subset of the actual user base - will give a good idea of the breadth of technology usage. A lot of the projects here have lifetimes over decades, some started in the early days of Ada in the mid 80’s carried all the way to the present, some have already planned lifetimes spanning over the next two decades. Projects range from massive air traffic management systems running on vast arrays of servers to embedded controllers running on aircraft engines, sensors, or satellite flight control systems with extremely stringent resource constraints. Some applications are still maintained today on hardware dating as far back as Motorola 68K or Intel i386 series, while others are deployed on the latest ARM Cortex or RISC-V cores. Most have some level of reliability constraints, up to the highest levels of the avionics DO-178B/C standard. Due to the nature of the domain, it is difficult to communicate specifically about these projects, and we only have scarce news. One measure of the increasing interest in Ada and SPARK can be inferred from defense-driven research projects which contain references to these language technologies. The most notable example is the recent UK-funded HICLASS project, focused on security, which involves a large portion of the UK defense industry. Some press releases are also available, in particular in the space domain (European Space Agency, AVIO and MDA). These data samples are representative of a very active and vibrant community which is committed to Ada and SPARK for decades to come - effectively guaranteeing their industrial future as far as we can reasonably guess. The Emerging Adopters The so-called “established user base” has fueled the Ada and SPARK community up until roughly the mid 2010s. At that point of time, a new trend started to emerge, from users and use cases that we’had never seen before. While each case is a story in its own right, some common patterns have emerged. The starting point is almost always either the increase of safety or security requirements, or a wish to reduce the costs of development of an application with some kind of high reliability needs. This is connected to the acknowledgement that the programming language in use - almost exclusively C or C++ - may not be the optimal language to reach these goals. This is well documented in the industry; C and C++ flaws have been the subject of countless papers, and the source of catastrophic vulnerability exploits and tools to work around issues. The technical merits of Ada and its ability to prevent many of these issues is also well documented - we even have access to some measurements of the effects. The most recent one is an independent study developed by VDC, which measured up to 38% cost-savings on Ada vs C in the context of high-integrity markets that have adopted Ada for a long time. We’re talking a lot about Ada here, but in fact new adopters are typically driven by a mix of SPARK and Ada. The promise that SPARK offers is automatic verification of software properties such as absence of buffer overflow, together with stringent mitigation of others - and this by design, early in the development process. This means that developers are able to self-check their code - not only is the code more reliable, it is also more reliable straight out as you write it, avoiding many mistakes that could otherwise pass through testing, integration or deployment phases. Some of the SPARK adopters motivated by these benefits come from academia. Over the past 2 years, over 40 universities have joined the GNAT Academic Program (“GAP”), with a mix of teaching and research activities, including for example FH Campus Wien train project, CubeSat and UPMSat-2. Many adopters can also be found in industry. Some of the following references highlight teams at the research phase, some others represent projects already deployed. They all however contribute to this solid wave of new Ada and SPARK adopters. The publications referenced in the following paragraphs have been published between 2018 and 2019. One obvious application for Ada and SPARK, where human lives are at risk, is the medical device domain. So it comes without surprise that this area is amongst those adopting the technology. Two interesting cases come to mind. The first one in RealHeart, a Scandinavian manufacturer that is developing an artificial heart with on-board software written in Ada and SPARK, who issued a press release and later made an in-depth presentation at SPARK & Frama-C days. The second reference comes from a large medical device corporation, Hillrom, who published a paper explaining the rationale for the selection of SPARK and Ada for development of ECG algorithms. Another domain is everything that relates to security. The French security agency ANSSI studied various languages to implement a secure USB key and selected SPARK as the best choice. They published a research paper, presentation and source code. Another interesting new application has been implemented by a German company Componolit developing proven communication protocols Of course, established markets are also at the party. The University of Colorado’s Laboratory for Atmospheric and Space Physics has recently adopted Ada to develop an application for the International Space Station. In the defense domain, the Air Force Research Labs is studying re-writing a drone framework from C++ to SPARK and doing functional proofs, with a public research paper and source code available. While all of these domains provide interesting adopter stories, the one single domain that has demonstrated the most interest in the recent past is undoubtedly automotive. This is probably coming from the increasing complexity of electronics systems in cars, with applications such as Advanced Driver Assistance Systems (ADAS) and autonomous vehicles. References in this domain ranges from tier 1 suppliers such as Denso or JTEKT as well as OEMs and autonomous vehicle companies like Volvo’s subsidiary Zenuity And there’s NVIDIA. In January of this year, we published with NVIDIA a press release and a blog post, followed-up this November by a presentation at our annual Tech Days conference, and an on-line webex (also see the slides for the webex). In many respects, this is a unique tipping point in the history of Ada adoption in terms of impact in a non-A&D domain, touching considerations ranging from security to automotive safety, all under the tight constraints of firmware development. The webex in particular provides a unique dive into the reasons behind the adoption of SPARK and Ada by a company that didn’t have any particular ties to it initially. It also gives key insights on the challenges and costs of such an adoption, together with the benefits already observed. In many respects, this is almost an adoption guide to the technology from a business standpoint. Wrapping Up Keep in mind that the above references are only those that are publicly available, which we know about. There are many more projects under the hood, and even more that we’re not even aware of. Everything considered, this is a very exciting time for the Ada and SPARK languages. Stay tuned, we have an array of new stories coming up for the months and years to come! ]]> AdaCore for HICLASS - Enabling the Development of Complex and Secure Aerospace Systems https://blog.adacore.com/adacore-for-hiclass Wed, 11 Dec 2019 07:27:00 -0500 Paul Butcher https://blog.adacore.com/adacore-for-hiclass What's changed? In 2019 AdaCore created a UK business unit and embarked on a new and collaborative venture researching and developing advanced UK aerospace systems. This blog introduces the reader to ‘HICLASS’, describes our involvement and explains how participation in this project is aligned with AdaCore’s core values. Introducing HICLASS The “High-Integrity, Complex, Large, Software and Electronic Systems” (HICLASS) project was created to enable the delivery of the most complex, software-intensive, safe and cyber-secure systems in the world. HICLASS is a strategic initiative to drive new technologies and best-practice throughout the UK aerospace supply chain, enabling the UK to affordably develop systems for the growing aircraft and avionics market expected over the coming decades. HICLASS includes key prime contractors, system suppliers, software tool vendors and Universities working together to meet the challenges of growing system complexity and size. HICLASS will allow the development of new, complex, intelligent and internet-connected electronic products that are safe and secure from cyber-attack and can be affordably certified. The HICLASS project is supported by the Aerospace Technology Institute (ATI) Programme, a joint Government and industry investment to maintain and grow the UK’s competitive position in civil aerospace design and manufacture. The programme, delivered through a partnership between the ATI, Department for Business, Energy & Industrial Strategy (BEIS) and Innovate UK, addresses technology, capability and supply chain challenges. The £32m investment program, led by Rolls-Royce Control Systems, focuses on the UK civil aerospace sector but also has direct engagement with the Defence, Science and Technology Laboratory (DSTL). The collaborative group, comprised of 16 funded partners and 2 unfunded partners, is made up of the following system developers, tool suppliers and academic institutions: AdaCore, Altran, BAE Systems, Callen-Lenz, Cobham, Cocotec, D-Risq, GE Aviation, General Dynamics UK, Leonardo, MBDA, University of Oxford, Rapita Systems, Rolls-Royce, University of Southampton, Thales, Ultra Electronics and University of York. As well as researching and developing advanced aerospace capabilities, the group aims to pool niche skills and build a highly collaborative community based around the enhanced understanding of shared problems. The project is split into 4 main work packages with 2 technology work packages focusing on integrated model based engineering, cyber-secure architectures and mechanisms, high integrity connectivity, networks and data distribution, advanced hardware platforms and smart sensors and advanced software verification capabilities. In addition, a work package will ensure domain exploitation and drive a cross-industry cyber-security regulatory approach for avionics. A final work package will see the development of integrated HICLASS technology demonstrators. Introducing ASSET HICLASS also aims to build, promote and manage the Aerospace Software Systems Engineering and Technology (ASSET) partnership. This community is open to all organisations undertaking technical work in aerospace software and systems engineering in the UK and operates in a manner designed to promote sharing, openness and accessibility. Unlike HICLASS, ASSET publications are made under a Creative Commons Licence, and the group operates without any non-disclosure or collaboration agreements. AdaCore's R&D Work in the UK Within HICLASS, AdaCore is working with partners across multiple work packages and is also leading a work package titled “SPARK for HICLASS”. This work package will develop and extend multiple SPARK-related technologies in order to satisfy industrial partner’s HICLASS requirements regarding safety and cyber-security. SPARK is a globally recognised safety and security profile of Ada and a software programming language defined by IEC/ISO 8652:2012. Born out of a UK MOD sponsored research project, the first version of SPARK, based on Ada 83, was initially produced at the University of Southampton. Since then the technology has been progressively extended and refined and the latest version SPARK 2014, based on Ada 2012, is now maintained and developed by AdaCore and Altran in partnership. Due to its rich pedigree, earnt at the forefront of high integrity software assurance, SPARK plays a big part in AdaCore’s safe and secure software development tool offerings. Through focused and collaborative research and development, AdaCore will guide the evolution of multiple SPARK-related technologies towards a level where they are suitable for building demonstrable, safe and secure cyber-physical systems that meet the software implementation and verification requirements of HICLASS developed by UK Plc. New extensions to the SPARK language, specific to HICLASS systems, will be developed, these will include the verification of cyber-safe systems and auto generated code. There is also a planned maturing of SPARK reusable code modules where AdaCore will be driven by the needs of our partners in providing high assurance reusable SPARK libraries resulting in the reduction of development time and reduced verification costs. QGen, a qualifiable and tuneable code generation and model verification tool suite for a safe subset of Simulink® and Stateflow® models, is as a game changer in Model Based Software Engineering (MBSE). For HICLASS, AdaCore will place an emphasis on the fusion of SPARK verification capabilities and HICLASS-related emerging MBSE tools, allowing code level verification to be achieved at the model level. The generation of SPARK code, from our QGen tool, as well as various HICLASS partner’s MBSE technologies, will be researched and developed. Collaborative case studies will be performed to assess and measure success. Collaboration is a key critical success factor in meeting this objective; multiple HICLASS partners are developing MBSE tools and SPARK evolution will be achieved in close partnership with them. The second, and complementary, objective of this work package is to research and develop cyber-secure counter measures and HICLASS verification strategies, namely in the form of compiler hardening and the development of a ‘fuzzing’ capability for Ada/SPARK. HICLASS case studies, produced within proceeding work packages, will be observed to ensure our SPARK work package is aligned with HICLASS specific standards, guidelines and recommendations and to ensure the relevancy of the work package deliverables. The third objective is for AdaCore, in collaboration with our HICLASS partners, to evaluate QGen, and associated formal approaches, for existing UK aerospace control systems and to make comparisons with existing Simulink code generation processes. In addition, AdaCore will promote processor emulation technology through a collaborative HICLASS case study. The final objective is to demonstrate the work package technology through the creation of a software stack capable of executing SPARK software on a range of (physical and emulated) target processors suitable for use in HICLASS. The ability to execute code generated from MBSE environments will also be demonstrated. Committing Investment into the UK AdaCore has a long history of working with partners within the UK aerospace industry on safety-related, security-related and mission-critical software development projects. Participation in the HICLASS research and development group complemented AdaCore’s commitment to invest within the UK. This four-year research project is also an excellent fit with AdaCore’s core values and its existing and future capabilities. In addition, the creation of a new UK business unit, ‘AdaCore Ltd’, created to rapidly grow into our UK Centre of Excellence, ensures that our existing and future UK aerospace customers will continue to receive the high level of technical expertise and quality products associated with AdaCore. History has shown that the UK aerospace industry isn’t afraid to be ambitious and has the technological capability to stay at the forefront of this rapidly growing sector. With HICLASS, the sky really is the limit, and AdaCore welcomes the opportunity to be a part of the journey and further extend our partnerships within this technologically advanced and continually growing market. Further information about the ATI, BEIS and IUK... Aerospace Technology Institute (ATI) The Aerospace Technology Institute (ATI) promotes transformative technology in air transport and supports the funding of world-class research and development through the multi-billion pound joint government-industry programme. The ATI stimulates industry-led R&D projects to secure jobs, maintain skills and deliver economic benefits across the UK. Setting a technology strategy that builds on the UK’s strengths and responds to the challenges faced by the UK civil aerospace sector; ATI provides a roadmap of the innovation necessary to keep the UK competitive in the global aerospace market, and complements the broader strategy for the sector created by the Aerospace Growth Partnership (AGP). The ATI provides strategic oversight of the R&T pipeline and portfolio. It delivers the strategic assessment of project proposals and provides funding recommendations to BEIS. Department for Business, Energy and Industrial Strategy (BEIS) Department for Business, Energy and industrial Strategy (BEIS) is the government department accountable for the ATI Programme. As the budget holder for the programme, BEIS, is accountable for the final decision regarding projects to progress and fund with Government resources, as well as performing Value for Money (VfM) assessment on all project proposals, one of the 3 ATI Programme assessment streams. Innovate UK (IUK) Innovate UK is the funding agency for the ATI Programme. It delivers the competitions process including independent assessment of project proposals, and provides funding recommendations to BEIS. Following funding award, Innovate UK manages the programme, from contracting projects, through to completion. Innovate UK is part of UK Research and Innovation (UKRI), a non-departmental public body funded by a grant-in-aid from the UK government. Innovate UK drives productivity and economic growth by supporting businesses to develop and realise the potential of new ideas, including those from the UK’s world-class research base. UKRI is the national funding agency investing in science and research in the UK. Operating across the whole of the UK with a combined budget of more than £6 billion, UKRI brings together the 7 Research Councils, Innovate UK and Research England. ]]> An Expedition into Libadalang https://blog.adacore.com/an-expedition-into-libadalang Thu, 07 Nov 2019 08:04:00 -0500 Martyn Pike https://blog.adacore.com/an-expedition-into-libadalang I’ve been telling Ada developers for a while now that Libadalang will open up the possibility of more-easily writing Ada source code analysis tools. (You can read more about Libadalang here and here and can also access the project on Github.) Along these lines, I recently had a discussion with a customer about whether there were any tools for detecting uses of access types in their code which got me thinking about possible ways to detect the use of Access Types in a set of Ada source code files. GNATcheck doesn't currently have a rule that prohibits the use of access types. Also, SPARK 2014 recently added support for Access Types, whereas previously they were banned. So earlier versions of GNATprove could detect them quite effectively, the latest and future versions may not. I decided to architect a solution to this problem and determined there were several implementation options open to me: 1. Use ‘grep’ on a set of Ada sources to find instances of the "access" Ada keyword 2. Use gnat2xml and then use ‘grep’ on the resulting output to search for certain tags 3. Use gnat2xml and then write an XML-aware search utility to search for certain tags 4. Use Libadalang to write my own Ada static analysis program Option 1 and 2 just feel too easy and would defeat the purpose of this blog post. Option 3 is perhaps a good topic for another post related to using XML/Ada, however I decided to put my money where my mouth is and go with Option 4! While I wrote this program in Ada, I could have written it in Python. So here is the program: with Ada.Text_IO; use Ada.Text_IO; with Libadalang.Analysis; use Libadalang.Analysis; with Libadalang.Common; use Libadalang.Common; with Ada.Strings.Fixed; with Ada.Strings; procedure ptrfinder1 is LAL_CTX : constant Analysis_Context := Create_Context; begin Read_Standard_Input: while not End_Of_File(Standard_Input) loop Process_Ada_Unit: declare Filename : constant String := Get_Line; Unit : constant Analysis_Unit := LAL_CTX.Get_From_File(Filename); function Process_Node(Node : Ada_Node'Class) return Visit_Status is begin if Node.Kind in Ada_Access_Def | Ada_Access_To_Subp_Def_Range | Ada_Base_Type_Access_Def | Ada_Anonymous_Type_Access_Def_Range | Ada_Type_Access_Def_Range then Put_Line( Ada.Strings.Fixed.Trim( Source => Filename & ":" & Node.Sloc_Range.Start_Line'Img, Side => Ada.Strings.Left ) ); end if; return Into; end Process_Node; begin if not Unit.Has_Diagnostics then Unit.Root.Traverse(Process_Node'Access); end if; end Process_Ada_Unit; end loop Read_Standard_Input; end ptrfinder1; I designed the program to read a series of fully qualified absolute filenames from standard input and process each of them in turn. This approach made the program much easier to write and test and, as you'll see, allowed the program to be integrated effectively with other tools. Let's deconstruct the code a little.... For each provided filename, the program creates a Libadalang Analysis_Unit for that filename. Read_Standard_Input: while not End_Of_File(Standard_Input) loop Process_Ada_Unit: declare Filename : constant String := Get_Line; Unit : constant Analysis_Unit := LAL_CTX.Get_From_File(Filename); As long as it has no issues, the Ada unit is traversed and the Process_Node subprogram is executed for each detected node. if not Unit.Has_Diagnostics then Unit.Root.Traverse(Process_Node'Access); end if; The Process_Node subprogram checks the Kind field of the detected Ada_Node'Class parameter to see if it is any of the access type related nodes. If so, the program outputs the fully qualified filename, a ":" delimiter, and the line number of the detected node. function Process_Node(Node : Ada_Node'Class) return Visit_Status is begin if Node.Kind in Ada_Access_Def | Ada_Access_To_Subp_Def_Range | Ada_Base_Type_Access_Def | Ada_Anonymous_Type_Access_Def_Range | Ada_Type_Access_Def_Range then Put_Line( Ada.Strings.Fixed.Trim( Source => Filename & ":" & Node.Sloc_Range.Start_Line'Img, Side => Ada.Strings.Left ) ); end if; return Into; end Process_Node; At the end of the Process_Node subprogram, the returned value allows the traversal to continue. To make the program a more useful tool within a development environment based on GNAT Pro, I integrated it with the piped output of the 'gprls' program. gprls is a tool that outputs information about compiled sources. It gives the relationship between objects, unit names, and source files. It can also be used to check source dependencies as well as various other characteristics. My program can then be invoked as part of a more complex command line:  gprls -s -P test.gpr | ./ptrfinder1

Given the following content of test.gpr:

project Test is

for Source_Dirs use (".");
for Object_Dir use "obj";

end Test;

Plus an Ada source code file called inc_ptr1.adb (in the same directory as test.gpr) containing the following:

procedure Inc_Ptr1 is

type Ptr is access all Integer;

begin

null;

end Inc_Ptr1;

The resulting output from the integration of gprls and my program is:

/home/pike/Workspace/access-detector/test/inc_ptr1.adb:3

This output correctly identified the access type usage on line 3 of inc_ptr1.adb.

But how do I know that my program or indeed Libadalang has functioned correctly?

I decided to stick in principle to the UNIX philosophy of "Do One Thing and Do it Well" and write a second program to verify the output of my first program using a simple algorithm.

This second program is given a filename and line number and verifies that the keyword "access" appears on the specified line number.

