A real-time (RT) system consists of three parts: the hardware, the operating system, and the application. All three parts must meet the latency deadline for the whole system to correctly function. I have covered how to deal with the hardware and the operating system (OS) in the previous post. The natural next question is: how do we ensure the application behaves correctly under RT constraints? The answer to this is surprising complicated and requires a few posts to cover the basics. In this post, I will discuss the sources of latency coming from the application. From these, best practices for writing an RT application can be established.
In the last post, I did promise to write some code. This unfortunately must be delayed for one last time, as this post is too long as is.
RT vs non-RT sections of the code
Before I begin, it is important to clarify that an RT application has two sections in its code: a non-RT section and a RT section. Code in the RT section must execute within the required deadline while code executing in the non-RT section do not have a (strict) deadline. The program flow of two typical RT applications are shown in Figure 1 below. When the program is initializing or shutting down, it is usually in the non-RT section. During initialization, multiple threads are typically created. Usually there are at least two threads: an RT one and a non-RT one. In the robot application example (left), there are two threads: a non-RT data logger thread and an RT PID controller thread. The initialization and shutdown sections of the RT PID controller thread are also not RT. Only the actual PID controller loop is. A similar architecture exists for the audio production program (right), and basically every other multi-threaded RT application. Most of the time, we only need to worry about sources of latency in the RT sections of the program. Thus, the practices discussed in this post are only necessary in the RT sections and usually unnecessary in the non-RT sections.
Avoid non-constant worst-case execution time
It turns out, there is a single rule that you need to follow to be successful at real-time programming: do not call code in the RT section if you don't know its worst-case execution time. Basically, this is saying that you cannot use non-constant time algorithms, unless the input size is low enough that the actual execution time of the algorithm remains acceptably bounded. Unfortunately, most algorithmic implementations favour throughput over latency, which usually means they are tailored to the average-case execution time. Worst-case execution time could be orders of magnitude worse than the average case and such occasions could occur infrequently (e.g. std::vector.push_back). The worst-case execution times also tend to be under-documented (if at all). If the RT system cannot tolerate any deadline misses, specialized libraries and algorithms are usually needed. However, most RT applications do not need such hard guarantees. In these situations, "normal" libraries such as the C++ standard library can be safely used if the worst-case execution time can be acceptably determined. There is a talk in 2021 by Timur Doumler that presents ways to avoid function calls with unbounded latency in the C++ standard library that I highly recommend for anyone working on RT C++ applications.
However, the implementation of algorithms is not the only type of latency. An additional source of latency exists in the interaction between the application and the OS. Such latency cannot be inferred from reading the application code by itself. As noted in part 1, programming on Linux and other general-purpose operating systems is almost like magic. Although the operating system is juggling between thousands of concurrent tasks with a limited amount of CPU and memory, the application can be written as if it is the only process consuming CPU and memory. The way this is accomplished on Linux (and other general-purpose OSes) is optimized for throughput instead of the worst-case latency. Since RT applications need to ensure that worst-case performance remain acceptable, we must understand the OS-level magic to ensure the application won't miss its deadline. The two main kernel subsystems that we need to understand and cooperate with are: (1) the virtual memory system and (2) the CPU scheduler.
Virtual memory: avoid page faults and use mlockall
When an application accesses memory on a general-purpose OS like Linux, it is almost certainly interacting with virtual memory instead of physical RAM. This is an abstraction that provides each process with its own private memory address space such that writes cannot accidentally overwrite values in memory addresses assigned to other processes. When allocating virtual memory using malloc in Linux, memory is not allocated until the addresses are accessed for the first time during a read or a write, due to a feature known as demand paging. Upon access, a page fault is generated, which causes the kernel to synchronously perform the actual memory allocation, which results in a delay. Kernel memory allocation can take an unbounded amount of time and the RT application may miss its deadlines. Thus, these page faults must be avoided. One way to avoid such page faults is to avoid dynamically allocating memory during the RT sections of the code.
