I'm trying to make sure that a unique user process executes as soon as possible after a particular hardware interrupt occurs.
One mechanism I'm aware of for doing this is to write a small kernel module that exports a device while sleeping inside the read handler. The module also registers an irq handler, which does nothing but wake the process. Then from the user's perspective, reads to that device block until the relevant interrupt occurs.
(1) On a modern CPU with a mainline kernel, can you reliably expect sub millisecond latency between the kernel seeing the interrupt and the user process regaining control with this?
(2) Are there any lower latency mechanisms on a mainline kernel?
Apply the PREEMPT_RT patch to the kernel and compile it configuring full preemptability through make menuconfig.
This will allow you to have threaded interrupts (i.e., interrupt handlers executed as kernel threads). Then, you can assign maximum priority (i.e., RT prio > 50) to your specific interrupt handler (check its PID using ps aux) and to your specific process, and a lower priority to anything else.
I have written a Kernel module that is interacting with net-filter hooks.
The net-filter hooks operate in Softirq context.
I am accessing a global data structure
"Hash Table" from the softirq context as well as from Process context. The process context access is due to a sysctl file being used to modify the contents of the Hash-table.
I am using spinlock_irq_save.
Is this choice of spin_lock api correct ?? In terms of performance and locking standards.
what would happen if an interrupt is scheduled on another processor? while on the current processor lock is already hold by a process context code?
Firstly:
So, with all the above details I concluded that my softirqs can run concurrently on both cores.
Yes, this is correct. Your softirq handler may be executed "simultaneously on more than one CPU".
Your conclusion to use spinlocks sounds correct to me. However, this assumes that the critical section (ie., that which is executed with the spinlock held) has the following properties:
It must not sleep (for example, acquire a blocking mutex)
It should be as short as possible
Generally, if you're just updating your hash table, you should be fine here.
If an IRQ handler tries to acquire a spinlock that is held by a process context, that's fine. As long as your process context does not sleep with that lock held, the lock should be released within a short amount of time, allowing the IRQ handler to make forward progress.
I think the solution is appropriate . Softirqs anyways runs with preemption disabled . To share a data with a process, the process must also disable both preemption and interrupts. In case of timer, which only reduces the time stamp of an entry can do it atomically i.e. the time stamp variable must be atomic. If in another core softirqs run and wants to acquire the spinlock, when it is already held in the other core,it must wait.
This is purely academic question, I don't really need to know this information for anything, but I would like to understand kernel a bit more :-)
According to kernel documentation http://www.tldp.org/LDP/tlk/kernel/processes.html processes in linux kernel have following states:
Running
The process is either running (it is the current process in the
system) or it is ready to run (it is waiting to be assigned to one of
the system's CPUs).
Waiting
The process is waiting for an event or for a resource. Linux
differentiates between two types of waiting process; interruptible and
uninterruptible. Interruptible waiting processes can be interrupted by
signals whereas uninterruptible waiting processes are waiting directly
on hardware conditions and cannot be interrupted under any
circumstances.
Stopped
The process has been stopped, usually by receiving a signal. A process
that is being debugged can be in a stopped state.
Zombie
This is a halted process which, for some reason, still has a
task_struct data structure in the task vector. It is what it sounds
like, a dead process.
As you can see, when I take a snapshot of processes state, using command like ps, I can see, if it's in Running state, that process either was literally Running or just waiting to be assigned to some CPU by kernel.
In my opinion, these 2 states (that are actually both represented by 1 state in task_struct) are quite different.
Why there is no state like "Ready" that would mean the process is "ready to run" but wasn't assigned to any CPU so far, so that the task_struct would be more clear about the real state? Is it even possible to retrieve this information, or is it secret for whatever reason which process is "literally running" on the CPU?
The struct task_struct contains a long to represent current state:
volatile long state; /* -1 unrunnable, 0 runnable, >0 stopped */
This simply indicates if a process is 'runnable'.
To see the currently executing process you should look at the runqueue. Specifically a struct rq (as defined in kernel/sched/sched.h) contains:
struct task_struct *curr, *idle, *stop;
The pointer *curr is the currently running process on this runqueue (there exists a runqueue per CPU).
You should consult files under kernel/sched/ to see how the Kernel determines which processes should be scheduled according to the different scheduling algorithms if you are interested in exactly how it arrives at the running state.
This is not a linux-kernel answer but a more general about scheduling ^^
A core part of any OS is the Scheduler: http://en.wikipedia.org/wiki/Process_scheduler
Many of them work giving every process a time slice of execution and letting each of them do a little bit of work before switching (referred as a context switch) to another process.
Since the length of a time slice is in the order of milliseconds by the time the information you requested is shown, the state has surely changed so differentiate between "Really Running" and "Ready-but-not-really-running" could result (most of the time) in inaccurate informations.
In ubuntu10.04 linux kernel if I insmod a module which runs
while(1);
in init_module part, entire system stops.
However, if I load a sys file in Windows 7
which runs while(1); in DriverEntry part,
system gets slow but still works.
can someone explain me why two system differs
and what is happening inside kernel?...