Of course,  I could also have embedded this verification into the first program,  but to illustrate a point about diversity I chose not to.

with Ada.Text_IO;       use Ada.Text_IO;

procedure ptrfinder2 is
begin

while not End_Of_File(Standard_Input)
loop

Process_Standard_Input:
declare

Std_Input : constant String := Get_Line;
Delimeter_Position : constant Natural := Index(Std_Input,":");
Line_Number_As_String : constant String :=  Std_Input(Delimeter_Position+1..Std_Input'Last);
Line_Number : constant Integer := Integer'Value(Line_Number_As_String);
Filename : constant String := Std_Input(Std_Input'First..Delimeter_Position-1);
The_File : File_Type;
Verified : Boolean := False;

begin

if Ada.Directories.Exists(Filename) and then Line_Number > 1
then

Open(File => The_File, Mode => In_File, Name => Filename);

Locate_Line:
for I in 1..Line_Number loop

Verified := Index(Get_Line(The_File)," access ") > 0;

exit when Verified or else End_Of_File(The_File);

end loop Locate_Line;

Close(File => The_File);

end if;

if Verified then
Put_Line("Access Type Verified on line #" & Line_Number_As_String & " of " & Filename);
else
Put_Line("Suspected Access Type *NOT* Verified on line #" & Line_Number_As_String & " of " & Filename);
end if;

end Process_Standard_Input;

end ptrfinder2;

I can then string the first and second program together:

$gprls -s -P test.gpr | ./ptrfinder1 | ./ptrfinder2 This produces the output: Access Type Verified on line #3 of /home/pike/Workspace/access-detector/test/inc_ptr1.adb It goes without saying that a set of Ada sources with no Access Type usage will result in no output from either the first or second program. This expedition into Libadalang has reminded me how extremely effective Ada can be at writing software development tools. The two programs described in this blog post were built and tested on 64-bit Ubuntu 19.10 using GNAT Pro and Libadalang. They are also known to build successfully with the 64-bit Linux version of GNAT Community 2019. The source code can be downloaded and built from GitHub. ]]> RecordFlux: From Message Specifications to SPARK Code https://blog.adacore.com/recordflux-from-message-specifications-to-spark-code Thu, 17 Oct 2019 09:08:23 -0400 Alexander Senier https://blog.adacore.com/recordflux-from-message-specifications-to-spark-code Software constantly needs to interact with its environment. It may read data from sensors, receive requests from other software components or control hardware devices based on the calculations performed. While this interaction is what makes software useful in the first place, processing messages from untrusted sources inevitably creates an attack vector an adversary may use to exploit software vulnerabilities. The infamous Heartbleed is only one example where a security critical software was attacked by specially crafted messages. Implementing those interfaces to the outside world in SPARK and proving the absence of runtime errors is a way to prevent such attacks. Unfortunately, manually implementing and proving message parsers is a tedious task which needs to be redone for every new protocol. In this article we'll discuss the challenges that arise when creating provable message parsers and present RecordFlux, a framework which greatly simplifies this task. Specifying Binary Messages Ethernet: A seemingly simple example At a first glance, Ethernet has a simple structure: A 6 bytes destination field, a 6 bytes source field and a 2 bytes type field followed by the payload: We could try to model an Ethernet frame as a simple record in SPARK: package Ethernet is type Byte is mod 2**8; type Address is mod 2**48; type Type_Length is mod 2**16; type Payload is array (1..1500) of Byte; type Ethernet_Frame is record Destination : Address; Source : Address; EtherType : Type_Length; Data : Payload; end record; end Ethernet; When looking closer, we realize that this solution is a bit short-sighted. Firstly, defining the payload as a fixed-size array as above will either waste memory when handling a lot of small (say, 64 bytes) frames or be too short when handling jumbo frames which exceed 1500 bytes. More importantly, the Ethernet header is not as simple as we pretended earlier. Looking at the standard, we realize that the EtherType field actually has a more complicated semantics to allow different frame types to coexist on the same medium. If the value of EtherType is greater or equal to 1536, then it contains an Ethernet II frame. The EtherType is treated as a type field which determines the protocol contained in Data. In that case, the Data field uses up the remainder of the Ethernet frame. If the value of EtherType is less or equal to 1500, then it contains a IEEE 802.3 frame and the EtherType field represents the length of the Data field. Frames with an EtherType value between 1501 and 1535 are considered invalid. To make things even worse, both variants may contain an optional IEEE 802.1Q tag to identify the frames priority and VLAN. The tag is inserted after the source field and itself is comprised of two 16 bit fields, TPID and TCI. It is present if the bytes that would contain the TPID field have a hexadecimal value of 8100. Otherwise these bytes contain the EtherType field. Lastly, the Data field will usually contain higher-level protocols. Which protocol is contained and how the payload is to be interpreted depends on the value of EtherType. With our naive approach above, we have to manually convert Data into the correct structured message format. Without tool support, this conversion will be another source of errors. Formal Specification with RecordFlux Next, we'll specify the Ethernet frame format using the RecordFlux domain specific language and demonstrate how the specification is used to create real-world parsers. RecordFlux, deliberately has a syntax similar to SPARK, but deviates where more expressiveness was required to specify complex message formats. Packages and Types Just like SPARK, entities are grouped into packages. By convention, a package contains one protocol like IPv4, UDP or TLS. A protocol will typically define many types and different message formats. Range types and modular types are identical to those found in SPARK. Just like in SPARK, entities can be qualified using aspects, e.g. to specify the size of a type using the Size aspect: package Ethernet is type Address is mod 2**48; type Type_Length is range 46 .. 2**16 - 1 with Size => 16; type TPID is range 16#8100# .. 16#8100# with Size => 16; type TCI is mod 2**16; end Ethernet; The first difference to SPARK is the message keyword which is similar to records, but has important differences to support the non-linear structure of messages. The equivalent of the naive Ethernet specification in RecordFlux syntax would be: type Simplified_Frame is message Destination : Address; Source : Address; Type_Length : Type_Length; Data : Payload; end message; Graph Structure As argued above, such a simple specification is insufficient to express the complex corner-cases found in Ethernet. Luckily, RecordFlux messages allow for expressing conditional, non-linear field layouts. While SPARK records are linear sequences of fixed-size fields, messages should rather be thought of as a graph of fields where the next field, its start position, its length and constraints imposed by the message format can be specified in terms of other message fields. To ensure that the parser generated by RecordFlux is deadlock-free and able to parse messages sequentially, conditionals must only reference preceding fields. We can extend our simple example above to express the relation of the value of Type_Length and length of the payload field: ... Type_Length : Type_Length then Data with Length => Type_Length * 8 if Type_Length <= 1500, then Data with Length => Message’Last - Type_Length’Last if Type_Length >= 1536; ... For a field, the optional then keyword defines the field to follow in the graph. If that keyword is missing, this defaults to the next field appearing in the message specification as in our Simplified_Frame example above. To have different fields follow under different conditions, an expression can be added using the if keyword. Furthermore, an aspect can be added using the with keyword, which can be used to conditionally alter properties like start or length of a field. If no succeeding field is specified for a condition, as for Type_Length in the range between 1501 and 1535, the message is considered invalid. In the fragment above, we use the value of the field Type_Length as the length of the Data field if its value is less or equal to 1500 (IEEE 802.3 case). If Type_Length is greater or equal to 1536, we calculate the payload length by subtracting the end of the Type_Length field from the end of the message. The 'Last (and also 'First and 'Length) attribute is similar to the respective SPARK attribute, but refer to the bit position (or bit length) of a field within the message. The Message field is special and refers to the whole message being handled. Optional Fields The graph structure described above can also be used to handle optional fields, as for the IEEE 802.1Q tag in Ethernet. Let's have a look at the full Ethernet specification first: package Ethernet is type Address is mod 2**48; type Type_Length is range 46 .. 2**16 - 1 with Size => 16; type TPID is range 16#8100# .. 16#8100# with Size => 16; type TCI is mod 2**16; type Frame is message Destination : Address; Source : Address; Type_Length_TPID : Type_Length then TPID with First => Type_Length_TPID’First if Type_Length_TPID = 16#8100#, then Type_Length with First => Type_Length_TPID’First if Type_Length_TPID /= 16#8100#; TPID : TPID; TCI : TCI; Type_Length : Type_Length then Data with Length => Type_Length * 8 if Type_Length <= 1500, then Data with Length => Message’Last - Type_Length’Last if Type_Length >= 1536; Data : Payload then null if Data’Length / 8 >= 46 and Data’Length / 8 <= 1500; end message; end Ethernet; Most concepts should look familiar by now. The null field used in the then expression of the Data field is just a way to state that the end of the message is expected. This way, we are able to express that the payload length must be between 46 and 1500. As there is only one then branch for payload (pointing to the end of the message), values outside this range will be considered invalid. This is the general pattern to express invariants that have to hold for a message. How can this be used to model optional fields of a message? We just need to cleverly craft the conditions and overlay the following alternatives. The relevant section of the above Ethernet specification is the following: ... Source : Address; Type_Length_TPID : Type_Length then TPID with First => Type_Length_TPID’First if Type_Length_TPID = 16#8100#, then Type_Length with First => Type_Length_TPID’First if Type_Length_TPID /= 16#8100#; TPID : TPID; TCI : TCI; Type_Length : Type_Length ... Remember that the optional IEEE 802.1Q tag consisting of the TPID and TCI fields is present after the Source only if the bytes that would contain the TPID field are equal to a hexadecimal value of 8100. We introduce a field Type_Length_TPID only for the purpose of checking whether this is the case. To avoid any confusion when using the parser later, we will overlay this field with properly named fields. If Type_Length_TPID equals 16#8100# (SPARK-style numerals are supported in RecordFlux), we define the next field to be TPID and set its first bit to the 'First attribute of the Type_Length_TPID field. If Type_Length_TPID does not equal 16#8100#, the next field is Type_Length, skipping TPID and TCI. As stated above, the specification actually is a graph with conditions on its edges. Here is an equivalent graph representation of the full Ethernet specification: Working with RecordFlux RecordFlux comes with the command line tool rflx which parses specification files, transforms them into an internal representation and subsequently generates SPARK packages that can be used to parse the specified messages: To validate specification files, which conventionally have the file ending .rflx, run RecordFlux in check mode. Note, that RecordFlux does not support search paths at the moment and all files need to be passed on the command line: $ rflx check ethernet.rflx
Parsing ethernet.rflx... OK

Code Generation

Code is generated with the generate subcommand which expects a list of RecordFlux specifications and an output directory for the generated SPARK sources. Optionally, a root package can be specified for the generated code using the -p switch:

$rflx generate -p Example ethernet.rflx outdir Parsing ethernet.rflx... OK Generating... OK Created outdir/example-ethernet.ads Created outdir/example-ethernet-generic_frame.ads Created outdir/example-ethernet-generic_frame.adb Created outdir/example-ethernet-frame.ads Created outdir/example-types.ads Created outdir/example-types.adb Created outdir/example-message_sequence.ads Created outdir/example-message_sequence.adb Created outdir/example-scalar_sequence.ads Created outdir/example-scalar_sequence.adb Usage To use the generated code, we need to implement a simple main program. It allocates a buffer to hold the received Ethernet packet and initializes the parser context with the pointer to that buffer. After verifying the message and checking whether it has the correct format, its content can be processed by the application. with Ada.Text_IO; use Ada.Text_IO; with Example.Ethernet.Frame; with Example.Types; procedure Ethernet is package Eth renames Example.Ethernet.Frame; subtype Packet is Example.Types.Bytes (1 .. 1500); Buffer : Example.Types.Bytes_Ptr := new Packet'(Packet'Range => 0); Ctx : Eth.Context := Eth.Create; begin -- Retrieve the packet Buffer (1 .. 56) := (16#ff#, 16#ff#, 16#ff#, 16#ff#, 16#ff#, 16#ff#, 16#04#, 16#d3#, 16#b0#, 16#ab#, 16#f9#, 16#31#, 16#08#, 16#06#, 16#00#, 16#01#, 16#08#, 16#00#, 16#06#, 16#04#, 16#00#, 16#01#, 16#04#, 16#d3#, 16#b0#, 16#9c#, 16#79#, 16#53#, 16#ac#, 16#12#, 16#fe#, 16#b5#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#ac#, 16#12#, 16#64#, 16#6d#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#, 16#00#); Eth.Initialize (Ctx, Buffer); Eth.Verify_Message (Ctx); if Eth.Structural_Valid_Message (Ctx) then Put_Line ("Source: " & Eth.Get_Source (Ctx)'Img); if Eth.Present (Ctx, Eth.F_TCI) and then Eth.Valid (Ctx, Eth.F_TCI) then Put_Line ("TCI: " & Eth.Get_TCI (Ctx)'Img); end if; end if; end Ethernet; The code can be proven directly using gnatprove (the remaining warning is a known issue in GNAT Community 2019 as explained in the Known Issues section of the README): $ gnatprove -P ethernet.gpr
Phase 1 of 2: generation of Global contracts ...
Phase 2 of 2: flow analysis and proof ...
ethernet.adb:22:25: warning: unused assignment to "Buffer"
gnatprove: error during flow analysis and proof

What guarantees can we obtain from this proof? Firstly, the absence of runtime errors is shown for the generated code as well as for the user code. No matter what input is read into the Buffer variable and presented to the parser, the program does not raise an exception at runtime. Furthermore, its control flow cannot be circumvented e.g. by buffer overflows or integer overflows. This is called "silver level" in SPARK parlance. Additionally, we prove key integrity properties ("gold level"), e.g. that optional fields are accessed if and only if all requirements defined in the RecordFlux specification are met.

Case Studies

RecordFlux greatly eases the specification and handling of binary messages. But is it suitable for real-world applications? We conducted a number of case studies to validate that it in fact is.

Packet Sniffer

Packet sniffers are tools often used by administrators to diagnose network problems. They capture and dissect all data received on a network interface to allow for a structured analysis of packet content. Famous open source examples are Wireshark and tcpdump. As packet sniffers need to handle a large amount of complex protocols, they also tend to be complex. There have been errors which allowed attackers to mount remote exploits against packet sniffers. For this reason, e.g. it is discouraged to run Wireshark under a privileged user account.

Obviously a formally verified packet sniffer is desirable to eliminate the risk of an attack when analyzing traffic from untrusted sources. We prototyped a very simple packet sniffer for IP/UDP on Linux for which we proved the absence of runtime errors. The output is similar to other packet sniffers:

$sudo ./obj/sniff_udp_in_ip IP: Version: 4 IHL: 5 DSCP: 0 ECN: 0 TLen: 53 Id: 9059 DF: 1 MF: 0 FOff: 0 TTL: 64 Proto: UDP HCSum: 6483 Src: 127.0.0.1 Dst: 127.0.0.1, UDP: SPort: 58423 DPort: 53 Len: 33 CSum: 65076 Payload: b9 7d 01 00 00 01 00 00 00 00 00 00 04 63 6f 61 70 02 6d 65 00 00 1c 00 01 TLS 1.3 Another area where correct parsers are essential are security protocols. The most important security protocol on the internet is TLS - whenever a browser connects to a remote server using the https protocol, in the background some version of TLS is used. We formalized the message format of the latest TLS version 1.3 according to RFC 8446 and generated a SPARK parser for TLS from the specification. An open question remained, though: Can the generated code handle real-world TLS traffic and is there a performance penalty compared to unverified implementations? While we are working on a verified component-based TLS 1.3 implementation completely done in SPARK, it is not yet available for this experiment. As an alternative, we used an open source TLS 1.3 library by Facebook named Fizz and replaced its TLS parser by our generated code. As the C++ types used by Fizz (e.g. vectors) could not easily be bound to Ada, glue code had to be written manually to translate between the C++ and the SPARK world. We ensured that all untrusted data the library comes in touch with is handled by SPARK code. For the SPARK part, we proved the absence of runtime errors and the invariants stated in the specification. Our constructive approach turned out to be effective for improving the security of existing software. In CVE-2019-3560 an integer overflow was found in the Fizz library. Just by sending a short, specially crafted sequence of messages, an attacker could mount a Denial of Service attack against an application using Fizz by putting it into an infinite loop. While Facebook fixed this bug by using a bigger integer type, the RecordFlux parser prevents this issue by rejecting packets with invalid length fields. Despite the required transformations, the performance overhead was surprisingly low. For the TLS handshake layer - the part that negotiates cryptographic keys when the communication starts - the throughput was 2.7% lower than for the original version. For the TLS record layer, which encrypts and decrypts packets when a connection is active, the throughput was only 1.1% lower: Conclusion and Further Information With RecordFlux, creating SPARK code that handles complex binary data has become a lot easier. With a proper specification the generated code can often be proven automatically. In the future, we will extend RecordFlux to support the generation of binary messages and the modeling of the dynamic behavior of binary protocols. For more information see our research paper and the language reference. If you have comments, found a bug or have suggestions for enhancements, feel free to open an issue on GitHub or write an email. ]]> The Power of Technology Integration and Open Source https://blog.adacore.com/the-power-of-technology-integration Tue, 15 Oct 2019 08:04:00 -0400 Arnaud Charlet https://blog.adacore.com/the-power-of-technology-integration Part of our core expertise at AdaCore is to integrate multiple technologies as smoothly as possible and make it a product. This started at the very beginning of our company by integrating a code generator (GCC) with an Ada front-end (GNAT) which was then followed by integrating a debugger engine (GDB) and led to today's rich GNAT Pro offering. Today we are going much further in this area and I am going to give you a few concrete examples in this post. For example take our advanced static analysis engine CodePeer and let's look at it from two different angles (a bit like bottom-up and top-down if you will): what does it integrate, and what other products integrate it? From the first perspective, CodePeer integrates many different and complex pieces of technology: the GNAT front-end, various "core" GNAT tools (including GNATmetric, GNATcheck, GNATpp), the AdaCore IDEs (GNAT Studio, GNATbench), GNATdashboard, a Jenkins plug-in, GPRbuild, as well as the codepeer engine itself and finally as of version 20, libadalang and light checkers based on libadalang (aka "LAL checkers"). Thanks to this complex integration, CodePeer users can launch various tools automatically and get findings stored in a common database, and get a common user interface to drive it. Indeed, from CodePeer you can get access to: GNAT warnings, GNATcheck messages, LAL checkers, CodePeer "advanced static analysis" messages. And the list will continue growing in the future, but that's for another set of posts! Now from the other perspective, the codepeer engine is also integrated as an additional prover in our SPARK Pro product, to complement the SMT solvers integrated in SPARK, and is also used in our QGen product as the back-end of the QGen Verifier which performs static analysis on Simulink models. Speaking of SPARK, this is also another good example of complex integration of many different technologies: the GNAT front-end, GPRbuild, GNAT Studio and GNATbench, GNATdashboard, Why3, CVC4, Z3, Alt-Ergo, the codepeer engine, and the SPARK engine itself. Many of these components are developed in house at AdaCore, and many other components are developed primarily outside AdaCore (such as GCC, GDB, Why3, CVC4, ...). Such complex integration is only possible because all these components are Open Source and precisely allow this kind of combination. Add on top of that AdaCore's expertise in such integration and productization, and you get our ever growing offering! ]]> Learning SPARK via Conway's Game of Life https://blog.adacore.com/learning-spark Thu, 10 Oct 2019 07:42:21 -0400 Michael Frank https://blog.adacore.com/learning-spark I began programming in Ada in 1990 but, like many users, it took me a while to become a fan of the language. A brief interlude with Pascal in the early 90’s gave me a better appreciation of the strong-typing aspects of Ada, but continued usage helped me learn many of the other selling points. I’ve recently started working at AdaCore as a Mentor, which means part of my job is to help companies transition from other languages to Ada and/or SPARK. At my last job, the development environment was mixed C++ and Ada (on multiple platforms) so, to simplify things, we were coding in Ada95. Some of our customers’ code involved Ada2005 and Ada2012 constructs, so I did learn the new language variants and see the benefits, but I hadn’t actually written anything using contracts or expression functions or the like. We all know the best way to learn is to do, so I needed to come up with something that would require me to increase my knowledge not only of the Ada language, but also to get some more experience with the AdaCore tools (GPS, CodePeer) as well as learning how to prove the correctness of my SPARK code. So I settled on Conway’s Game of Life (Wikipedia). This is an algorithmic model of cell growth on a 2-dimensional grid following four simple rules about cell life and death: • A live cell with fewer than two live neighbors dies (underpopulation) • A live cell with two or three live neighbors stays alive (life) • A live cell with more than three live neighbors dies (overpopulation) • A dead cell with exactly three live neighbors becomes a live cell (birth) This little application gave me everything I needed to experiment with, and even a little more. Obviously, this was going to be written in Ada following SPARK rules, so I could program using the latest language variant. The development environment was going to be AdaCore’s GPS Integrated Development environment, and I could use CodePeer for static analysis on the code I was writing, so I could get some experience with these tools. I then could use the SPARK provers to prove that my code was correct. In addition, as Life has a visual aspect, I was able to do a little work with GtkAda to make a graphical representation of the algorithm. I started with the lower level utilities needed to perform the algorithm – counting the number of live neighbors. Easy, right? Yes, it is easy to code, but not so easy to prove using the SPARK provers. If I wasn’t worried about provability, I could just write a simple doubly-nested loop that would iterate over rows and columns around the target cell, and count how many neighboring cells are alive. The issue is that, with loops in SPARK, we need to “remind” SPARK what has happened in previous loop iterations (typically using “pragma Loop_Invariant”) so it has all the knowledge it needs to prove the current iteration. This issue is made more difficult when we use nested loops. So, I rewrote my “count all neighbors” routine into two routines - “count neighbors in a row” and “total neighbors in all rows”. Each routine has a single loop, which made specifying the loop invariant much easier. And that “rewriting” helped reinforce my understanding of “good coding practices”. Most of us were taught that high-complexity subroutines were problematic because they were difficult to test. So we write code that tries to keep the complexity down. The next level in safety-critical software – provability – is also made more difficult by higher complexity subroutines. In my case, my counting routines needed to check to see if a cell was alive and, if so, increment the count. That’s just a simple “if” statement. But, to make it easier on the prover, I created a new routine that would return 1 if the cell was alive and 0 otherwise. Very simple, and therefore very easy to prove correctness. Now, my counting routine could just iterate over the cells, and sum the calls to the function – once again, easier to prove. Once the low-level functions were created the “real” coding began. And in trying to write pre- and post-conditions for higher level algorithms, I found a great help in the CodePeer tool. This tool performs static analysis on the code base, from simple and quick to exhaustive and slow. The correlation between this tool and the prover tools is that, if I was having a problem writing a valid postcondition, more often than not a CodePeer report would show some warning indicating my preconditions were not complete. For example, every time you add numbers, CodePeer would remind you that a precondition would have to be set to ensure no overflow. Without that warning, everything looks fine, but the prover will not be able to validate the routine. In addition, when using GPS, CodePeer will insert annotations into the Source Code View detailing many of the aspects of each subprogram. These annotations typically include pre- and post-conditions, making it easier to determine which of these aspects could be written into the code to make proving easier. So, a development cycle was created. First, write a subprogram to “do something” (usually this means writing other subprograms to help!) Next, run CodePeer to perform static analysis. In addition to finding run-time errors, the annotations would help determine pre-conditions for the routines. I would study these annotations, and, for each pre- and post-condition, I had to determine if it made sense or not. If the condition did not make sense, I needed to modify my subprogram to get the correct results. If the condition made sense, then I needed to decide if I wanted to encode that condition using aspects, or just ignore – because sometimes the annotations were more detailed than I needed. Once I got past CodePeer analysis, I would try to prove the subprograms I created. If they could be proven, I could go onto the next step in building my application. If not, I needed to modify my code and start the cycle over. This process is basically a one-person Code Review! In a large software project, these steps would be the same on some portion of the code, and only then would the code go into a peer review process – with the benefit that the reviewers are only concerned with design issues, and don’t have to be worried about dealing with run-time errors or verifying that the implementation was correct. So, not only did I achieve my objective of learning aspects of the languages and the tools, I got first-hand evidence of the benefits of writing software “the right way!” To access the source code for this project, please checkout my github repository here: https://github.com/frank-at-adacore/game_of_life ]]> Pointer Based Data-Structures in SPARK https://blog.adacore.com/pointer-based-data-structures-in-spark Tue, 08 Oct 2019 08:19:34 -0400 Claire Dross https://blog.adacore.com/pointer-based-data-structures-in-spark As seen in a previous post, it is possible to use pointers (or access types) in SPARK provided the program abides by a strict memory ownership policy designed to prevent aliasing. In this post, we will demonstrate how to define pointer-based data-structures in SPARK and how to traverse them without breaking the ownership policy. Pointer-based data structures can be defined in SPARK as long as they do not introduce cycles. For example, singly-linked lists and trees are supported whereas doubly-linked lists are not. As an example for this post, I have chosen to use a map encoded as a singly-linked list of pairs of a key and a value. To define a recursive data structure in Ada, you need to use an access type. This is what I did here: I declared an incomplete view of the type Map, and then an access type Map_Acc designating this view. Then, I could give the complete view of Map as a record with a component Next of type Map_Acc: type Map; type Map_Acc is access Map; type Element_Acc is not null access Element; type Map is record Key : Positive; Value : Element_Acc; Next : Map_Acc; end record; Note that I also used an access type for the element that is stored in the map (and not for the key). Indeed, we assume here that the elements stored in the map can be big, so that we may want to modify them in place rather than copying them. We will see later how this can be done. To be able to write specifications on my maps, I need to define some basic properties for them. To describe precisely my maps, I will need to speak about: • whether there is a mapping for a given key in my map, and • the value associated to a key by the map, knowing that there is a mapping for this key. These concepts are encoded as the following functions: function Model_Contains (M : access constant Map; K : Positive) return Boolean; function Model_Value (M : access constant Map; K : Positive) return Element with Pre => Model_Contains (M, K); Model_Contains should return True when there is an occurrence of the key K in the map M, and Model_Value should retrieve the element associated with the first occurrence of K in M. Now, I need to give a definition to my functions. The natural way to express the meaning of properties on linked data-structures is through recursion. This is what I have done here for both Model_Contains and Model_Value: function Model_Contains (M : access constant Map; K : Positive) return Boolean is (M /= null and then (M.Key = K or else Model_Contains (M.Next, K))); -- A key K has a mapping in a map M if either K is the key of the first mapping of M or if -- K has a mapping in M.Next. function Model_Value (M : access constant Map; K : Positive) return Element is (if M.Key = K then M.Value.all else Model_Value (M.Next, K)); -- The value mapped to key K by a map M is either the value of the first mapping of M if -- K is the key this mapping or the value mapped to K in M.Next otherwise. Note that we do not need to care about cases where there is no mapping for K in M in the definition of Model_Value, as we have put as a precondition that it cannot be used in such cases. As these functions are recursive, GNATprove cannot determine that they will terminate. As a result, it will not use their definition in some cases, lest they may be incorrect. To avoid this problem, I have added Terminating annotations to their declarations. These annotations cannot be verified by the tool, but they can easily be discharged by review, so I accepted the check messages with appropriate justification. I have also marked the two functions as ghost, as I will assume that, for efficiency reasons or to control the stack usage, we don't want to use recursive functions in normal code: function Model_Contains (M : access constant Map; K : Positive) return Boolean with Ghost, Annotate => (GNATprove, Terminating); pragma Annotate (GNATprove, False_Positive, "subprogram ""Model_Contains"" might not terminate", "Recursive calls occur on strictly smaller structure"); function Model_Value (M : access constant Map; K : Positive) return Element with Ghost, Annotate => (GNATprove, Terminating), Pre => Model_Contains (M, K); pragma Annotate (GNATprove, False_Positive, "subprogram ""Model_Value"" might not terminate", "Recursive calls occur on strictly smaller structure"); Now that we have defined our specification properties for maps, let's try to implement some functionality. We will start with the easiest one, a Contains function which checks whether there is mapping for a given key in a map. In terms of functionality, it should really return the same value as Model_Contains. However, implementation-wise, we would like to use a loop to avoid introducing recursion. Here, we need to be careful. Indeed, traversing a linked data-structure using a loop generally involves an alias between the traversed structure and the pointer used for the traversal. SPARK supports this use case through the concepts of local observers and borrowers (borrowed from Rust's regular and mutable borrows). What we need here is an observer. Intuitively, it is a local variable which is granted for the duration of its lifetime read-only access to a component of another object. When the observer stays in scope, the actual owner of the object also has read-only access on the object. When the observed goes out of scope, it regains complete ownership (provided it used to have it, and there are no other observers in scope). Let us see how we can use it in our example: function Contains (M : access constant Map; K : Positive) return Boolean with Post => Contains'Result = Model_Contains (M, K); function Contains (M : access constant Map; K : Positive) return Boolean is C : access constant Map := M; begin while C /= null loop pragma Loop_Invariant (Model_Contains (C, K) = Model_Contains (M, K)); if C.Key = K then return True; end if; C := C.Next; end loop; return False; end Contains; Because it is declared with an anonymous access-to-constant type, the assignment to C at its declaration is not considered to be a move, but an observe. In the body of Contains, C will be a read-only alias of M. M itself will retain read-only access to the data it designates, and will regain full ownership at the end of Contains (note that here, it does not change much, as M itself is an access-to-constant type, so it has read-only access to its designated data to begin with). In the body of Contains, we use a loop which searches for a key equal to K in M. We see that, even though C is an observer, we can still assign to it. It is because it is the data designated by C which is read-only, and not C itself. When we assign directly into C, we do not modify the underlying structure, we simply modify the handle. Note that, as usual with loops in SPARK, I had to provide a loop invariant to help GNATprove verify my code. Here it states that K is contained in M, if and only if, it is contained in C. This is true because we have only traversed values different from K until now. We see that, in the invariant, we are allowed to mention M. This is because M retains read-only access to the data it designated during the observe. Let's now write a function to retrieve in normal code the value associated to a key in a map. Since elements can be big, we don't want to copy them, so we should not return the actual value, but rather a pointer to the object stored in the map. For now, let's assume that we are interested in a read-only access to the element. As we have seen above, in the ownership model of SPARK, a read-only access inside an existing data-structure is an observer. So here, we want a function which computes and returns an observer of the input data-structure. This is supported in SPARK, provided the traversed data-structure is itself an access type, using what we call "traversal functions". An observing traversal function takes an access type as its first parameter and returns an anonymous access-to-constant object which should be a component of this first parameter. function Constant_Access (M : access constant Map; K : Positive) return not null access constant Element with Pre => Model_Contains (M, K), Post => Model_Value (M, K) = Constant_Access'Result.all; function Constant_Access (M : access constant Map; K : Positive) return not null access constant Element is C : access constant Map := M; begin while C /= null loop pragma Loop_Invariant (Model_Contains (C, K) = Model_Contains (M, K)); pragma Loop_Invariant (Model_Value (C, K) = Model_Value (M, K)); if C.Key = K then return C.Value; end if; C := C.Next; end loop; raise Program_Error; end Constant_Access; The return type of Constant_Access is rather verbose. It states that it computes an anonymous access-to-constant object (an observer in SPARK) which is not allowed to be null. Indeed, since we know that we have a mapping for K in M when calling Constant_Access, we are sure that there will always be an element to return. This also explains why we have a raise statement as the last statement of Constant_Access: GNATprove is able to prove that this statement is unreachable, and we need that statement to confirm that the bottom of the function cannot be reached without returning a value. The contract of Constant_Access is straightforward, as we are again reimplementing a concept that we already had in the specification (Constant_Access returns a pointer to the result of Model_Value). In the body of Constant_Access, we create a local variable C which observes the data-structure designated by M, just like we did for Contains. When the correct mapping is found, we return an access to the corresponding value in the data-structure. GNATprove will make sure that the structure returned is a part of the first parameter (here M) of the function. Note that this function does not breach the ownership policy of SPARK as what is computed by the Constant_Access is an observer, so it will not be possible to assign it to a component inside another data-structure for example. Now let's assume that we want to modify the value associated to a given key in the data-structure. We can provide a Replace_Element procedure which takes a map M, a key K, and an element V and replaces the value associated to K in M by V. We will write in its postcondition that the value associated to K in M after the call is V (this is not complete, as we do not say anything about other mappings, but for the sake of the explanation, let's stick to something simple). procedure Replace_Element (M : access Map; K : Positive; V : Element) with Pre => Model_Contains (M, K), Post => Model_Contains (M, K) and then Model_Value (M, K) = V; In its body, we want to loop, find the matching pair, and replace the element by the new value V. Here we cannot use an observer to search for the key K, as we want to modify the mapping afterward. Instead, we will use a local borrower. Just like an observer, a borrower takes the ownership of a part of an existing data-structure for the duration of its life-time, but it takes full ownership, in the sense that a borrow has the right to both read and modify the borrowed object. While the borrower is in scope, the borrowed object cannot be accessed directly (there is a provision for reading it in the RM that is used by the tool in particular cases, we will see later). At the end of the scope of the borrower, the ownership returns to the borrowed object. A borrower in SPARK is introduced by the declaration of an object of an anonymous access-to-variable type (note the use of "access Map" instead of "access constant Map"): procedure Replace_Element (M : access Map; K : Positive; V : Element) is X : access Map := M; begin while X /= null loop pragma Loop_Invariant (Model_Contains (X, K)); pragma Loop_Invariant (Pledge (X, (if Model_Contains (X, K) then Model_Contains (M, K) and then Model_Value (M, K) = Model_Value (X, K)))); if X.Key = K then X.Value.all := V; return; end if; X := X.Next; end loop; end Replace_Element; The body of Replace_Element is similar to the body of Contains, except that we modify the borrower before returning from the procedure. However, we see that the loop invariant is more involved. Indeed, as we are modifying M using X, we need to know how the modification of X will affect M. Usually, GNATprove can track this information without help, but when loops are involved, this information needs to be supplied in the loop invariant. To describe how a structure and its borrower are affected, I have used a pledge. The notion of pledges was introduced by researchers from ETH Zurich to verify Rust programs (see the preprint of their work to be published at OOPSLA this year). Conceptually, a pledge is a property which will always hold between the borrowed object and the borrower, no matter the modifications that are made to the borrower. As pledges are not yet supported at a language level in SPARK, it is possible to mark (a part of) an assertion as a pledge by using an expression function which is annotated with a Pledge Annotate pragma: function Pledge (Borrower : access constant Map; Prop : Boolean) return Boolean is (Prop) with Ghost, Annotate => (GNATProve, Pledge); -- Pledge for maps Note that the name of the function could be something other than "Pledge", but the annotation should use the string "Pledge". A pledge function is a ghost expression function which takes a borrower and a property and simply returns the property. When GNATprove encounters a call to such a function, it knows that the property given as a second parameter to the call must be handled as a pledge of the local borrower given as a first parameter. It will attempt to verify that, no matter the modification which may be done to Borrower, the property will still hold. Inside Prop, it is necessary to be able to mention the borrowed object. This is why there is provision for reading it in the SPARK RM, and in fact, it is the only case where the tool will allow a borrowed object to be mentioned. In our example, the pledge of X states that, no matter how X will be modified afterward, if it happens that X has a mapping for K, then M will have the same mapping for K. This is true because we have not encountered K in the previous iterations of the loop, so if we find a mapping in X for K, it will be the first such mapping in M too. Note that a more precise pledge, like Pledge (X, Model_Contains (X, K)), would not be correct. Given a borrower aliasing the root of a subtree in a borrowed object, the pledge relationship expresses what necessarily holds for that subtree, independently of any modifications that may occur to it through the borrower. This is why we need to state the pledge here with an if-expression: "if the borrower X still contains the key K, then M necessarily contains the key K, and both agree on the associated value". After the loop, M has been modified through X. GNATprove does not know anything about it but what can be deduced from the current value of X and information about the relation between M and X supplied by the loop-invariants that contain pledges. These invariants don't completely define the relation between X and M, but they give enough information to deduce the postcondition when a mapping for K is found in X. Since at the last iteration X.Key = K, GNATprove can deduce that Model_Contains (X, K) and Model_Value (X, K) = V holds after the loop. Using the pledges from the loop-invariants, it can infer that we also have Model_Contains (M, K) and Model_Value(M, K) = V. Before we reach the end of this post, we will go one step further. We now have a way to replace an element with another one in the map. It can be used if we want to do a modification inside an element of the map, but it won't be efficient, as the element will need to be copied. We would like to provide a way to find the value associated to a given key in a map, and return an access to it so that it can be modified in-place. As for Constant_Reference, this can be done using a traversal function: function Pledge (Borrower : access constant Element; Prop : Boolean) return Boolean is (Prop) with Ghost, Annotate => (GNATProve, Pledge); -- Pledge for elements function Reference (M : access Map; K : Positive) return not null access Element with Pre => Model_Contains (M, K), Post => Model_Value (M, K) = Reference'Result.all and then Pledge (Reference'Result, Model_Contains (M, K) and then Model_Value (M, K) = Reference'Result.all); Reference returns a mutable access inside the data-structure M. We can see that, in its postcondition, I have added a pledge. Indeed, since the result of Reference is an access-to-variable object, a user of my function can use it to modify M. If I want the tool to be able to deduce anything about such a modification, I need to describe the link between the result of a call to the Reference function, and its first parameter. Here my pledge gives the same information as the postcondition of Replace_Element, that is, that the value designated by the result of the call will be the one mapped from K in M. function Reference (M : access Map; K : Positive) return not null access Element is X : access Map := M; begin while X /= null loop pragma Loop_Invariant (Model_Contains (X, K)); pragma Loop_Invariant (Model_Value (X, K) = Model_Value (M, K)); pragma Loop_Invariant (Pledge (X, (if Model_Contains (X, K) then Model_Contains (M, K) and then Model_Value (M, K) = Model_Value (X, K)))); if X.Key = K then return X.Value; end if; X := X.Next; end loop; raise Program_Error; end Reference; The body of Reference contains the same loop as Constant_Reference, except that I have added a loop invariant, similar to the one in Replace_Element to supply a pledge to X. The specification and verification of pointer-based data-structures is a challenge in deductive verification whether the "pointer" part is implemented as a machine pointer or as an array index (as an example, see our previous post about verifying insertion inside a red-black tree). In addition, SPARK has a strict ownership policy which will prevent completely the use of some (doubly-linked) data-structures, and complicate the writing of usual algorithms on others. However, I think I have demonstrated in this post that is still feasible to write and verify in SPARK some of these algorithms, with comparatively few user-supplied annotations. ]]> Combining GNAT with LLVM https://blog.adacore.com/combining-gnat-with-llvm Tue, 01 Oct 2019 08:24:29 -0400 Arnaud Charlet https://blog.adacore.com/combining-gnat-with-llvm Presenting the GNAT LLVM project At AdaCore labs, we have been working for some time now on combining the GNAT Ada front-end with a different code generator than GCC. The GNAT front-end is particularly well suited for this kind of exercise. Indeed, we've already plugged many different code generators into GNAT in the past, including a Java byte code generator (the old "JGNAT" product for those who remember it), a .NET byte code generator derived from the Java one, a Why3 generator, used at the heart of the SPARK formal verification technology, a SCIL (Statically Checkable Intermediate Language) generator used in our advanced static analyzer CodePeer, and a C back-end used in GNAT CCG (Common Code Generator). This time, we're looking at another general purpose code generator, called LLVM, in order to expand the outreach of Ada to the LLVM ecosystem (be it the compiler itself or other components such as static analysis tools). This work-in-progress research project is called "GNAT LLVM" and is meant to show the feasibility of generating LLVM bytecode for Ada and to open the LLVM ecosystem to Ada, including tools such as KLEE, that we are also planning to work with and add Ada support for. Note that we are not planning on replacing any existing GNAT port based on GCC, so this project goes in addition rather than in replacement. Technical Approach We decided on a "pure" LLVM approach that's as easy to integrate and fit into the LLVM ecosystem as possible, using the existing LLVM API directly, while at the same time doing what we do best: write Ada code! So we are using the LLVM "C" API and generate automatically Ada bindings via the GCC -fdump-ada-spec switch and a bit of postprocessing done in a python script, that we can then call directly from Ada, which allows us to both easily traverse the GNAT tree and generate LLVM instructions, all in Ada. By the way, if you know about the DragonEgg project then a natural question would be "why are you starting a GNAT LLVM project from scratch instead of building on top of DragonEgg?". If you want to know the answer, check the file README.dragonegg in the repository! Next Steps We have just published the GNAT LLVM tool sources licensed under GPLv3 on GitHub for hobbyists and researchers to experiment with it and give us their feedback. If you are interested, give it a try, and let us know via Issues and Pull Requests on GitHub how it works for you, or by leaving comments at the bottom of this page! One exciting experiment would be to compile Ada code using the webassembly LLVM back-end for instance! ]]> The Make with Ada competition is back! https://blog.adacore.com/the-make-with-ada-competition-is-back Thu, 12 Sep 2019 08:50:57 -0400 Emma Adby https://blog.adacore.com/the-make-with-ada-competition-is-back AdaCore’s fourth annual Make with Ada competition launched this week with over$8K in cash and prizes to be awarded for the most innovative embedded systems projects developed using Ada and/or SPARK.