However, simply not allocating memory in RT will not get rid of all page faults. In Linux (and other general-purpose OSes), the sum of memory allocated (via malloc) by all applications running on the system can be greater than the available physical RAM. This feature is known as memory overcommitment. When physical RAM is exhausted, the kernel has two options: (1) kill the process using the out-of-memory (OOM) killer or (2) move blocks of memory from physical RAM into a much-slower secondary storage, like the disk. Obviously, the first case is catastrophic for RT, as a terminated process cannot satisfy latency requirements. The second case is much more subtle: the swapping of memory from physical RAM to disk may occur while the application is sleeping if the system is facing memory exhaustion. When the swapped memory addresses are eventually accessed again by the application, a page fault is generated, which causes the OS to load the values from the disk. Since disk latency can be on the order of milliseconds and such page faults are synchronously processed, it can introduce a delay to the RT sections of the application. This is completely unpredictable by the application and catastrophic for real-time. While running out of memory is something that should be generally avoided for RT applications, the OS kernel may choose to swap physical RAM into the disk even in the absence of memory pressure. To avoid all of this, the application can instruct the OS to lock all its memory addresses in physical RAM, using the mlockall(MCL_CURRENT | MCL_FUTURE) function call.
In the literature, there are a number of references that claim the need to "pre-fault" the memory addresses assigned to the RT process. This usually means writing a value to all memory addresses allocated to the application during initialization. The reason to do this is to counteract the effects of demand paging and ensure that all memory addresses are locked in physical RAM. However, the man page for mlockall shows this to be unnecessary. The following snippet comes from the man page of mlockall:
MCL_CURRENT Lock all pages which are currently mapped into the address space of the process.
MCL_FUTURE Lock all pages which will become mapped into the address space of the process in the future. These could be, for instance, new pages required by a growing heap and stack as well as new memory-mapped files or shared memory regions.
If these two flags are passed to mlockall, all existing memory addresses (pages) will be mapped into physical RAM after the call completes; and all future memory allocations will immediately lock the memory addresses into physical RAM. Combining these two options effectively turns off demand paging for the calling application. Indeed, experiments show that page faults are immediately generated upon memory allocation for both the heap and the stack. Thus, contrary to the numerous resources available online, pre-faulting is unnecessary.
- To avoid page faults generated during memory allocation, avoid dynamic memory allocation in the RT sections entirely. Memory should be allocated before the application enters the RT section, or in the non-RT thread (something that will be discussed more next time). There are other programming techniques, like the object pool pattern, that can be used instead of dynamic memory allocation.
- To avoid memory swapping, lock down all the virtual memory addresses needed by the RT sections to physical RAM with the mlockall(MCL_CURRENT | MCL_FUTURE) function call. This prevents the operating system from swapping the RT application's memory into secondary storage at the OS's discretion.
- mlockall(MCL_CURRENT | MCL_FUTURE) also turns off demand paging, at least for Linux. There is thus no need to pre-fault the stack, despite numerous literature to the contrary.
As a note, code for everything presented here and in the subsequent sections will be presented as a part of the small RT app framework in the next post.
|||For example, the OS may need to free some RAM elsewhere (by possibly moving it to the disk) to be able to satisfy your application's memory allocation request, which make take a while.|
|||Technically, it is possible to perform dynamic memory allocation via malloc if you already reserved a block of memory from the OS. In practise, most malloc implementations are not constant time and may occasionally take a long time even if free memory is already reserved to the application. While it is certainly possible to get a constant-time allocator, it's likely better to keep it simple and not perform any dynamic memory allocations, thus avoiding this problem alltogether.|
|||It is also not clear to me if prefaulting works at all. A quick reading at some of the code that prefaults the stack suggests that it may be optimized out by the compiler, as it has no side effects.|
CPU scheduler: Avoid priority inversion
By default, threads created on Linux are not scheduled using a RT scheduler. The behavior of the default Linux scheduler is quite complex and is not suitable for RT. Thus, threads that require RT behavior must request the RT scheduler through the pthreads API. For brevity, I am not going to present the code that does this now, as it will be presented in the next post. Instead, I want to focus on a much more subtle problem that can cause unbounded latency involving the interaction between the CPU scheduler and the application's mutexes. This bug is famous for affecting the Mars Pathfinder Rover despite the fact that the application is deployed on a hard RTOS (VxWorks).