The contest runs from September 10, 2019, to January 31, 2020, and participants can register on the Hackster.io developer platform here.

What’s new?

Based on feedback from previous participants, we changed the competition evaluation criteria so projects will now be judged on:

• Software quality - Does the software meet its requirements?;

• Openness - Is the project open source?; and

• “Buzz factor” - Does it have the wow effect to appeal to the software community?

Further information about the judging criteria is available here

We’ve also increased the amount of prizes to give more projects a chance to win:

• One First Prize, in the amount of 2000 (two thousand) USD

• Ten Finalist Prizes, in the amount of 600 (six hundred) USD each

• One Student-only Prize (an Analog Discovery 2 Pro Bundle worth 299.99 USD) will go to the best-ranking student finalist. A project submitted by a student is eligible for both the Student-only Prize and the cash prizes.

Award winners will be announced in March 2020 and project submissions will be evaluated by a judging panel consisting of Bill Wong, Senior Technology Editor at Electronic Design, and Fabien Chouteau, AdaCore software engineer, and author of the Make with Ada blog post series.

Don’t forget that the new and enhanced GNAT Community 2019 is also available for download for use in your projects!

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The Ada Community gathered recently around a new exciting initiative - an Ada Virtual Conference, to present Ada-related topics in a 100% remote environment.

The first such conference took place on August, 10th 2019, around the topic of the new features in Ada 202x. The conference took the form of a video/audio chat based on the open source platform jitsi.org. No registration required, just access over a web browser or mobile application, with the possibility to participate anonymously. The presentation is just short of 25 minutes and is available on YouTube, Vimeo or via DropBox.

Note that while the talk presents the current draft for Ada 202x, some features are still in discussion and may not make it to the standard, or with a different syntax or set of rules. That's in particular the case for all parallelism-related features and iterators (that is, up to slide 16 in the talk) which are being revisited by a group of people within AdaCore, in order to submit recommendations to the Ada Rapporteur Group next year.

As the talk concludes, please contribute to the new Ada/SPARK RFCs website if you have ideas about the future of the language!

Now the Ada Virtual Conference has a dedicated website, where anyone can vote for the topic of the next event. We invite you to participate!

Image by Tomasz Mikołajczyk from Pixabay.

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Secure Use of Cryptographic Libraries: SPARK Binding for Libsodium https://blog.adacore.com/secure-use-of-cryptographic-libraries-spark-binding-for-libsodium Tue, 03 Sep 2019 07:54:00 -0400 Isabelle Vialard https://blog.adacore.com/secure-use-of-cryptographic-libraries-spark-binding-for-libsodium

The challenge faced by cryptography APIs is to make building functional and secure programs easy for the user. Even with good documentation and examples, this remains a challenge, especially because incorrect use is still possible. I made bindings for two C cryptography libraries, TweetNaCl (pronounce Tweetsalt) and Libsodium, with the goal of making this binding easier to use than the original API by making it possible to detect automatically a large set of incorrect uses. In order to do this, I did two bindings for each library: a low-level binding in Ada, and a higher level one in SPARK, which I call the interface. I used Ada strong-typing characteristics and SPARK proofs to enforce a safe and functional use of the subprograms in the library.

In this post I will explain the steps I took to create these bindings, and how to use them.

Steps to create a binding

I will use one program as example: Crypto_Box_Easy from Libsodium, a procedure which encrypts a message.

At first I generated a binding using the Ada spec dump compiler:

gcc -c -fdump-ada-spec -C ./sodium.h

Which gives me that function declaration:

function crypto_box_easy
(c : access unsigned_char;
m : access unsigned_char;
mlen : Extensions.unsigned_long_long;
n : access unsigned_char;
pk : access unsigned_char;
sk : access unsigned_char) return int  -- ./sodium/crypto_box.h:61
with Import => True,
Convention => C,
External_Name => "crypto_box_easy";

Then I modified this binding: First I changed the types used and I added in and out parameters. I removed the access parameter. Scalar parameters with out mode are passed using a temporary pointer on the C side, so it works even without explicit pointers. For unconstrained arrays like Block8 it is more complex. In Ada unconstrained arrays are represented by what is called fat pointers, that is to say a pointer to the bounds of the array and the pointer to the first element of the array. In C the expected parameter is a pointer to the first element of the array. So a simple binding like this one should not work. What saves the situation is this line which forces passing directly the pointer to the first element:

with Convention => C;

Thus we go from a low-level language where the memory is indexed by pointers to a typed language like Ada:

function crypto_box_easy
(c    :    out Block8;
m    : in     Block8;
mlen : in     Uint64;
n    : in out Block8;
pk   : in     Block8;
sk   : in     Block8) return int  -- ./sodium/crypto_box.h:61
with Import => True,
Convention => C,
External_Name => "crypto_box_easy";

After this, I created an interface in SPARK that uses this binding: the goal is to make the same program as in the binding with some modifications. Some useless parameters will be deleted. For instance the C program often asks for an array and the length of this array (like m and mlen), which is useless since the length can be found with the attribute 'Length. Functions with out parameters must be changed into procedures to comply with SPARK rules. Finally new types can be created to take advantage of strong typing, as well as preconditions and postconditions.

procedure Crypto_Box_Easy
(C  :    out Cipher_Text;
M  : in     Plain_Text;
N  : in out Box_Nonce;
PK : in     Box_Public_Key;
SK : in     Box_Secret_Key)
with
Pre => C'Length = M'Length + Crypto_Box_MACBYTES
and then Is_Signed (M)
and then Never_Used_Yet (N);

How to use strong typing, and why

Most of the parameters required by these programs are arrays. Some arrays for messages, others for key, etc. The type Block8 could be enough to represent them all:

type Block8 is array (Index range <>) of uint8;

But then anyone could use a key as a message, or a message as a key. To avoid that, in Libsodium I derived new types from Block8. For instance, Box_Public_Key and Box_Secret_Key are the key types used by the Crypto_Box_* programs. The type for messages is Plain_Text, and the type for messages after encryption is Cipher_Text. Thus I take advantage of Ada's strong-typing characteristics in order to enforce the right use of the programs and their parameters.

With TweetNaCl, I did things a bit differently: I created the different types directly in the binding, for the same result. Since TweetNaCl is a very small library, it was faster that way. In Libsodium I chose to let my first binding stay as close as possible to the generated one, and to focus on the interface where I use strong-typing and contracts (preconditions and postconditions).

Preconditions and postconditions

Preconditions and postconditions serve the same purpose as derived types: they  enforce a specific use of the programs. There are two kinds of conditions:

The first kind are mostly conditions on array's length. They are here to ensure the program will not fail. For instance:

Pre => C'Length = M'Length + Crypto_Box_MACBYTES;

This precondition says that the cipher text should be exactly Crypto_Box_MACBYTES bytes longer than the message that we want to encrypt. If this condition is not filled then execution will fail. Note that we can reference C'Length in the precondition, even though C is an out parameter, because the length attribute of an out parameter that is of an array type is available when the call is made, so we can reason about it in our precondition.

The other kind of conditions is used to avoid an unsafe use of the programs. For instance, Crypto_Box uses a Nonce. A Nonce is a small array used as a complement to a key. In theory, to be safe, a key should be long, and used only for one message. However, it is costly to generate a new long key for each message. So we use a long key for every message with a Nonce which is different for each message, but easy to generate. Thus the encryption is safe, but only if we remember to use a different Nonce for each message. To ensure that, I wrote this precondition:

Pre => Never_Used_Yet (N)

Never_Used_Yet is a ghost function, it means a function that doesn't affect the program's behavior.

type Box_Nonce is limited private;
function Never_Used_Yet (N: Box_Nonce) return Boolean with Ghost;

When GNATprove is used, it sees that procedure Randombytes (Box_Nonce: out N), the procedure that randomly generates the Nonce, has the postcondition Never_Used_Yet (N). So it deduces that when the Nonce is first generated, Never_Used_Yet (N) is true. Thus the first time N is used by Crypto_Box_Open, the precondition is valid. But if N is used a second time, GNATprove cannot prove Never_Used_Yet (N) is still true (because N as the parameter in out, so it could have been changed). That's why it cannot prove a program that calls Crypto_Box_Open twice with the same Nonce:

There are ways around this condition: for instance one could copy a random generated Nonce many times to use it on different message. To avoid that, Box_Nonce is declared as limited private: it cannot be copied and GNATprove cannot prove a copied Nonce has the same Never_Used_Yet property as a generated one.

Box_Nonce and Never_Used_Yet are declared in the private part of the package with SPARK_Mode Off so that GNATprove treats them as opaque entities:

private
pragma SPARK_Mode (Off);

type Box_Nonce is new Block8 (1 .. Crypto_Box_NONCEBYTES);

function Never_Used_Yet (N : Box_Nonce) return Boolean is (True);

Never_Used_Yet always returns true. It is a fake implementation, what matters is that it is hidden for the proof. It works at runtime because it is always used as a positive condition. As a program requirement it is always used as "Never_Used_Yet (N)" and never "not Never_Used_Yet (N)" so the conditions are always valid, and Never_Used_Yet doesn't affect the program's behavior, even if contracts are executed at runtime.

Another example of preconditions made with a ghost function is the function Is_Signed (M : Plain_Text). When you want to send an encrypted message to someone, you want this person to be able to check if this message is from you, so no one will be able to steal your identity. To do this, you have to sign your message with Crypto_Sign_Easy, before encrypting it with Crypto_Box_Easy. Trying to skip the signing step leads to a proof error:

How to use Libsodium_Binding

The repository contains:

• The project file libsodium.gpr

• The library directory lib

• The common directory which contains:

◦ The libsodium_binding package, a low-level binding in Ada made from the files generated using the Ada spec dump compiler.

◦ The libsodium_interface package, a higher level binding in SPARK which uses libsodium_binding.

• include, a directory which contains the headers of libsodium.

• libsodium_body, a directory which contains the bodies.

• outside_src, a directory which contains the headers that were removed from include, to fix a problem of double definition.

• thin_binding, a directory which contains the binding generated using the Ada spec dump compiler.

• The test directory which contains tests for each group of functions.

• A testsuite which verifies the same tests as the ones in the test directory.

• The examples directory, with examples that use different groups of programs together. It also contains a program where a Nonce is used twice, and as expected it fails at proof stage.

outside_src and thin_binding are not used for the binding, but I let them in the repository because it shows what I changed from the original libsodium sources and the Ada generated binding.

This project is a library project so directory lib is the only thing necessary.

How to use TweetNaCl_Binding

The repository contains:

• The project file tweetnacl.gpr

• The common directory which contains:

◦ The tweetnacl_binding package, a low-level binding in Ada made from the files generated using the Ada spec dump compiler.

◦ The tweetnacl_interface package, a higher level binding in SPARK which uses tweetnacl_binding.

◦ tweetnacl.h and tweetnacl.c, the header and the body of the library

◦ randombytes.c, which holds randombytes, a program to generate arrays.

• The test directory: test1 and test1b are functional examples of how to use tweetnacl main programs, the others are examples of what happens if you give an array with the wrong size, if you try to use the same nonce twice etc. They fail either at execution or at proof stage.

To use this binding, you just have to include the common directory in the Sources of your project file.

]]>
Proving a simple program doing I/O ... with SPARK https://blog.adacore.com/proving-a-simple-program-doing-io Tue, 09 Jul 2019 08:37:05 -0400 Joffrey Huguet https://blog.adacore.com/proving-a-simple-program-doing-io

The functionality of many security-critical programs is directly related to Input/Output (I/O). This includes command-line utilities such as gzip, which might process untrusted data downloaded from the internet, but also any servers that are directly connected to the internet, such as webservers, DNS servers and so on. However, I/O has received little attention from the formal methods community, and SPARK also does not help the programmer much when it comes to I/O. SPARK has been used to debug functions from the standard library Ada.Text_IO, but this approach lacked support for error handling and didn't allow going up to the application level.

As an example, take a look at the current specification of Ada.Text_IO.Put, which only recently has been annotated with some SPARK contracts:

procedure Put
(File : File_Type;
Item : String)
with
Pre    => Is_Open (File) and then Mode (File) /= In_File,
Global => (In_Out => File_System);

(We have suppressed the postcondition of this function, which talks about line and page length, a functionality of Ada.Text_IO which is not relevant to this blog post.)

We can see that Put has a very light contract that protects against some errors, such as calling it on a file that has not been opened for writing, but not against the other many possible errors related to writing, such as a full disk or file permissions. If such an error occurs, Put raises an exception, but SPARK does not allow catching exceptions. So in the above form, even a proved SPARK program that uses Put may terminate with an unhandled exception.

Moreover, Put does not specify what exactly is written. For example, one cannot prove in SPARK that two calls to Put with the same File argument write the concatenation of the two Item arguments to the file.

This second problem can probably be solved using a stronger contract (though it would be difficult to do so in Ada.Text_IO, whose interface must respect the Ada standard), but together with the first point it means that, even if our program is annotated with a suitable contract and proved, we can only say something informal like “if the program doesn’t crash because of unexpected errors, it respects its contract”.

In this blogpost, we propose to solve these issues as follows. We replace Put by a procedure Write which reports errors via its output Has_Written:

procedure Write
(Fd          : int;
Buf         : Init_String;
Num_Bytes   : Size_T;
Has_Written : out ssize_t);

We now can annotate Write with a suitable postcondition that explains exactly what has been written to the file descriptor, and that a negative value of Has_Written signals an error.

It is no accident that the new procedure looks like the POSIX system call write. In fact we decided to base this experiment on the POSIX API that consists of open, close, read and write, and simply added very thin SPARK wrappers around those system calls. The advantage of this approach is that system calls never crash (unless there are bugs in the kernel, of course) - they simply flag errors via their return value, or in our case, the out parameter Has_Written. This means that, assuming our program contains a Boolean variable Error that is updated accordingly after invoking the system calls, we can now formally write down the contract of our program in the form:

if not Error then ...

And if one thinks about it, this is really the best any program dealing with I/O can hope for, because some errors cannot be predicted or avoided by the programmer, such as trying to write to a full disk or trying to open a file that does not exist.

We could further refine the above condition to distinguish the behavior of the program depending on the error that was encountered, but we did not do this in the work described here.

The other advantage of using a POSIX-like API is that porting existing programs from C to this API would be simpler.

We validated this approach by writing a SPARK clone of the “cat” utility. We were able to prove that cat copies the content of the files in its argument to stdout, unless errors were encountered. The project is available following this link.

How to represent the file system?

The main interest of the library is to be able to write properties about content, which means that we have to represent this content through something (global variables, state abstraction…). We decided to use maps, that link a file descriptor to an unbounded string representing the content of the corresponding file. Formal_Hashed_Maps were convenient to use because they come with many functions that compare two different maps (e.g Keys_Included_Except, or Elements_Equal_Except).

In Ada, even Unbounded_String (from Ada.Strings.Unbounded) are somehow bounded: the maximum length of Unbounded_String is Natural’Last, a length that could be exceeded by certain files. We had to create another library, that relies on Ada.Strings.Unbounded, but would accept appending two unbounded strings whose lengths are maximal. The choice we made was to drop any character that would make the length overflow, e.g. the append function has the following contract:

function "&" (L, R : My_Unbounded_String) return My_Unbounded_String with
Global         => null,
Contract_Cases =>
(Length (L) = Natural'Last               =>
To_String ("&"'Result) = To_String (L),
--  When L already has maximal length

Length (L) < Natural'Last
and then
Length (R) <= Natural'Last - Length (L) =>
To_String ("&"'Result) = To_String (L) & To_String (R),
--  When R can be appended entirely

others                                  =>
To_String ("&"'Result)
= To_String (L) & To_String (R) (1 .. Natural'Last - Length (L)));
--  When R cannot be appended entirely

Also, what you can see here is that all properties of unbounded strings are expressed with the conversion to String type: this allows to use the existing axiomatization of strings (through arrays) instead of redefining one.

Another design choice was made here: what’s the content of a file we opened in read mode? In stdio.ads, we considered that the content of this file is strictly what we read from it. Because there is no way to know the content of the file, we decided to implement cat as a procedure that would “write whatever it reads”. Any given file could be modified while reading, or, for the case of stdin, parent process may also read part of the data.

The I/O library

As said before, the library consists of thin bindings to system calls. Interfacing scalar types (int, size_t, ssize_t) was easy. We also wanted to model more precisely instantiation of buffers. Indeed, when calling Read procedure, the buffer might not be initialized entirely after the call; it is also possible to call Write on a partially initialized String, to copy the first n values that are initialized. The current version of SPARK requires that arrays (and thus Strings) are fully initialized. More details are available in the SPARK User's Guide. A new proposed evolution of the language (see here) is already prototyped in SPARK and allows manipulating safely partially initialized data. In the prototype, we declare a new type, Init_Char, that will have the Init_By_Proof annotation.

subtype Init_Char is Character;
pragma Annotate(GNATprove, Init_By_Proof, Init_Char);

type Init_String is array (Positive range <>) of Init_Char;

With this type declaration, it is now possible to use the attribute ‘Valid_Scalars on Init_Char variables or slices of Init_String variables in Ghost code to specify which scalars have been initialized

The next part was writing the bindings to C functions. An example, for Read, is the following:

function C_Read (Fd : int; Buf : System.Address; Size : size_t; Offset : off_t) return ssize_t;

procedure Read (Fd : int; Buf : out Init_String; Has_Read : out ssize_t) is
begin
end Read;

Other than hiding the use of addresses from SPARK, this part was not very difficult.

The final part was adding the contracts to our procedures.

Firstly, there are no preconditions. System calls may return errors, but they will accept any parameter in input. The only precondition we have in the library is a precondition on Write, that states that the characters that we want to write in the file are initialized.

Secondly, every postcondition is a case expression, where we give properties for each possible return value, e.g:

procedure Open (File : char_array; Flags : int; Fd : out int) with
Global => (In_Out => (FD_Table, Errors.Error_State, Contents)),
Post   =>
(case Fd is
when -1                      => Contents'Old = Contents,
when 0 .. int (OPEN_MAX - 1) =>
Length (Contents'Old) + 1 = Length (Contents)
and then
Contains (Contents, Fd)
and then
Length (Element (Contents, Fd)) = 0
and then
not Contains (Contents'Old, Fd)
and then
Model (Contents'Old) <= Model (Contents)
and then
M.Keys_Included_Except (Model (Contents),
Model (Contents'Old),
Fd),
when others                  => False);

The return value of Open will be either -1, which corresponds to an error, or a natural value (in my case, OPEN_MAX is equal to 1023, 1024 being the maximum number of files that can be open at the same time on my machine). If an error occured, the Contents map is the same as before. If an appropriate file descriptor is returned, the postcondition states that the new Contents map has the same elements as before, plus a new empty unbounded string associated with the file descriptor. With regard to the functional model, these contracts are complete.

The Cat program

The cat program is split in two different parts: the main program that opens/closes file(s) in its argument, and calls Copy_To_Stdout. This procedure will read from the input file and write to stdout what it read.

Since errors in I/O can happen, we propagate them using a status flag from nested subprograms to the main program and handle them there. This point is the first difference with Ada.Text_IO.

The other difference is the presence of postconditions about data, for example in postcondition of Copy_To_Stdout:

procedure Copy_To_Stdout (Input : int; Err : out Int) with
Post =>
(if Err = 0 then Element (Contents, Stdout)
= Element (Contents’Old, Stdout) & Element (Contents, Input));

This postcondition is the only functional contract we have for cat, this is why it is so important. It states that if no error occurred, the content of stdout is equal to its value before calling Copy_To_Stdout appended to the content we read from the input.

If we wanted to write a more precise contract, the main difficulty would be to handle the cases where an error occured. For example, if we call cat on three different files, and one error occurs when copying the second file to stdout, we have no contract about the content of Stdout, and everything becomes unprovable. Adding contracts for these cases would require work on slices and sliding and to do even more case-splitting, which would add more difficulty for the provers.

Type definitions and helper subprograms to define the library take about 200 lines of code. The library itself has around 100 lines of contracts. The cat program has 100 lines of implementation and 1200 lines of Ghost code in order to prove everything. Around 1000 verification conditions for the entire project (I/O library + cat + lemmas) are discharged by the solvers to prove everything with auto-active proof.

Conclusion

Cat looks like a simple program, and it is. But being able to prove correctness of cat shows that we are able to reason about contents and copying of data (here, between file descriptors), something which is necessary and largely identical for more complex applications like network drivers or servers that listen on sockets. We will be looking at extending our approach to code manipulating (network) sockets.

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I am an Associate Professor at Polytechnic University of Madrid’s (Universidad Politécnica de Madrid / UPM) in the Department of Architecture and Technology of Computer Systems. For the past several years I have been directing a team of colleagues and students in the development of a UPMSat-2 microsatellite. The project originally started in 2013 as a follow-to the UPM-SAT 1, launched by an Ariane-4 in 1995.

The UPMSat-2 weighs 50kg, and its geometric envelope is a parallelepiped with a base measuring 0.5m x 0.5m and a height of 0.6m. The microsatellite is scheduled to be launched September 9, 2019 on a Vega launcher, and is expected to be operational for two years.

The primary goals of the project were:

• to improve the knowledge of the project participants, both professors and students;
• to demonstrate UPM’s capabilities in space technology;
• to design, develop, integrate, test, launch and operate a microsatellite in orbit from within a university environment; and
• to develop a qualified space platform that can be used for general purpose missions aimed at educational, scientific and technological demonstration applications.

The project encompasses development of the software together with the platform, thermal control, attitude control, and other elements. In 2014, we selected AdaCore’s GNAT cross-development environment for the UPMSat-2 microsatellite project’s real-time on-board and ground control software.

While Java is the primary language used to teach programming at UPM, Ada was chosen as the main programming language for our project because we considered it the most appropriate to develop high-integrity software. In total, the on-board software consists of more than 100 Ada packages, comprising over 38K lines of code. For the altitude control subsystem, we used C code that was generated automatically from Simulink® models (there are 10 source files in C, with a total of about 1,600 lines of code). We also used database and XML interfaces to support the development of the ground control software.

Since most of our students are last-year or graduate students, they generally have programming experience. However, they did not have experience with embedded systems or real-time programming, and none of them had any previous experience with GNAT, SPARK or Ada.

To teach Ada to our students, we provided them with John Barnes’ Programming in Ada 2012 textbook and spent a fair amount of time with it in the laboratory. The difficult part was not in understanding and using Ada, but rather in understanding the software issues and programming style associated with concurrency, exceptions, real-time scheduling, and large system design. Fortunately, Ada’s high-level concurrency model, simple exception facility, real-time support, and its many features for “programming in the large” (packages, data abstraction, child libraries, generics, etc.) helped to address these difficulties.

We also used a GNAT feature that saved us a lot of tedious coding time - the Scalar_Storage_Order attribute. The UPMSat-2’s on-board computer is big-endian and the ground computer is little-endian. Therefore, we had to decode and encode every telemetry and telecommand message to deal with the endianness mismatch. I learned about the Scalar_Storage_Order feature at an AdaCore Tech Day, and it works really well, even for 12-bit packet types.

Although it would have been nice to have some additional tool support for things like database and XML interfacing, we found the GNAT environment very intuitive and especially appreciated the GPS IDE; it’s a great tool for developing software.

I have been in love with Ada for a long time. I learned programming with Pascal and concurrent Pascal (as well as Fortran and COBOL); I find it frustrating that, at many academic institutions, modular, strongly-typed, concurrent languages such as Ada have sometimes been replaced by others that have much weaker support for good programming practices.

While I do not teach programming at UPM, my research group tries to use Ada whenever possible, because we consider it the most appropriate programming language for illustrating the concepts of real-time and embedded systems. I have to say that most of my students have also fallen in love with Ada. Our graduate students in-particular appreciate the value and reliability that Ada brings to their final projects.

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Ever wished that Ada was more this and less that? Or that SPARK had such-and-such feature for specifying your programs? Then you're not alone. The Ada-Comment mailing list is one venue for Ada language discussions, but many of us at AdaCore have felt the need for a more open discussion and prototyping of what goes in the Ada and SPARK languages. That's the main reason why we've set up a platform to collect, discuss and process language evolution proposals for Ada and SPARK. The platform is hosted on GitHub, and uses GitHub built-in mechanisms to allow people to propose fixes or evolutions for Ada & SPARK,or give feedback on proposed evolutions.

For SPARK, the collaboration between Altran and AdaCore allowed us to completely redesign the language as a large subset of Ada, including now object orientation (added in 2015), concurrency (added in 2016) and even pointers (now available!), but we're reaching the point where catching up with Ada cannot lead us much further, and we need a broader involvement from users to inform our strategic decisions.

Regarding Ada, the language evolution has always been a collective effort by an international committee, but here too we feel that more user involvement would be beneficial to drive future evolution, including for the upcoming Ada 202X version. Note that there is no guarantee that changes discussed and eventually prototyped & implemented will ever make it into the Ada standard, even though AdaCore will do its best to collaborate with the Ada Rapporteur Group (ARG).

You will see that we've started using the RFC process internally. That's just the beginning. We plan to use this platform much more broadly within AdaCore and the Ada community to evolve Ada and SPARK in the future. Please join us in that collective effort if you are interested!

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Using Pointers in SPARK https://blog.adacore.com/using-pointers-in-spark Thu, 06 Jun 2019 08:22:00 -0400 Claire Dross https://blog.adacore.com/using-pointers-in-spark

I joined the SPARK team during the big revamp leading to the SPARK 2014 version of the proof technology. Our aim was to include in SPARK all features of Ada 2012 that did not specifically cause problems for the formal verification process. Since that time, I have always noticed the same introduction to the SPARK 2014 language: a subset of Ada, excluding features not easily amenable to formal analysis. Following this sentence was a list of (more notable) excluded features. Over the years, this list of features has started to shrink, as (restricted) support for object oriented programming, tasking and others were added to the language. Up until now, the most notable feature that was still missing, to my sense, was pointer support (or support for access types as they are called in Ada). I always thought that this was a feature we were never going to include. Indeed, absence of aliasing is a key assumption of the SPARK analysis, and removing it would induce so much additional annotation burden for users, that it would make the tool hardly usable. This was what I believed, and this was what we kept explaining to users, who were regretting the absence of an often used feature of the language.

I think it was work on the ParaSail language, and the emergence of the Rust language, that first made us look again into supporting this feature. Pointer ownership models did not appear explicitly within Rust, but Rust made the restrictions associated with them look tractable, and maybe even desirable from a safety point of view. So what is pointer ownership? Basically, the idea is that an object designated by a pointer always has a single owner, which retains the right to either modify it, or (exclusive or) share it with others in a read-only way. Said otherwise, we always have either several copies of the pointer which allow only reading, or only a single copy of the pointer that allows modification. So we have pointers, but in a way that allows us to ignore potential aliases… What a perfect fit for SPARK! So we began to look into how ownership rules were enforced in Rust and ParaSail, and how we could adapt some of them for Ada without introducing too many special cases and new annotations. In this post, I will show you what we came up with. Don’t hesitate to comment and tell us what you like / don’t like about this feature.

The main idea used to enforce single ownership for pointers is the move semantics of assignments. When a pointer is copied through an assignment statement, the ownership of the pointer is transferred to the left hand side of the assignment. As a result, the right hand side loses the ownership of the object, and therefore loses the right to access it, both for writing and reading. On the example below, the assignment from X to Y causes X to lose ownership on the value it references. As a result, the last assertion, which reads the value of X, is illegal in SPARK, leading to an error message from GNATprove:

procedure Test is
type Int_Ptr is access Integer;
X : Int_Ptr := new Integer'(10);
Y : Int_Ptr;                --  Y is null by default
begin
Y := X;                     --  ownership of X is transferred to Y
pragma Assert (Y.all = 10); --  Y can be accessed
Y.all := 11;                --  both for reading and writing
pragma Assert (X.all = 11); --  but X cannot, or we would have an alias
end Test;
test.adb:9:20: insufficient permission on dereference from "X"
test.adb:9:20: object was moved at line 6

In this example, we can see the point of these ownership rules. To correctly reason about the semantics of a program, SPARK needs to know, when a change is made, what are the objects that are potentially impacted. Because it assumes that there can be no aliasing (at least no aliasing of mutable data), the tool can easily determine what are the parts of the environment that are updated by a statement, be it a simple assignment, or for example a procedure call. If we were to break this assumption, we would need to either assume the worst (that all references can be aliases of each other) or require the user to explicitly annotate subprograms to describe which references can be aliased and which cannot. In our example, SPARK can deduce that an assignment to Y cannot impact X. This is only correct because of ownership rules that prevent us from accessing the value of X after the update of Y.

Note that a variable which has been moved is not necessarily lost for the rest of the program. Indeed, it is possible to assign it again, restoring ownership. For example, here is a piece of code that swaps the pointers X and Y:

declare
Tmp : Int_Ptr := X; --  ownership of X is moved to Tmp
--  X cannot be accessed.
begin
X := Y; --  ownership of Y is moved to X
--  Y cannot be accessed
--  X is unrestricted.
Y := Tmp; --  ownership of Tmp is moved to Y
--  Tmp cannot be accessed
--  Y is unrestricted.
end;

This code is accepted by the SPARK tool. Intuitively, we can see that writing at top-level into X after it has been moved is OK, since it will not modify the actual owner of the moved value (here Tmp). However, writing in X.all is forbidden, as it would affect Tmp (don’t hesitate to look at the SPARK Reference Manual if you are interested in the formal rules of the move semantics). For example, the following variant is rejected:

declare
Tmp : Int_Ptr := X; --  ownership of X is moved to Tmp
--  X cannot be accessed.
begin
X.all := Y.all;
insufficient permission on dereference from "X"
object was moved at line 2

Moving is not the only way to transfer ownership. It is also possible to borrow the ownership of (a part of) an object for a period of time. When the borrower disappears, the borrowed object regains the ownership, and is accessible again. It is what happens for example for mutable parameters of a subprogram when the subprogram is called. The ownership of the actual parameter is transferred to the formal parameter for the duration of the call, and should be returned when the subprogram terminates. In particular, this disallows procedures that move some of their parameters away, as in the following example:

type Int_Ptr_Holder is record
Content : Int_Ptr;
end record;

procedure Move (X : in out Int_Ptr_Holder; Y : in out Int_Ptr_Holder) is
begin
X := Y; --  ownership of Y.Content is moved to X.Content
end Move;
insufficient permission for "Y" when returning from "Move"
object was moved at line 3

Note that I used a record type for the type of the parameters. Indeed, SPARK RM has a special wording for in out parameters of an access type, stating that they are not borrowed but moved on entry and on exit of the subprogram. This allows us to move in out access parameters, which otherwise would be forbidden, as borrowed top-level access objects cannot be moved.

The SPARK RM also allows declaring local borrowers in a nested scope by using an anonymous access type:

declare
Y : access Integer := X;    --  Y borrows the ownership of X
--  for the duration of the declare block
begin
pragma Assert (Y.all = 10); --  Y can be accessed
Y.all := 11;                --  both for reading and writing
end;
pragma Assert (X.all = 11);   --  The ownership of X is restored,
--  it can be accessed again

But this is not supported yet by the proof tool, as it raises the complex issue of tracking modifications of X that were done through Y during its lifetime:

local borrower of an access object is not yet supported

It is also possible to share a single reference between several readers. This mechanism is called observing. When a variable is observed, both the observed object and the observer retain the right to read the object, but none can modify it. As for borrowing, when the observer disappears, the observed object regains the permissions it had before (read-write or read-only). Here is an example. We have a list L, defined as a recursive pointer-based data structure in the usual way.  We then observe its tail by introducing a local observer N using an anonymous access to constant type. We then do it again to observe the tail of N:

declare
N : access constant List := L.Next; -- observe part of L
begin
declare
M : access constant List := N.Next; -- observe again part of N
begin
pragma Assert (M.Val = 3); --  M can be read
pragma Assert (N.Val = 2); --  but we can still read N
pragma Assert (L.Val = 1); --  and even L
end;
end;
L.Next := null; --  all observers are out of scope, we can modify L

We can see that the three variables retain the right to read their content. But it is OK as none of them is allowed to update it. When no more observers exist, it is again possible to modify L.

In addition to single ownership, SPARK restricts the use of access types in several ways. The most notable one is that SPARK does not allow general access types. The reason is that we did not want to deal with accesses to variables defined on the stack and accessibility levels. Also, access types cannot be stored in subcomponents of tagged types, to avoid having access types hidden in record extensions.

To get convinced that the rules enforced by SPARK still allow common use cases, I think the best is to look at an example.

A common use case for pointers in Ada is to store indefinite types inside data-structures. Indefinite types are types whose subtype is not known statically. It is the case for example for unconstrained arrays. Since the size of an indefinite type is not known statically, it is not possible to store it inside a data-structure, such as another array, or a record. For example, as strings are arrays, it is not possible to create an array that can hold strings of arbitrary length in Ada. The usual work-around consists in adding an indirection via the use of pointers, storing pointers to indefinite elements inside the data structure. Here is an example of how this can now be done in SPARK, for a minimal implementation of a dictionary. A simple vision of a dictionary is an array of strings. Since strings are indefinite, I need to define an access type to be allowed to store them inside an array:

type Word is not null access String;
type Dictionary is array (Positive range <>) of Word;

We can then search for a word in a dictionary. The function below is successfully verified in SPARK. In particular, SPARK is able to verify that no null pointer dereference may happen, due to Word being an access type with null exclusion:

function Search (S : String; D : Dictionary) return Natural with
Post => (Search'Result = 0 and then
(for all I in D'Range => D (I).all /= S))
or else (Search'Result in D'Range and then D (Search'Result).all = S) is
begin
for I in D'Range loop
pragma Loop_Invariant
(for all K in D'First .. I - 1 => D (K).all /= S);
if D (I).all = S then
return I;
end if;
end loop;
return 0;
end Search;

Now imagine that I want to modify one of the words stored in my dictionary. The words may not have the same length, so I need to replace the pointer in the array. For example:

My_Dictionary (1) := new String'("foo");
pragma Assert (My_Dictionary (1).all = "foo");

But this is not great, as now I have a memory leak. Indeed, the value previously stored in My_Dictionary is no longer accessible and it has not been deallocated. The SPARK tool does not currently complain about this problem, even though the SPARK definition says it should (it has not been implemented yet). But let’s try to correct our code nevertheless by storing the value previously in dictionary in a temporary and deallocating it afterward. First I need a deallocation function. In Ada, they can be obtained by instantiating the generic procedure Ada.Unchecked_Deallocation with the appropriate types (note that as access objects are set to null after deallocation, I had to introduce a base type for Word without the null exclusion constraint):

type Word_Base is access String;
subtype Word is not null Word_Base;
procedure Free is new Ada.Unchecked_Deallocation (String, Word_Base);

Then, I can try to do the replacement:

declare
Temp : Word_Base := My_Dictionary (1);
begin
My_Dictionary (1) := new String'("foo");
Free (Temp);
pragma Assert (My_Dictionary (1).all = "foo");
end;

Unfortunately this does not work, the SPARK tool complains with:

test.adb:36:37: insufficient permission on dereference from "My_Dictionary"
test.adb:36:37: object was moved at line 3

Where line 31 is the line where Temp is defined and 36 is the assertion. So, what is happening? In fact, this is due to the way checking of single ownership is done in SPARK. As the analysis used for this verification is not value dependent, when a cell of an array is moved, the tool is never able to determine whether or not ownership to an array cell has been regained. As a result, if an element of an array is moved away, the array will never become readable again unless it is assigned as a whole. Better to avoid moving elements of an array in these conditions, right? So what can we do? If we cannot move, what about borrowing... Let us try with an auxiliary Swap procedure:

procedure Swap (X, Y : in out Word_Base) with
Pre => X /= null and Y /= null,
Post => X /= null and Y /= null
and X.all = Y.all'Old and Y.all = X.all'Old
is
Temp : Word_Base := X;
begin
X := Y;
Y := Temp;
end Swap;

declare
Temp : Word_Base := new String'("foo");
begin
Swap (My_Dictionary (1), Temp);
Free (Temp);
pragma Assert (My_Dictionary (1).all = "foo");
end;

Now everything is fine. The ownership on My_Dictionary (1) is temporarily transferred to the X formal parameter of Swap for the duration of the call, and it is restored at the end. Now the SPARK tool can ensure that My_Dictionary indeed has the full ownership of its content after the call, and the read inside the assertion succeeds. This small example is also verified by the SPARK tool.

I hope this post gave you a taste of what it would be like to program using pointers in SPARK. If you now feel like using them, a preview is available in the community 2019 edition of GNAT+SPARK. Don’t hesitate to come back to us with your findings, either on GitHub or by email.

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GNAT Community 2019 is here! https://blog.adacore.com/gnat-community-2019-is-here Wed, 05 Jun 2019 08:56:00 -0400 Nicolas Setton https://blog.adacore.com/gnat-community-2019-is-here

This release is supported on the same platforms as last year:

• Windows, Linux, and Mac 64-bit native

• RISC-V hosted on Linux

• ARM 32 bits hosted on 64-bit Linux, Mac, and Windows

GNAT Community now includes a number of fixes and enhancements, most notably:

• The installer for Windows and Linux now contains pre-built binary distributions of Libadalang, a very powerful language tooling library for Ada and SPARK.

We hope you enjoy using SPARK and Ada!

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Introduction

C is the dominant language of the embedded world, almost to the point of exclusivity. Due to its age, and its goal of being a “portable assembler”, it deliberately lacks type-safety that languages like Ada provide. The lack of type-safety in C is one of the reasons for the commonness of embedded device exploits. Proposed solutions are partitioning the application into smaller intercommunicating blocks, designed with the principle of least privilege in mind; and rewriting the application in a type-safe language. We believe that both approaches are complementary and want to show you how to combine separation and isolation provided by MultiZone together with iteratively rewriting parts in Ada.

We will take the MultiZone SDK demo and rewrite one of the zones in Ada. The full demo simulates an industrial application with a robotic arm. It runs on the Arty A7-35T board and interfaces with the PC and a robotic arm (OWI-535 Robotic Arm) via an SPI to USB converter. More details are available from the MultiZone Security SDK for Ada manual (https://github.com/hex-five/multizone-ada/blob/master/manual.pdf). We will just be focusing on the porting process here.

MultiZone Security

MultiZone(TM) Security is the first Trusted Execution Environment for RISC-V - it enables development of a simple, policy-based security environment for RISC-V that supports rich operating systems through to bare metal code. It is a culmination of the embedded security best practices developed over the last decade and now applied to RISC-V processors. Instead of splitting into the secure and non-secure domain, MultiZoneTM Security provides policy-based hardware-enforced separation for an unlimited number of security domains, with full control over data, code and peripherals.