Non-trivial RT applications usually require both RT and non-RT threads that communicate with each other. Multi-threaded communication require some form of synchronization to avoid data races. In non-RT programming, one simple solution to this problem is to protect access to the shared variables with a mutex. In C++, this is usually coded with std::mutex as defined by the C++ standard library. When such a program runs, access to the shared variable may be serialized in the following sequence:
- Initially, the shared variable has the value of v1.
- Thread 1 acquires lock on the mutex and begins reading/writing to the shared variable with value v2.
- Thread 2 attempts to acquire the lock on the same mutex and is blocked as it is held by Thread 1.
- Thread 1 finishes writing to the variable and releases the lock.
- Thread 2 is unblocked, reads the shared variable that has a value of v2.
This is perfectly acceptable for an application without a bounded latency requirement (i.e. all non-RT apps) as the average latency is likely to be quite low. However, the worst-case latency is unbounded on Linux (and other "general-purpose" operating systems). Thus, mutexes are unacceptable for RT. The root cause for this is the priority inversion problem as demonstrated in Figure 2 below:
The figure depicts three processes with three different priority levels sharing a single CPU. The colour of the rectangles shows the original priority levels of the threads. The colour of the lock status line shows the thread that currently owns the mutex. Each process executes for a duration, which is denoted by the width of the rectangles. Finally, the vertical axis denotes the current priority level of the code executing on the CPU. An application is shown in the top plot which uses a regular mutex (e.g. std::mutex). The low-priority thread of this application, shown in green, acquires a lock via the mutex. Then, the high-priority thread, shown in red, preempts the low-priority thread (at A) as it is scheduled to wake up. The high-priority thread attempts to acquire a lock on the same mutex, which blocks (at B). At this point, the OS scheduler noticed that the high-priority thread is blocked and thus puts it back to sleep. The scheduler then reschedules the low-priority thread, allowing it to finish with its work and release the lock. As this work occurs, an unrelated thread (or even another process) with a slightly higher priority level, shown in orange, preempts the low-priority thread (at C) until it is put back to sleep, which can take an unbounded amount of time. Throughout this time, the high-priority thread cannot resume as it remains blocked by the low-priority thread. In effect, the medium-priority thread is able to block the execution of the high-priority thread indefinitely due to the usage of the regular mutex. Such unbounded latency is always unacceptable for RT.
One way to solve this problem is via mutexes with priority inheritance. The bottom plot of Figure 2 demonstrates this approach. As with the original case, the low-priority thread acquires a lock. The high-priority thread preempts it (at A) and tries to lock the same mutex (at B). This blocks, prompting the OS to put the high-priority thread back to sleep. Noticing that the high-priority thread is blocked on the mutex currently being held by the low-priority thread, the OS switches to the low-priority thread with a temporarily boosted priority level equaling that of the high-priority thread. In effect, the low-priority thread inherited the priority level of the high priority thread, which forbids the OS from interrupting its execution by the medium-priority thread. Once the originally-low-priority thread releases the lock, its priority level is reverted to the original value and the high-priority thread can continue with its execution (at C). Thus, the overall latency remains bounded as long as the code in the critical section (i.e. the duration when it held the lock) of the low-priority thread is bounded.
There are several drawbacks to this approach. Notably, code within the critical sections protected by mutexes on the low-priority thread might occasionally run with RT priority. Thus, such code must be treated as if they are RT code. This requires the code within the mutex's critical sections to follow the best practices outlined in this article. In many situations, this is not desirable. Coupled with other mutex-related problems, lock-free (or more strongly, wait-free) programming can potentially be a more appealing way to get around the need for a mutex. However, this topic is way too big for me to cover now, so I will defer it to a future post. For the time being, you can look into the boost::lockfree package and atomic variables.