MultiZoneTM Security consists of the following components:

• MultiZone(TM) nanoKernel - lightweight, formally verifiable, bare metal kernel providing policy-driven hardware-enforced separation of ram, rom, i/o and interrupts.
• InterZone(TM) Messenger - communications infrastructure to exchange secure messages across zones on a no- shared memory basis.
• MultiZone(TM) Configurator - combines fully linked zone executables with policies and kernel to generate the signed firmware image.
• MultiZone(TM) Signed Boot - 2-stage signed boot loader to verify integrity and authenticity of the firmware image (sha-256 / ECC)

Contrary to traditional solutions, MultiZone(TM) Security requires no additional hardware, dedicated cores or clunky programming models. Open source libraries, third party binaries and legacy code can be configured in minutes to achieve unprecedented levels of safety and security. See https://hex-five.com/ for more details or check out the MultiZone SDK repository on GitHub - https://github.com/hex-five/multizone-sdk.

We port zone 3, the zone controlling the robotic arm, to Ada. The zone communicates with other zones via MultiZone APIs and with the robotic arm by bitbanging GPIO pins.

New Runtime

MultiZone zones differ from a bare metal applications as access to resources is restricted – a zone has only a portion of the RAM and FLASH and can only access some of the peripherals. Looking at our configuration (https://github.com/hex-five/multizone-ada/blob/master/bsp/X300/multizone.cfg), zone3 has the following access privileges:

Zone = 3 #
base = 0x20430000; size =   64K; rwx = rx # FLASH
base = 0x80003000; size =    4K; rwx = rw # RAM
base = 0x0200BFF8; size =   0x8; rwx = r  # RTC
base = 0x10012000; size = 0x100; rwx = rw # GPIO

In the Ada world, this translates to having a separate runtime that we need to create. Luckily, AdaCore has released sources to their existing runtimes on GitHub - https://github.com/adacore/bb-runtimes and they have also included a how-to for creating new runtimes - https://github.com/AdaCore/bb-runtimes/tree/community-2018/doc/porting_runtime_for_cortex_m. Here’s how we create our customized HiFive1 runtime:

./build_rts.py --bsps-only --output=build --prefix=lib/gnat hifive1

This creates sources for building a runtime using a mix of sources from bb-runtimes and from the compiler itself thanks to the --bsps-only flag. Without this switch, we would need the original GNAT repository, which is not publicly available. Notice we don’t use --link, so our runtime sources are a proper copy and can be checked into a new git repo - https://github.com/hex-five/multizone-ada/tree/master/bsp/X300/runtime. Our runtime needs to be compiled and installed before it can be used, and we do that automatically as part of the Makefile for zone3 - https://github.com/hex-five/multizone-ada/blob/master/zone3/Makefile:

BSP_BASE := ../bsp
PLATFORM_DIR := $(BSP_BASE)/$(BOARD)
RUNTIME_DIR := $(PLATFORM_DIR)/runtime GPRBUILD :=$(abspath $(GNAT))/bin/gprbuild GPRINSTALL :=$(abspath $(GNAT))/bin/gprinstall .PHONY: all all:$(GPRBUILD) -p -P $(RUNTIME_DIR)/zfp_hifive1.gpr$(GPRINSTALL) -f -p -P $(RUNTIME_DIR)/zfp_hifive1.gpr --prefix=$(RUNTIME_DIR)
$(AR) cr$(RUNTIME_DIR)/lib/gnat/zfp-hifive1/adalib/libgnat.a
$(RUNTIME_DIR)/hifive1/zfp/obj/*.o$(GPRBUILD) -f -p -P zone3.gpr
\$(OBJCOPY) -O ihex obj/main zone3.hex --gap-fill 0x00

Note that we use a combination of GPRbuild and Make to minimize code differences between the two repositories as much as possible. GPRbuild is limited to zone3 only.

Board support

We use the Ada Drivers Library on GitHub (https://github.com/AdaCore/Ada_Drivers_Library) as a starting point as it provides examples for a variety of boards and architectures. Our target is the X300, itself a modified HiFive1/FE310/FE300. The differences between FE310/FE300 and X300 are detailed on the multizone-fpga GitHub repository (https://github.com/hex-five/multizone-fpga).

We need to change the drivers for our application slightly - we want LD0 for indicating the status of the robotic arm, red blink when disconnected and green when connected:

with FE310.Device; use FE310.Device;
with SiFive.GPIO; use SiFive.GPIO;

package Board.LEDs is

subtype User_LED is GPIO_Point;

Red_LED   : User_LED renames P01;
Green_LED : User_LED renames P02;
Blue_LED  : User_LED renames P03;

procedure Initialize;
-- MUST be called prior to any use of the LEDs

procedure Turn_On (This : in out User_LED) renames SiFive.GPIO.Set;
procedure Turn_Off (This : in out User_LED) renames SiFive.GPIO.Clear;
procedure Toggle (This : in out User_LED) renames SiFive.GPIO.Toggle;

procedure All_LEDs_Off with Inline;
procedure All_LEDs_On with Inline;
end Board.LEDs;

Having the package named Board allows us to modify the target board at compile-time by just providing the folder containing the Board package. This goes in line with what multizone-sdk does.

Code

MultiZone support

MultiZone nanoKernel offers trap and emulate so existing applications can be provided to MultiZone directly unmodified and will work as expected. Access to RAM, FLASH and peripherals needs to be allowed in the configuration file, though. MultiZone does provide functionality for increased performance, better power usage and interzone communication via the API in Libhexfive. Here we create an Ada wrapper around libhexfive32.a, providing the MultiZone specific calls:

package MultiZone is

procedure ECALL_YIELD; -- libhexfive.h:8
pragma Import (C, ECALL_YIELD, "ECALL_YIELD");

procedure ECALL_WFI; -- libhexfive.h:9
pragma Import (C, ECALL_WFI, "ECALL_WFI");
...
end MultiZone;

A typical MultiZone optimized application will yield whenever it doesn’t have anything to do, to save on processing time and power:

with MultiZone; use MultiZone;

procedure Main is
begin
-- Application initialization
loop
-- Application code
ECALL_YIELD;
end loop;
end Main;

If a zone wants to communicate with other zones, such as receiving commands and sending back replies, it needs to use the ECALL_SEND/ECALL_RECV. These send/receive a chunk of 16 bytes and return a status whether the send/receive was successful. The prototypes are a bit special, as they take a void * parameter, which translates to System.Address:

function ECALL_SEND_C (arg1 : int; arg2 : System.Address) return int; -- libhexfive.h:11
pragma Import (C, ECALL_SEND_C, "ECALL_SEND");
function ECALL_RECV_C (arg1 : int; arg2 : System.Address) return int; -- libhexfive.h:12
pragma Import (C, ECALL_RECV_C, "ECALL_RECV");

We wrap these to provide a more Ada idiomatic alternative:

type Word is new Unsigned_32;
type Message is array (0 .. 3) of aliased Word;
pragma Pack (Message);
subtype Zone is int range 1 .. int'Last;

function Ecall_Send (to : Zone; msg : Message) return Boolean;
function Ecall_Recv (from : Zone; msg : out Message) return Boolean;

We hide the System.Address usage and provide a safer subtype for the source/destination zone, since zone cannot be negative or 0.

function Ecall_Send (to : Zone; msg : Message) return Boolean is
begin
return ECALL_SEND_C (to, msg'Address) = 1;
end Ecall_Send;

function Ecall_Recv (from : Zone; msg : out Message) return Boolean is
begin
return ECALL_RECV_C (from, msg'Address) = 1;
end Ecall_Recv;

With all the primitives in place, we can make a simple MultiZone optimized application that can respond to a ping from another zone:

with HAL;       use HAL;
with MultiZone; use MultiZone;

procedure Main is
begin
-- Application initialization
loop
-- Application code
declare
msg : Message;
Status : Boolean := Ecall_Recv (1, msg);
begin
if Status then
if msg(0) = Character'Pos('p') and
msg(1) = Character'Pos('i') and
msg(2) = Character'Pos('n') and
msg(3) = Character'Pos('g') then
Status := Ecall_Send (1, msg);
end if;
end if;
end;

ECALL_YIELD;
end loop;
end Main;

Keeping some legacy (OWI Robot)

We keep the SPI functionality and the Owi Task as C files and just create Ada bindings for them. spi_c.c implements the SPI protocol by bit-banging GPIO pins:

package Spi is

procedure spi_init; -- ./spi.h:8
pragma Import (C, spi_init, "spi_init");

function spi_rw (cmd : System.Address) return UInt32; -- ./spi.h:9
pragma Import (C, spi_rw, "spi_rw");

end Spi;

The OwiTask (owi_task.c) is a state machine containing different robot sequences. The main function is owi_task_run, while others are change the active state. owi_task_run returns the next SPI command to send via SPI for a given moment in time:

package OwiTask is

end OwiTask;

The following Ada code runs the state machine for the OWI robot. Since the target functions are written in C, we see how Ada can interact with legacy C code:

-- OWI sequence run
if usb_state = 16#12670000# then
declare
cmd_bytes : Cmd;
begin
if cmd_word /= -1 then
cmd_bytes(0) := UInt8 (cmd_word and 16#FF#);
cmd_bytes(1) := UInt8 (Shift_Right (cmd_word,  8) and 16#FF#);
cmd_bytes(2) := UInt8 (Shift_Right (cmd_word, 16) and 16#FF#);
ping_timer := CLINT.Machine_Time + PING_TIME;
end if;
end;
end if;

We now port over the owi sequence selection:

declare
Status : Boolean := Ecall_Recv (1, msg);
begin
if Status then
-- OWI sequence select
if usb_state = 16#12670000# then
case msg(0) is
when others => null;
end case;
end if;
...

We leave it as an exercise for the reader to port the OWI sequence handler code to Ada.

We want the LED color to change as a result of the robot connected or disconnected events. Each of the LED colors is a separate pin on the board:

Red_LED   : User_LED renames P01;
Green_LED : User_LED renames P02;
Blue_LED  : User_LED renames P03;

One way of handling this is to create an Access Type that can store the currently selected LED color:

procedure Main is
type User_LED_Select is access all User_LED;
...
LED : User_LED_Select := Red_LED'Access;
...
begin
...
-- Detect USB state every 1sec
if CLINT.Machine_Time > ping_timer then
ping_timer := CLINT.Machine_Time + PING_TIME;
end if;

-- Update USB state
declare
Status : int;
begin
if rx_data /= usb_state then
usb_state := rx_data;

if rx_data = UInt32'(16#12670000#) then
LED := Green_LED'Access;
else
LED := Red_LED'Access;
end if;
end if;
end;

The actual blinking is then just a matter of dereferencing the Access Type and calling the right function:

-- LED blink
if CLINT.Machine_Time > led_timer then
if GPIO.Set (Red_LED) or GPIO.Set (Green_LED) or GPIO.Set (Blue_LED) then
All_LEDs_Off;
led_timer := CLINT.Machine_Time + LED_OFF_TIME;
else
All_LEDs_Off;
Turn_On (LED.all);
led_timer := CLINT.Machine_Time + LED_ON_TIME;
end if;
end if;

Closing

If you would like to know more about MultiZone and Ada, please reach out to Hex-Five Security - https://hex-five.com/contact-2/. I will also be attending the RISC-V Workshop in Zurich - https://tmt.knect365.com/risc-v-workshop-zurich/ if you would like to grab a coffee and discuss MultiZone and Ada on RISC-V.

]]>

Using SPARK to prove absence of run time errors on a hobby project.

The Danish Technical University has a yearly RoboCup where autonomous vehicles solve a number of challenges. Each solved challenge gives points and the team with most points wins. The track is available for testing two weeks before the qualification round.

The idea behind and creation of RoadRunner for DTU RoboCup 2019.

RoadRunner is a 3D printed robot with wheel suspension, based on the BeagleBone Blue ARM-based board and the Pixy 1 camera with custom firmware enabling real-time line detection. Code is written in Ada and formally proved correct with SPARK at Silver level. SPARK Silver level proves that the code will execute without run-time errors like overflows, race-conditions and deadlocks. During development SPARK prevented numerous hard-to-debug errors to reach the robot. The time-boxed testing period of DTU RoboCup made early error detection even more valuable. In our opinion the estimated time saved on debugging more than outweighs the time we spent on SPARK. The robot may still fail, but it will not be due to run-time errors in our code...

Ada was the obvious choice. We both have years of experience with Ada. The BeagleBone Blue runs Debian Linux and we use the default GNAT FSF compiler version from Debian Testing. It is easy to cross compile from a Debian laptop and it is not too old compared to GNAT Community Edition. The BeagleBone Blue runs Debian Stable.

We decided initially not to use SPARK. According to some internet articles, SPARK had problems proving properties about floating points, and the Ada subset for tasking seemed to be too restricted. Support for floating points has improved since then, and tasking was extended to the Jorvik (extended Ravenscar) profile.

A race condition in the code changed the SPARK decision. A lot of valuable time was spent chasing it. The GNAT runtime on Debian armhf has no traceback info, so it was difficult to find what caused a Storage_Error. SPARK flow analysis detects race conditions, so that would prevent that kind of issues in the future.

It took a lot of effort before SPARK would be able to do flow analysis without errors. SPARK rejects code with exception handlers, fixed-point to floating-point conversions, the standard Ada.Containers and task entries. Hopefully, future versions of SPARK will be able to analyze the above or at least give a warning and continue analysis, even if it contains unsupported Ada features. When the code base is in SPARK, it is quite easy to add new code in SPARK.

Experience with SPARK.

I had to start from scratch, with no previous experience with SPARK, formal methods or safety critical software. I read the online documentation to find examples on how to prove different types of properties.

SPARK is a tough code reviewer. It detects bad design decisions and forces a rewrite. Bad design is impossible to test and prove and SPARK does not miss any of it. Very valuable for small projects without any code review.

SPARK code is easier to read. Defensive code is moved from exception handlers and if statements to preconditions. The resulting code has less branching and is easier to understand and test.

SPARK code is also more difficult to read. Loop invariants and assertions to prove loops can outnumber the Ada statements in the loop.

We used saturation as a shortcut to prove absence of overflows, in locations where some values could not be proved to stay in bounds. But this is not an ideal solution: if it does saturate, then the result is not correct.

In our experience, proving absence of run time errors also detects real bugs, not only division by zero and overflows.

Two examples.

SPARK could not prove a resulting value in range. A scaling factor had been divided with instead of multiplied with.

SPARK could not prove array index in range. X and Y on camera image calculation had been swapped.

Stability.

If code is in SPARK, then all changes must be analyzed by SPARK. Almost all coding error gives a run time exception. That happened several times on the track during testing. But when tested with SPARK, it always found the exact line with the bug. After learning the lesson we found that code analyzed with SPARK always worked the first time.

Having a stable build environment with SPARK and automated tests, made it possible to make successful last-minute changes to the software just before the final run in the video link.

Here is the presentation video for our robot:

and the video of the final winning challenge:

]]>
Using SPARK to prove 255-bit Integer Arithmetic from Curve25519 https://blog.adacore.com/using-spark-to-prove-255-bit-integer-arithmetic-from-curve25519 Tue, 30 Apr 2019 15:00:00 -0400 Joffrey Huguet https://blog.adacore.com/using-spark-to-prove-255-bit-integer-arithmetic-from-curve25519

In 2014, Adam Langley, a well-known cryptographer from Google, wrote a post on his personal blog, in which he tried to prove functions from curve25519-donna, one of his projects, using various verification tools: SPARK, Frama-C, Isabelle... He describes this attempt as "disappointing", because he could not manage to prove "simple" things, like absence of runtime errors. I will show in this blogpost that today, it is possible to prove what he wanted to prove, and even more.

Algorithms in elliptic-curve cryptography compute numbers which are too large to fit into registers of typical CPUs. In this case, curve25519 uses integers ranging from 0 to 2255-19. They can't be represented as native 32- or 64-bit integers. In the original implementations from Daniel J. Bernstein and curve25519-donna, these integers are represented by arrays of 10 smaller, 32-bit integers. Each of these integers is called a "limb". Limbs have alternatively 26-bit and 25-bit length. This forms a little-endian representation of the bigger integer, with low-weight limbs first. In section 4 of this paper, Bernstein explains both the motivation and the details of his implementation. The formula to convert an array A of 32-bit integers into the big integer it represents is: $\sum_{i=0}^{9} A (i) \times 2^{\lceil 25.5i \rceil}$ where the half square brackets represent the ceiling function. I won't be focusing on implementation details here, but you can read about them in the paper above. To me, the most interesting part about this project is its proof.

First steps

Types

The first step was to define the different types used in the implementation and proof.

I'll use two index types. Index_Type, which is range 0 .. 9, is used to index the arrays that represent the 255-bit integers. Product_Index_Type is range 0 .. 18, used to index the output array of the Multiply function.

Integer_Curve25519 represents the arrays that will be used in the implementation, so 64-bit integer arrays with variable bounds in Product_Index_Type.  Product_Integer is the array type of the Multiply result. It is a 64-bit integer array with Product_Index_Type range. Integer_255 is the type of the arrays that represent 255-bit integers, so arrays of 32-bit integers with Index_Type range.

Big integers library

Proving Add and Multiply in SPARK requires reasoning about the big integers represented by the arrays. In SPARK, we don't have mathematical integers (i.e. integers with infinite precision), only bounded integers, which are not sufficient for our use case. If I had tried to prove correctness of Add and Multiply in SPARK with bounded integers, the tool would have triggered overflow checks every time an array is converted to integer.

To overcome this problem, I wrote the specification of a minimal big integers library, that defines the type, relational operators and basic operations. It is a specification only, so it is not executable.

SPARK is based on why, which has support for mathematical integers, and there is a way to link a why3 definition to a SPARK object directly. This is called external axiomatization; you can find more details about how to do it here.

Using this feature, I could easily provide a big integers library with basic functions like "+", "*" or To_Big_Integer (X : Integer). As previously mentioned, this library is usable for proof, but not executable (subprograms are marked as Import).  To avoid issues during binding, I used a feature of SPARK named Ghost code. It takes the form of an aspect: "Ghost" indicates to the compiler that this code cannot affect the values computed by ordinary code, thus can be safely erased. It's very useful for us, since this aspect can be used to write non-executable functions which are only called in annotations.

Conversion function

One of the most important functions used in the proof is a function that converts the array of 10 integers to the big integer it represents, so when X is an Integer_255 then +X is its corresponding big integer.

function "+" (X : Integer_Curve25519) return Big_Integer
renames To_Big_Integer;

After this, I can define To_Big_Integer recursively, with the aid of an auxiliary function, Partial_Conversion.

function Partial_Conversion
(X : Integer_Curve25519;
L : Product_Index_Type)
return Big_Integer
is
(if L = 0
then (+X(0)) * Conversion_Array (0)
else
Partial_Conversion (X, L - 1) + (+X (L)) * Conversion_Array (L))
with
Ghost,
Pre => L in X'Range;

function To_Big_Integer (X : Integer_Curve25519)
return Big_Integer
is
(Partial_Conversion (X, X'Last))
with
Ghost,
Pre => X'Length > 0;

With these functions defined, it was much easier to write specifications of Add and Multiply. All_In_Range is used in preconditions to bound the parameters of Add and Multiply in order to avoid overflow.

function All_In_Range
(X, Y     : Integer_255;
Min, Max : Long_Long_Integer)
return     Boolean
is
(for all J in X'Range =>
X (J) in Min .. Max
and then Y (J) in Min .. Max);

function Add (X, Y : Integer_255) return Integer_255 with
Pre  => All_In_Range (X, Y, -2**30 + 1, 2**30 - 1),
Post => +Add'Result = (+X) + (+Y);

function Multiply (X, Y : Integer_255) return Product_Integer with
Pre  => All_In_Range (X, Y, -2**27 + 1, 2**27 - 1),
Post => +Multiply'Result = (+X) * (+Y);

Proof

Proof of absence of runtime errors was relatively simple: no annotation was added to Add, and a few, very classical loop invariants for Multiply. Multiply is the function where Adam Langley stopped because he couldn't prove absence of overflow. I will focus on proof of functional correctness, which was a much more difficult step. In SPARK, Adam Langley couldn't prove it for Add because of overflow checks triggered by the tool when converting the arrays to the big integers they represent. This is the specific part where the big integers library is useful: it is possible to manipulate big integers in proof without overflow checks.

The method I used to prove both functions is the following:

1.   Create a function that allows tracking the content of the returned array.
2.   Actually track the content of the returned array through loop invariant(s) that ensure equality with the function in point 1.
3.  Prove the equivalence between the loop invariant(s) at the end of the loop and the postcondition in a ghost procedure.
4.  Call the procedure right before return; line.

I will illustrate this example with Add. The final implementation of Add is the following.

function Add (X, Y : Integer_255) return Integer_255 is
Sum : Integer_255 := (others => 0);
begin
for J in Sum'Range loop
Sum (J) := X (J) + Y (J);

pragma Loop_Invariant (for all K in 0 .. J =>
Sum (K) = X (K) + Y (K));
end loop;
return Sum;
end Add;

Points 1 and 2 are ensured by the loop invariant. No new function was created since the content of the array is simple (just an addition). At the end of the loop, the information we have is: "for all K in 0 .. 9 => Sum (K) = X (K) + Y (K)" which is an information on the whole array. Points 3 and 4 are ensured by Prove_Add. Specification of Prove_Add is:

procedure Prove_Add (X, Y, Sum : Integer_255) with
Ghost,
Pre  => (for all J in Sum'Range => Sum (J) = X (J) + Y (J)),
Post => To_Big_Integer (Sum) = To_Big_Integer (X) + To_Big_Integer (Y);

This is what we call a lemma in SPARK. Lemmas are ghost procedures, where preconditions are the hypotheses, and postconditions are the conclusions. Some lemmas are proved automatically, while others require solvers to be guided. You can do this by adding a non-null body to the procedure, and guide them differently, depending on the conclusion you want to prove.

I will talk about the body of Prove_Add in a few paragraphs.

Multiply

Multiply was much harder than Add had been. But I had the most fun proving it. I refer to the implementation in curve25519-donna as the "inlined" one, because it is fully explicit and doesn't go through loops. This causes a problem in the first point of my method: it is not possible to track what the routine does, except by adding assertions at every line of code. To me, it's not a very interesting point of view, proof-wise. So I decided to change the implementation, to make it easier to understand. The most natural approach for multiplication, in my opinion, is to use distributivity of product over addition. The resulting implementation would be similar to TweetNaCl's implementation of curve25519 (see M function):

function Multiply (X, Y : Integer_255) return Product_Integer is
Product : Product_Integer := (others => 0);
begin
for J in Index_Type loop
for K in Index_Type loop
Product (J + K) :=
Product (J + K)
+ X (J) * Y (K) * (if J mod 2 = 1 and then K mod 2 = 1 then 2 else 1);
end loop;
end loop;
return Product;
end Multiply;

Inside the first loop, J is fixed and we iterate over all values of K, so on all the values of Y. It is repeated over the full range of J, so the entire content of X.  With this implementation, it is possible to track the content of the array in loop invariants through an auxiliary function, Partial_Product, that will track the value of a certain index in the array at each iteration. We add the following loop invariant at the end of the loop:

pragma Loop_Invariant (for all L in 0 .. K =>
Product (J + L) = Partial_Product (X, Y, J, L));

The function Partial_Product is defined recursively, because it is a sum of several factors.

function Partial_Product
(X, Y : Integer_255;
J, K : Index_Type)
return Long_Long_Integer
is
(if K = 9 or else J = 0
then (if J mod 2 = 1 and then K mod 2 = 1 then 2 else 1) * X (J) * Y (K)
else Partial_Product (X, Y, J - 1, K + 1)
+ (if J mod 2 = 1 and then K mod 2 = 1 then 2 else 1) * X (J) * Y (K));

As seen in the loop invariant above, the function returns Product (J + K), which is modified for all indexes L and M when J + K = L + M. The recursive call will be on the pair J + 1, K - 1, or J - 1, K + 1. Given how the function is designed in the implementation, the choice J - 1, K + 1 is preferable, because it follows the evolution of Product (J + K). The base case is when J = 0 or K = 9.

Problems and techniques

Defining functions to track content was rather easy to do for both functions. Proving that the content is actually equal to the function was a bit more difficult for the Multiply function, but not as difficult as proving the equivalence between this and the postcondition. With these two functions, we challenge the provers in many ways:

• We use recursive functions in contracts, which means provers have to reason inductively and they struggle with this. That's why we need to /guide/ them in order to prove properties that involve recursive functions.
• The context size is quite big, especially for Multiply. Context represent all variables, axioms, theories that solvers are given to prove a check. The context grows with the size of the code, and sometimes it is too big. When this happens, solvers may be lost and not be able to prove the check. If the context is reduced, it will be easier for solvers, and they may be able to prove the previously unproved checks.
• Multiply contracts need proof related to nonlinear integers arithmetics theory, which provers have a lot of problems with.  This problem is specific to Multiply, but it explains why this function took some time to prove. Solvers need to be guided quite a bit in order to prove certain properties.

What follows is a collection of techniques I found interesting and might be useful when trying to prove problems, even very different from this one.

Axiom instantiation

When I tried to prove loop invariants in Multiply to track the value of Product (J + K), solvers were unable to use the definition of Partial_Product. Even asserting the exact same definition failed to be proven. This is mainly due to context size: solvers are not able to find the axiom in the search space, and the proof fails. The workaround I found is to create a Ghost procedure which has the definition as its postcondition, and a null body, like this:

procedure Partial_Product_Def
(X, Y : Integer_255;
J, K : Index_Type)
with
Pre  => All_In_Range (X, Y, Min_Multiply, Max_Multiply),
Post =>
(if K = Index_Type'Min (9, J + K)
then Partial_Product (X, Y, J, K)
= (if J mod 2 = 1 and then K mod 2 = 1 then 2 else 1)
* X (J) * Y (K)
else Partial_Product (X, Y, J, K)
= Partial_Product (X, Y, J - 1, K + 1)
+ X (J) * Y (K)
* (if J mod 2 = 1 and then K mod 2 = 1 then 2 else 1));
procedure Partial_Product_Def
(X, Y : Integer_255;
J, K : Index_Type)
is null;

In this case, the context size is reduced considerably, and provers are able to prove it without any body. During the proof process, I had to create other recursive functions that needed this *_Def procedure in order to use their definition. It has to be instantiated manually, but it is a way to "remind" solvers this axiom. A simple but very useful technique.

Manual induction

When trying to prove Prove_Add, I encountered one of the cases where the provers have to reason inductively: a finite sum is defined recursively. To help them prove the postcondition, the conversion is computed incrementally in a loop, and the loop invariant tracks the evolution.

procedure Prove_Add (X, Y, Sum : Integer_255) with
Ghost,
Pre  => (for all J in Sum'Range => Sum (J) = X (J) + Y (J)),
Post => To_Big_Integer (Sum) = To_Big_Integer (X) + To_Big_Integer (Y);
--  Just to remember

procedure Prove_Add (X, Y, Sum : Integer_255) with
Ghost
is
X_255, Y_255, Sum_255 : Big_Integer := Zero;
begin
for J in Sum'Range loop

X_255 := X_255 + (+X (J)) * Conversion_Array (J);
Y_255 := Y_255 + (+Y (J)) * Conversion_Array (J);
Sum_255 := Sum_255 + (+Sum (J)) * Conversion_Array (J);

pragma Loop_Invariant (X_255 = Partial_Conversion (X, J));
pragma Loop_Invariant (Y_255 = Partial_Conversion (Y, J));
pragma Loop_Invariant (Sum_255 = Partial_Conversion (Sum, J));
pragma Loop_Invariant (Partial_Conversion (Sum, J) =
Partial_Conversion (X, J) +
Partial_Conversion (Y, J));
end loop;
end Prove_Add;

What makes this proof inductive is the treatment of Loop_Invariants by SPARK: it will first try to prove the first iteration, as an initialization, and then it will try to prove iteration N knowing that the property is true for N - 1.

Here, the final loop invariant is what we want to prove, because when J = 9, it is equivalent to the postcondition. The other loop invariants and variables are used to compute incrementally the values of partial conversions and facilitate proof.

Fortunately, in this case, provers do not need more guidance to prove the postcondition.  Prove_Multiply also follows this proof scheme, but is much more difficult. You can access its proof following this link.

Another case of inductive reasoning in my project is a lemma, whose proof present a very useful technique when proving algorithms using recursive ghost functions in contracts.

procedure Equal_To_Conversion
(A, B : Integer_Curve25519;
L    : Product_Index_Type)
with
Pre  =>
A'Length > 0
and then A'First = 0
and then B'First = 0
and then B'Last <= A'Last
and then L in B'Range
and then (for all J in 0 .. L => A (J) = B (J)),
Post => Partial_Conversion (A, L) = Partial_Conversion (B, L);

It states that given two Integer_Curve25519, A and B, and a Product_Index_Type, L, if A (0 .. L) = B (0 .. L), then Partial_Conversion (A, L) = Partial_Conversion (B, L). The proof for us is evident because it can be proved by induction over L, but we have to help SPARK a bit, even though it's usually simple.

procedure Equal_To_Conversion
(A, B : Integer_Curve25519;
L    : Product_Index_Type)
is
begin
if L = 0 then
return;                         --  Initialization of lemma
end if;
Equal_To_Conversion (A, B, L - 1); --  Calling lemma for L - 1
end Equal_To_Conversion;

The body of a procedure proved by induction actually looks like induction: first thing to add is initialization, with an if statement ending with a return;. Then, the following code is the general case. We call the same lemma for L - 1, and we add assertions to prove postcondition if necessary. In this case, calling the lemma at L - 1 was sufficient to prove the postcondition.

Guide with assertions

Even if the solvers know the relation between Conversion_Array (J + K) and Conversion_Array (J) * Conversion_Array (K), it is hard for them to prove properties requiring non-linear arithmetic reasoning. The following procedure is a nice example:

procedure Split_Product
(Old_Product, Old_X, Product_Conversion : Big_Integer;
X, Y                                   : Integer_255;
J, K                                   : Index_Type)
with
Ghost,
Pre  =>
Old_Product
= Old_X * (+Y)
+ (+X (J))
* (if K = 0
then Zero
else Conversion_Array (J) * Partial_Conversion (Y, K - 1))
and then
Product_Conversion
= Old_Product
+ (+X (J)) * (+Y (K))
* (+(if J mod 2 = 1 and then K mod 2 = 1 then 2 else 1))
* Conversion_Array (J + K),
Post =>
Product_Conversion
= Old_X * (+Y)
+ (+X (J)) * Conversion_Array (J)
* Partial_Conversion (Y, K);

The preconditions imply the postcondition by arithmetic reasoning. With a null body, the procedure is not proved. We can guide provers through assertions, by splitting the proof into two different cases:

procedure Split_Product
(Old_Product, Old_X, Product_Conversion : Big_Integer;
X, Y                                   : Integer_255;
J, K                                   : Index_Type)
is
begin
if J mod 2 = 1 and then K mod 2 = 1 then
pragma Assert (Product_Conversion
= Old_Product
+ (+X (J)) * (+Y (K))
* Conversion_Array (J + K) * (+2));
pragma Assert ((+2) * Conversion_Array (J + K)
= Conversion_Array (J) * Conversion_Array (K));
--  Case where Conversion_Array (J + K) * 2
--             = Conversion_Array (J) * Conversion_Array (K).
else
pragma Assert (Conversion_Array (J + K)
= Conversion_Array (J) * Conversion_Array (K));
--  Other case
end if;
if K > 0 then
pragma Assert (Partial_Conversion (Y, K)
= Partial_Conversion (Y, K - 1)
+ (+Y (K)) * Conversion_Array (K));
--  Definition of Partial_Conversion, needed for proof
end if;
end Split_Product;

What is interesting in this technique is that it shows that to prove something with auto-active proof, you need to understand it yourself first. I proved this lemma by hand on paper, which was not difficult, and then I rewrote my manual proof through assertions. If my manual proof is easier when I split two different cases, so it should be for provers. It allows them to choose the first /good/ steps to the proof.

Split into subprograms

Another thing I have noticed, is that provers were really quickly overwhelmed by context size in my project. I had a lot of recursive functions in my contracts, but also quantifiers... I did not hesitate to split proofs into Ghost procedures in order to reduce context size, but also wrap some expressions into functions.

I have 7 subprograms that enable me to prove Prove_Multiply, which is not a lot in my opinion. It also increases readability, which is important if you want other people to read your code.

There is also another method I want to share to overcome this problem. When there is only one assertion to prove, but it requires a lot of guidance, it is possible to put all assertions, including the last one in a begin end block, and the last one has a Assert_And_Cut pragma, just like this:

code ...
--  It is possible to add the keyword declare to declare variables
--  before the block, if it helps for proof.
begin
pragma Assert (one assert needed to prove the last assertion);
pragma Assert (another one);
pragma Assert_And_Cut (Assertion to prove and remember);
end;

Assert_And_Cut asks the provers to prove the property inside them, but also will remove all context that has been created in the begin end; block. It will keep just this assertion, and might help to reduce context.

Sadly, this workaround didn't work for my project, because context was already too big to prove the first assertions. But adding a lot of subprograms also has drawbacks, e.g. you have to write preconditions to the procedures, and this may be more difficult than just writing the block with an Assert_And_Cut. Surely, both are useful in various projects, and I think it's nice to have these methods in mind.

Takeaways

For statistics only: Add has a 8-line implementation for around 25 lines of Ghost code, to prove it. Multiply has 12 lines of implementation for more than 500 lines of proof. And it's done with auto-active proof, which means that verification was all done automatically! In comparison, verification of TweetNaCl's curve25519 has more than 700 lines of Coq in order to prove Multiply for an implementation of 5 lines, but they have a carry operation to prove on top of the product.

I would say this challenge is at an intermediary level, because it is not difficult to understand the proof of Multiply when you write it down. But it presents several techniques that can apply to various other problems, and this is the main interest of the project to me.

As done in the blogpost by Adam Langley, I tried to prove my project with alt-ergo only, because it was the only prover available for SPARK back in 2014. Even today, alt-ergo alone is unable to prove all code. However, it doesn't make alt-ergo a bad prover, in fact, none of the provers available in SPARK (CVC4, Z3 and Alt-Ergo) are able to prove the entire project on their own. I think it shows that having multiple provers available increases greatly chances of code to be proved.

At the beginning, working on this project was just a way to use my big integers library in proof. But in the end, I believe it is an interesting take on the challenge of verifying elliptic curves functions, especially when projects like this appear for example with fiat-crypto or verification of TweetNaCl's curve25519, and I had a lot of fun experimenting with SPARK to prove properties that usually are badly handled by provers. You can access full proof of my project in this repository.

]]>

This course is geared to software professionals looking for a practical introduction to the Ada language with a focus on embedded systems, including real-time features as well as critical features introduced in Ada 2012. By attending this course you will understand and know how to use Ada for both sequential and concurrent applications, through a combination of live lectures from AdaCore's expert instructors and hands-on workshops using AdaCore's latest GNAT technology. AdaCore will provide an Ada 2012 tool-chain and ARM-based target boards for embedded workshops. No previous experience with Ada is required.

The course will be conducted in English.

Prerequisite: Knowledge of a programming language (Ada 83, C, C++, Java…)

Each participant should come with a computer running Windows.

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A question that our users sometimes ask us is "do you use CodePeer at AdaCore and if so, how?". The answer is yes! and this blog post will hopefully give you some insights into how we are doing it for our own needs.

First, I should note that at AdaCore we are in a special situation since we are both developing lots of Ada code, and at the same time we are also developing and evolving CodePeer itself. One of the consequences is that using CodePeer on our own code has actually two purposes: one is to improve the quality of our code by using an advanced static analyzer, and the other is to eat our own dog food and improve the analyzer itself by finding limitations, sub-optimal usage, etc...

In the past few years, we've gradually added automated runs of CodePeer on many of our Ada code base, in addition to the systematic use of light static analysis performed by the compiler: GNAT already provides many useful built-in checks such as all the static Ada legality rules, as well as many clever warnings fine tuned over the past 25 years and available via the -gnatwa compiler switch and coupled with -gnatwe that transforms warnings into errors, not mentioning also the built-in style checks available via the -gnaty switch.

GNAT

For GNAT sources (Ada front-end and runtime source code) given that this is a large and complex piece of Ada code (compilers do come with a high level of algorithmic complexity and code recursion!) and that we wanted to favor rapid and regular feedback, we've settled on running CodePeer at level 1 with some extra fine tuning: we've found that some categories of messages didn't generate any extra value for the kind of code found in GNAT, so we've disabled these categories via the --be-messages switch. We've also excluded some files from analysis which were taking a long time compared to other files, for little benefits, via the Excluded_Source_Files project attribute as explained in the Partial Analysis section of the User's Guide.

Given that the default build mechanism of GNAT is based on Makefiles, we've written a separate project file for performing the CodePeer analysis, coupled with a separate simple Makefile that performs the necessary setup phase (automatic generation of some source files in the case of GNAT sources).

After this fine tuning, CodePeer generated a few dozen useful messages that we analyzed and allowed us to fix potential issues, as well as improve the general quality of the code. You'll find a few examples below. The most useful findings on the GNAT code base are related to redundant or dead code and potential access to uninitialized variables when complex conditions are involved, which is often the case in a compiler! CodePeer also detected useful cases of code cleanups in particular related to wrong parameter modes: in out parameters that should be out or inout parameters that should be in out.

Once we've addressed these findings by improving the source code, we ended up in a nice situation where the CodePeer run was completely clean: no new messages found. So we've decided on a strict policy similar to our no warning policy: no CodePeer messages should be left alone moving forward. To ensure that, we've put in place a continuous run that triggers after each commit in the GNAT repository and reports back its findings within half an hour.

One of the most useful categories of potential issues found by CodePeer in our case is related to not always initializing a variable before using it (what CodePeer calls validity checks). For example in gnatchop.adb we had the following code:

Offset_FD : File_Descriptor;
[...]

exception
when Failure | Types.Terminate_Program =>
Close (Offset_FD);
[...]

CodePeer detected that at line 6 Offset_FD may not have been initialized (precisely because an exception may have been raised before assigning Offset_FD. We fixed this potential issue by explicitly assigning a default value and testing for it:

Offset_FD : File_Descriptor := Invalid_FD;
[...]

exception
when Failure | Types.Terminate_Program =>
if Offset_FD /= Invalid_FD then
Close (Offset_FD);
end if;
[...]

CodePeer also helped detect suspicious or dead code that should not have been there in the first place. Here is an example of code smell that CodePeer detected in the file get_scos.adb:

procedure Skip_EOL is
C : Character;
begin
loop
Skipc;
C := Nextc;
exit when C /= LF and then C /= CR;

if C = ' ' then
Skip_Spaces;
C := Nextc;
exit when C /= LF and then C /= CR;
end if;
end loop;
end Skip_EOL;

in the code above, CodePeer complained that the test at line 9 is always false because C is either LF or CR. And indeed, if you look closely at line 7, when the code goes past this line, then C will always be either CR or LF, and therefore cannot be a space character. The code was simplified into:

procedure Skip_EOL is
C : Character;
begin
loop
Skipc;
C := Nextc;
exit when C /= LF and then C /= CR;
end loop;
end Skip_EOL;

GPS

Analysis of GPS sources with CodePeer is used at AdaCore both for improving the code quality and also to test our integration with the SonarQube tool via our GNATdashboard integration.

For GPS, we are using CodePeer in two modes: one mode where developers can manually run CodePeer on their local set up at level 0. This mode runs in about 3 to 5 minutes to analyze all the GPS sources and runs the first level of checks provided by CodePeer, which is a very useful complement to the compiler warnings and style checks and allows to stay 100% clean of messages after an initial set of code cleanups.