- If mutexes are required for RT, always enable priority inheritance. This is not possible with std:mutex and requires the use of pthread mutex directly.
- Investigate into lock-free (wait-free) programming techniques to share data between threads. This is something I'll explore in a future post.
|||There are also examples of using this API in this wiki page.|
|||Data races occur when two or more threads attempt to access the same memory location, where at least one thread performs a write. For interested readers, I recommend this series on memory models across languages and hardware, which goes into these ideas in more detail.|
|||There are some debates over whether using priority inheritance is even a good idea, especially since mutexes with priority inheritance can be quite difficult to implement correctly. Additionally, mutexes can introduce a number of issues, such as deadlocks and performance issues involving thread preemption even in the absence of priority inversion. Notably, audio production software appears to employ lock-free programming heavily, as priority inheritance is not availble on Windows until very recently.|
|||I will try to cover this in a future post. For now, interested readers can take a look at my rt::mutex implementation here.|
Don't trust the OS? Avoid system calls
When an application runs, it usually performs a lot of system calls to instruct the OS kernel to do some work on its behalf, usually synchronously during the application's execution. We have already seen two of them: malloc (via sbrk and mmap) and mlockall. Others may include writing to files and interacting with USB devices. Most of these system calls are hidden behind libraries commonly used by applications. Since Linux was not originally designed to be a RTOS, there are generally no guarantees that a particular system call won't cause page faults or priority inversion problems internally. It might even block the process (such as calls like accept) which causes the process to be scheduled out of the CPU until the call is unblocked. Further, system calls may result in a full context switch, which is associated with a small CPU overhead that may be problematic in some situations.
There are a few solutions to these problems:
- Use an OS where all system calls are documented with worst-case execution time.
- Audit the kernel source code to determine worst-case execution time and ensure the calls used do not block. Alternatively, obtains some sort of "soft" guarantee from someone else that has audited the code.
- Don't trust the kernel, be defensive, and avoid system calls unless absolutely necessary (such as for IO, and getting the current time in high-resolution).
If you want to write an RT application for Linux, the number of distinct system calls used should be kept to a minimum, so it is feasible to audit them and make sure they cannot cause problems. This might feel somewhat shaky, but RT applications in robotics and audio domains have been developed for Linux with (presumably) acceptable performance.
|||I'm not aware of a list of safe and unsafe system calls for Linux. Presumably commercial hard-RTOSes have such a list.|
In the third part of this series, we determined a list of potential sources of latency and came up with the following "best practices":
- Avoid code with non-constant worst-case execution time
- Avoid non constant-time algorithms
- Write fast code
- Profile and optimize the code as necessary
- Avoid excessive context switches
- Avoid CPU cache invalidation if the code relies on cache for speed
- Avoid page faults due to either demand paging or swapping by calling
mlockall(MCL_CURRENT | MCL_FUTURE) and reserving all memory needed before
the RT code sections start.
- There is no need to prefault the stack nor the heap after allocation, contrary to numerous online literature.
- Avoid standard mutexes such as std::mutex by either using priority-inheriting mutexes or lock-free programming instead.
- Avoid system calls where possible, in case the kernel suffers from any of the three issues mentioned above
In the next post, we will see these in action, with an RT application framework as well as an example application.
These are some of the more relevant materials I've reviewed as I wrote this post:
- Challenges Using Linux as a Real-Time Operating System - Michael Madden
- Real-time programming with the C++ standard library - Timur Doumler
- System Performance, 2nd Edition - Brendan Gregg
- Code demonstrating that prefaulting is not needed
- Make multiprocessor computer correctly execute multiprocess programs - Leslie Lamport
- Series on memory model - Russ Cox