In addition, an automated run is performed nightly on a server using level 1, further tuned in a similar way to what we did for GNAT. Here we have some remaining messages under analysis and we use SonarQube to track and analyze these messages.

Here is an example of code that looked suspicious in GPS sources:

View : constant Registers_View := Get_Or_Create_View (...);
begin
View.Locked := True;

if View /= null then
[...]
end if;

CodePeer complained at line 5 that the test is always True since View cannot be null at this point. Why? Because at line 3 we are already dereferencing View, so CodePeer knows that after this point either an exception was raised or, if not, View cannot be null anymore.

In this case, we've replaced the test by an explicit assertion since it appears that Get_Or_Create_View can never return null:

View : constant Registers_View := Get_Or_Create_View (...);
begin
pragma Assert (View /= null);
View.Locked := True;
[...]

Run Time Certification

As part of a certification project of one of our embedded runtimes for a bare metal target, we ran CodePeer at its highest level (4) in order to detect all potentially cases of a number of vulnerabilities and in particular: validity checks, divide by zero, overflow checks, as well as confirming that the runtime did not contain dead code or unused assignments. CodePeer was run manually and then all messages produced were reviewed and justified, as part of our certification work.

GNAT LLVM

As part of the GNAT LLVM project (more details on this project if you're curious in a future post!) in its early stages, we ran CodePeer manually on all the GNAT LLVM sources - excluding the sources common to GNAT, already analyzed separately - initially at level 1 and then at level 3, and we concentrated on analyzing all (and only) the potential uninitialized variables (validity checks). In this case we use the same project file used to build GNAT LLVM itself and added the CodePeer settings which basically looked like:

package CodePeer is
for Switches use ("-level", "3", "--be-messages=validity_check", "--no-lal-checkers");
for Excluded_Source_Dirs use ("gnat_src", "obj");
end CodePeer;

which allowed us to perform a number of code cleanups and review more closely the code pointed by CodePeer.

CodePeer

Last but not least, we also run CodePeer on its own code base! For CodePeer given that this is another large and complex piece of Ada code and we wanted to favor rapid and regular feedback, we've settled on a setting similar to GNAT: level 1 with some extra fine tuning via --be-messages. We also added a few pragma Annotate to both justify some messages - pragma Annotate (CodePeer, False_Positive) - as well as skipping analysis of some subprograms or files where CodePeer was taking too long to analyze, for little benefits (via either the Excluded_Source_Files project attribute or pragma Annotate (CodePeer, Skip_Analysis)).

A CodePeer run is triggered after each change in the repository in a continuous builder and made available to the team within 30 minutes. We've found that in this case the most interesting messages where: validity checks on local variables and out parameters, test always true/false, duplicated code and potential wrong parameter mode.

We also run CodePeer on other code bases in a similar fashion, such as analysis of the SPARK tool sources.

As part of integrating CodePeer in our daily work, we also took the opportunity to improve the documentation and describe many possible workflows corresponding to the various needs of teams wanting to analyze Ada code, with an explanation on how to put these various scenarios in place, check Chapter 5 of the User's Guide if you're curious.

What about you? Do not hesitate to tell us how you are using CodePeer and what are the most useful benefits you are getting by commenting below.

]]>
Ten Years of Using SPARK to Build CubeSat Nano Satellites With Students https://blog.adacore.com/ten-years-of-using-spark-to-build-cubesat-nano-satellites-with-students Fri, 01 Mar 2019 14:31:01 -0500 Peter Chapin https://blog.adacore.com/ten-years-of-using-spark-to-build-cubesat-nano-satellites-with-students

My colleague, Carl Brandon, and I have been running the CubeSat Laboratory at Vermont Technical College (VTC) for over ten years. During that time we have worked with nearly two dozen students on building and programming CubeSat nano satellites. CubeSats are small (usually 10cm cube), easily launched spacecraft that can be outfitted with a variety of cameras, sensing instruments, and communications equipment. Many CubeSats are built by university groups like ours using students at various skill levels in the design and production process.

Students working in the CubeSat Laboratory at VTC have been drawn from various disciplines including computer engineering, electro-mechanical engineering, electrical engineering, and software engineering. VTC offers a masters degree in software engineering, and two of our MSSE students have completed masters projects related to CubeSat flight software. In fact, our current focus is on building a general purpose flight software framework called CubedOS.

Like all spacecraft CubeSats are difficult to service once they are launched. Because of the limited financial resources available to university groups, and because of on-board resource constraints, CubeSats typically don't support after-launch uploading of software updates. This means the software must be fully functional and fault-free, with no possibility of being updated, at the time of launch

Many university CubeSat missions have failed due to software errors. This is not surprising considering that most flight software is written in C, a language that is difficult to use correctly. To mitigate this problem we use the SPARK dialect of Ada in all of our software work. Using the SPARK tools we work toward proving the software free of runtime error, meaning that no runtime exceptions will occur. However, in general we have not attempted to prove functional correctness properties, relying instead on conventional testing for that level of verification.

Although we do have some graduate students working in the CubeSat Laboratory, most of our students are third and fourth year undergraduates. The standard curriculum for VTC's software engineering program includes primarily the Java and C languages. Although I have taught Ada in the classroom, it has only been in specialized classes that a limited number of students take. As a result most of the students who come to the CubeSat Laboratory have no prior knowledge of SPARK or Ada, although they have usually taken several programming courses in other languages.

Normally I arrange to give new students an intensive SPARK training course that spans three or four days, with time for them to do some general exercises. After that I assign the students a relatively simple introductory task involving our codebase so they can get used to the syntax and semantics of SPARK without the distraction of complex programming problems. Our experience has been that good undergraduate students with a reasonable programming background can become usefully productive with SPARK in as little as two weeks. Of course, it takes longer for them to gain the skills and experience to tackle the more difficult problems, but the concern commonly expressed by some that the lack of SPARK skills in a programming team is a barrier to the adoption of SPARK is not validated by our experience

Students are, of course, novice programmers almost by definition. Many of our students are in the process of learning basic software engineering principles such as the importance of requirements, code review, testing, version control, continuous integration, and many other things. Seeing these ideas in the context of our CubeSat work gives them an important measure of realism that can be lacking in traditional courses.

However, because of their general inexperience, and because of the high student turnover rate that is natural in an educational setting, our development process is often far from ideal. Here SPARK has been extremely valuable to us. What we lack in rigor of the development process we make up for in the rigor of the SPARK language and tools. For example, if a student refactors some code, perhaps without adequately communicating with the authors of the adjoining code, the SPARK tools will often find inconsistencies that would otherwise go unnoticed. This has resulted in a much more disciplined progression of the code than one would expect based on the team's overall culture. Moreover the discipline imposed by SPARK and its tools can serve to educate the students about the kinds of issues that can go wrong by following an overly informal approach to development.

For example, one component of CubedOS is a module that sends “tick” messages to other modules on request. These messages are largely intended to trigger slow, non-timing critical housekeeping tasks in the other modules. For flexibility the module supports both one-shot tick messages that occur only once and periodic tick messages. Many “series” of messages can be active at once. The module can send periodic tick messages with different periods to several receivers with additional one-shot tick messages waiting to fire as well.

The module has been reworked many times. At one point it was decided that tick messages should contain a counter that represents the number of such messages that have been sent in the series. Procedure Next_Ticks below is called at a relatively high frequency to scan the list of active series and issue tick messages as appropriate. I asked a new student to add support for the counters to this system as a simple way of getting into our code and helping us to make forward progress. The version produced was something like:

procedure Next_Ticks is
begin
-- Iterate through the array to see who needs a tick message.
for I in Series_Array'Range loop
declare
Current_Series : Series_Record renames Series_Array(I);
begin
-- If we need to send a tick from this series...
if Current_Series.Is_Used and then Current_Series.Next <= Current_Time then
Route_Message
Request_ID => 0,
Series_ID  => Current_Series.ID,
Count      => Current_Series.Count));

-- Update the current record.
case Current_Series.Kind is
when One_Shot =>
-- TODO: Should we reinitialize the rest of the series record?
Current_Series.Is_Used := False;

when Periodic =>
Current_Series.Count := Current_Series.Count + 1;
Current_Series.Next :=
Current_Series.Next + Current_Series.Interval;
end case;
end if;
end;
end loop;
end Next_Ticks;

The code iterates over an array of Series_Record objects looking for the records that are active and are due to be ticked (Current_Series.Next <= Current_Time). For those records, it invokes the Route_Message procedure on an appropriately filled in Tick Reply message as created by function Tick_Reply_Encode. The student had to modify the information saved for each series to contain a counter, modify the Tick Reply message to contain a count, and update Tick_Reply_Encode to allow for a count parameter. The student also needed to increment the count as needed for periodic messages. This involved relatively minor changes to a variety of areas in the code.

The SPARK tools ensured that the student dealt with initialization issues correctly. But they also found an issue in the code above with surprisingly far reaching consequences. In particular, SPARK could not prove that no runtime error could ever occur on the line

Current_Series.Count := Current_Series.Count + 1;

Of course, this will raise Constraint_Error when the count overflows the range of the data type used. If periodic tick messages are sent for too long, the exception would eventually be raised, potentially crashing the system. Depending on the rate of tick messages, that might occur in as little as a couple of years, a very realistic amount of time for the duration of a space mission. It was also easy to see how the error could be missed during ordinary testing. A code review might have discovered the problem, but we did code reviews infrequently. On the other hand, SPARK made this problem immediately obvious and led to an extensive discussion about what should be done in the event of the tick counters overflowing.

In November 2013 we launched a low Earth orbiting CubeSat. The launch vehicle contained 13 other university built CubeSats. Most were never heard from. One worked for a few months. Ours worked for two years until it reentered Earth's atmosphere as planned in November 2015. Although the reasons for the other failures are not always clear, software problems were known to be an issue in at least one of them and probably for many others. We believe the success of our mission, particularly in light of the small size and experience of our student team, is directly attributable to the use of SPARK.

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MISRA C is the most widely known coding standard restricting the use of the C programming language for critical software. For good reasons. For one, its focus is entirely on avoiding error-prone programming features of the C programming language rather than on enforcing a particular programming style. In addition, a large majority of rules it defines are checkable automatically (116 rules out of the total 159 guidelines), and many tools are available to enforce those. As a coding standard, MISRA C even goes out of its way to define a consistent sub-language of C, with its own typing rules (called the "essential type model" in MISRA C) to make up for the lack of strong typing in C.

That being said, it's still a long shot to call it a security coding standard for C. Even when taking into account the 14 additional guidelines focusing on security of the "MISRA C:2012 - Amendment 1: Additional security guidelines for MISRA C:2012". MISRA C is first and foremost focusing on software quality, which has obvious consequences for security, but programs in MISRA C remain for the most part vulnerable to the major security vulnerabilities that plague C programs.

In particular, it's hard to state what guarantees are obtained when respecting the MISRA C rules (which means essentially respecting the 116 decidable rules enforced automatically by analysis tools). In order to clarify this, and to present at the same time how guarantees can be obtained using a different programming language, we have written a book available online. Even better, we host on our e-learning website an interactive version of the book where you can compile C or Ada code, and analyze SPARK code, to experiment how a different language with its associated analysis toolset can go beyond what MISRA C allows.

So that, even if MISRA C is the best thing that could happen to C, you can decide if C is really the best thing that could happen to your software.

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Like last year, we've sent a squad of AdaCore engineers to participate in the celebration of Open Source software at FOSDEM. Like last year, we had great interactions with the rest of the Ada and SPARK Community in the Ada devroom on Saturday. You can check the program with videos of all the talks here. This year's edition was particularly diverse, with an academic project from Austria for an autonomous train control in Ada, two talks on Database development and Web development made type-safe with Ada, distributed computing, libraries, C++ binding, concurrency, safe pointers, etc.

We also had a talk in the RISC-V devroom:

And there was a related talk in the Security devroom on the use of SPARK for security:

Hope to see you at FOSDEM next year!

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In Part 1 of this blog post I discussed why I chose to implement this application using the Ada Web Server to serve the computed fractal to a web browser. In this part I will discuss a bit more about the backend of the application, the Ada part.

Why do we care about performance?

The Ada backend will compute a new image each time one is requested from the front-end. The front-end immediately requests a new image right after it receives the last one it requested. Ideally, this will present to the user as an animation of a Mandelbrot fractal changing colors. If the update is too slow, the animation will look slow and terrible. So we want to minimize the compute time as much as possible. We have a few ways to do that: optimize the computation, and parallelize the computation.

Parallelizing the Fractal Computation

An interesting feature of the Mandelbrot calculation is that the computation of any pixel is independent of any other pixel. That means we can completely parallelize the computation of each pixel! So if our requested image is 1920x1280 pixels we can spawn 2457600 tasks right? Theoretically, yes we can do that. But it doesn’t necessarily speed up our application compared to, let’s say, 8 or 16 tasks, each of which computes a row or selection of rows. Either way, we know we need to create what’s called a task pool, or a group of tasks that can be queued up as needed to do some calculations. We will create our task pool by creating a task type, which will implement the actual activity that the task will complete, and we will use a Synchronous_Barrier to synchronize all of the tasks back together.

task type Chunk_Task_Type is
pragma Priority (0);
entry Go (Start_Row : Natural;
Stop_Row  : Natural;
Buf       : Stream_Element_Array_Access);

Start_Row : Natural;
Stop_Row  : Natural;
end record;

S_Sync_Obj  : Synchronous_Barrier (Release_Threshold => Task_Pool_Size + 1);

This snippet is from fractal.ads. S_Task_Pool will be our task pool and S_Sync_Obj will be our synchronization object. If you notice, each Chunk_Task_Type in the S_Task_Pool takes an access to a buffer in its Go entry procedure. In our implementation, each task will have an access to the same buffer. Isn’t this a race condition? Shouldn’t we use a protected object?

The downside of protected object

The answer to both of those questions, is yes. This is a race condition and we should be using a protected object. If you were to run CodePeer on this project, it identifies this as a definite problem. But using a protected object is going to destroy our performance. The reason for this is because under the hood, the protected object will be using locks each time we access data from the buffer. Each lock, unlock, and wait-on-lock call is going to make the animation of our fractal look worse and worse. However, there is a way to get around this race-condition issue. By design, we can guarantee that each task is going to access the buffer from Start_Row to Stop_Row. So, by design we can make sure that each task’s rows don’t overlap another task’s rows thereby avoiding a race condition.

Parallelization Implementation

Now that we understand the specification of the task pool, let’s look at the implementation.

task body Chunk_Task_Type
is
Start   : Natural;
Stop    : Natural;
Buffer  : Stream_Element_Array_Access;
Notified : Boolean;
begin

loop
accept Go (Start_Row : Natural;
Stop_Row : Natural;
Buf : Stream_Element_Array_Access) do
Start := Start_Row;
Stop := Stop_Row;
Buffer := Buf;
end Go;

for I in Start .. Stop loop

Calculate_Row (Y      => I,
Idx    => Buffer'First +
Stream_Element_Offset ((I - 1) *
Get_Width * Pixel'Size / 8),
Buffer => Buffer);
end loop;
Wait_For_Release (The_Barrier => S_Sync_Obj,
Notified    =>  Notified);
end loop;
end Chunk_Task_Type;

Now we come back to the size of the task pool. Based on the implementation above we can chunk the processing up to the granularity of at least the number of rows being requested. So in the case of a 1920x1280 image, we could have 1280 tasks! But we have to ask ourselves, is that going to give us better performance than 8 or 16 tasks? The answer, unfortunately, is probably not. If we create 8 tasks, and we have an 8 core processor, we can assume that some of our tasks are going to execute on different cores in parallel. If we create 1280 tasks and we use the same 8 core processor, we don’t get much more parallelization than with 8 tasks. This is a place where tuning and best judgement will give you the best performance.

Fixed vs Floating Point

Now that we have the parallelization component, let’s think about optimizing the maths. In most fractal computations floating point complex numbers are used. Based on our knowledge of processors, we can assume that in most cases floating point calculations will be slower than integer calculations. So theoretically, using fixed point numbers might give us better performance. For more information on Ada fixed point types check out the Fixed-point types section of the Introduction to Ada course on learn.adacore.com .

The Generic Real Type

Because we are going to use the same algorithm for both floating and fixed point math, we can implement the algorithm using a generic type called Real. The Real type is defined in computation_type.ads.

generic
type Real is private;
with function "*" (Left, Right : Real) return Real is <>;
with function "/" (Left, Right : Real) return Real is <>;
with function To_Real (V : Integer) return Real is <>;
with function F_To_Real (V : Float) return Real is <>;
with function To_Integer (V : Real) return Integer is <>;
with function To_Float (V : Real) return Float is <>;
with function Image (V : Real) return String is <>;
with function "+" (Left, Right : Real) return Real is <>;
with function "-" (Left, Right : Real) return Real is <>;
with function ">" (Left, Right : Real) return Boolean is <>;
with function "<" (Left, Right : Real) return Boolean is <>;
with function "<=" (Left, Right : Real) return Boolean is <>;
with function ">=" (Left, Right : Real) return Boolean is <>;
package Computation_Type is

end Computation_Type;

We can then create instances of the Julia_Set package using a floating point and fixed point version of the computation_type package.

type Real_Float is new Float;

function Integer_To_Float (V : Integer) return Real_Float is
(Real_Float (V));

function Float_To_Integer (V : Real_Float) return Integer is
(Natural (V));

function Float_To_Real_Float (V : Float) return Real_Float is
(Real_Float (V));

function Real_Float_To_Float (V : Real_Float) return Float is
(Float (V));

function Float_Image (V : Real_Float) return String is
(V'Img);

D_Small : constant := 2.0 ** (-21);
type Real_Fixed is delta D_Small range -100.0 .. 201.0 - D_Small;

function "*" (Left, Right : Real_Fixed) return Real_Fixed;
pragma Import (Intrinsic, "*");

function "/" (Left, Right : Real_Fixed) return Real_Fixed;
pragma Import (Intrinsic, "/");

function Integer_To_Fixed (V : Integer) return Real_Fixed is
(Real_Fixed (V));

function Float_To_Fixed (V : Float) return Real_Fixed is
(Real_Fixed (V));

function Fixed_To_Float (V : Real_Fixed) return Float is
(Float (V));

function Fixed_To_Integer (V : Real_Fixed) return Integer is
(Natural (V));

function Fixed_Image (V : Real_Fixed) return String is
(V'Img);

package Fixed_Computation is new Computation_Type (Real       => Real_Fixed,
"*"        => Router_Cb."*",
"/"        => Router_Cb."/",
To_Real    => Integer_To_Fixed,
F_To_Real  => Float_To_Fixed,
To_Integer => Fixed_To_Integer,
To_Float   => Fixed_To_Float,
Image      => Fixed_Image);

package Fixed_Julia is new Julia_Set (CT               => Fixed_Computation,
Escape_Threshold => 100.0);

package Fixed_Julia_Fractal is new Fractal (CT              => Fixed_Computation,
Calculate_Pixel => Fixed_Julia.Calculate_Pixel,

package Float_Computation is new Computation_Type (Real       => Real_Float,
To_Real    => Integer_To_Float,
F_To_Real  => Float_To_Real_Float,
To_Integer => Float_To_Integer,
To_Float   => Real_Float_To_Float,
Image      => Float_Image);

package Float_Julia is new Julia_Set (CT               => Float_Computation,
Escape_Threshold => 100.0);

package Float_Julia_Fractal is new Fractal (CT              => Float_Computation,
Calculate_Pixel => Float_Julia.Calculate_Pixel,
Task_Pool_Size  => Task_Pool_Size);

We now have the Julia_Set package with both the floating point and fixed point implementations. The AWS URI router is set up to serve a floating point image if the GET request URI is “/floating_fractal” and a fixed point image is the request URI is “/fixed_fractal”.</