I think in first case(infinite loop in init_module),
there is no reason the system stops. because
even if I make while(1); in init_module, it is running
in context of insmod user application program.
so the flow infinite loop has to be scheduled by hardware interrupt signal.
This is just my opinion, I want to know the details if I am wrong...
init_module() is a system call, it runs in kernel space and not in user space.
From what you have observed, it looks like the NT kernel performs module initialization in parallel, whereas the Linux kernel does it sequentially. It might have to do with their respective architectures, NT being a hybrid kernel and Linux being monolithic.
Adding to Frédéric's answer: on Windows the DriverEntry function runs at IRQL PASSIVE_LEVEL (same as virtually all user mode code, all if we exclude APCs). Which means that it can be interrupted by any code running at a higher IRQL at any point. So what you probably encounter here is that the thread that goes into the infinite loop is still being scheduled (thus consuming CPU time), but due to its (low) IRQL it isn't able to starve the system threads or much of the other code that is running. It will, however, be able to starve user mode threads. The effect can be anything from a slowdown to a perceived hanging system.
I am reading following article by Robert Love
http://www.linuxjournal.com/article/6916
that says
"...Let's discuss the fact that work queues run in process context. This is in contrast to the other bottom-half mechanisms, which all run in interrupt context. Code running in interrupt context is unable to sleep, or block, because interrupt context does not have a backing process with which to reschedule. Therefore, because interrupt handlers are not associated with a process, there is nothing for the scheduler to put to sleep and, more importantly, nothing for the scheduler to wake up..."
I don't get it. AFAIK, scheduler in the kernel is O(1), that is implemented through the bitmap. So what stops the scehduler from putting interrupt context to sleep and taking next schedulable process and passing it the control?
So what stops the scehduler from putting interrupt context to sleep and taking next schedulable process and passing it the control?
The problem is that the interrupt context is not a process, and therefore cannot be put to sleep.
When an interrupt occurs, the processor saves the registers onto the stack and jumps to the start of the interrupt service routine. This means that when the interrupt handler is running, it is running in the context of the process that was executing when the interrupt occurred. The interrupt is executing on that process's stack, and when the interrupt handler completes, that process will resume executing.
If you tried to sleep or block inside an interrupt handler, you would wind up not only stopping the interrupt handler, but also the process it interrupted. This could be dangerous, as the interrupt handler has no way of knowing what the interrupted process was doing, or even if it is safe for that process to be suspended.
A simple scenario where things could go wrong would be a deadlock between the interrupt handler and the process it interrupts.
Process1 enters kernel mode.
Process1 acquires LockA.
Interrupt occurs.
ISR starts executing using Process1's stack.
ISR tries to acquire LockA.
ISR calls sleep to wait for LockA to be released.
At this point, you have a deadlock. Process1 can't resume execution until the ISR is done with its stack. But the ISR is blocked waiting for Process1 to release LockA.
I think it's a design idea.
Sure, you can design a system that you can sleep in interrupt, but except to make to the system hard to comprehend and complicated(many many situation you have to take into account), that's does not help anything. So from a design view, declare interrupt handler as can not sleep is very clear and easy to implement.
From Robert Love (a kernel hacker):
http://permalink.gmane.org/gmane.linux.kernel.kernelnewbies/1791
You cannot sleep in an interrupt handler because interrupts do not have
a backing process context, and thus there is nothing to reschedule back
into. In other words, interrupt handlers are not associated with a task,
so there is nothing to "put to sleep" and (more importantly) "nothing to
wake up". They must run atomically.
This is not unlike other operating systems. In most operating systems,
interrupts are not threaded. Bottom halves often are, however.
The reason the page fault handler can sleep is that it is invoked only
by code that is running in process context. Because the kernel's own
memory is not pagable, only user-space memory accesses can result in a
page fault. Thus, only a few certain places (such as calls to
copy_{to,from}_user()) can cause a page fault within the kernel. Those
places must all be made by code that can sleep (i.e., process context,
no locks, et cetera).
Because the thread switching infrastructure is unusable at that point. When servicing an interrupt, only stuff of higher priority can execute - See the Intel Software Developer's Manual on interrupt, task and processor priority. If you did allow another thread to execute (which you imply in your question that it would be easy to do), you wouldn't be able to let it do anything - if it caused a page fault, you'd have to use services in the kernel that are unusable while the interrupt is being serviced (see below for why).
Typically, your only goal in an interrupt routine is to get the device to stop interrupting and queue something at a lower interrupt level (in unix this is typically a non-interrupt level, but for Windows, it's dispatch, apc or passive level) to do the heavy lifting where you have access to more features of the kernel/os. See - Implementing a handler.
It's a property of how O/S's have to work, not something inherent in Linux. An interrupt routine can execute at any point so the state of what you interrupted is inconsistent. If you interrupted the thread scheduling code, its state is inconsistent so you can't be sure you can "sleep" and switch threads. Even if you protect the thread switching code from being interrupted, thread switching is a very high level feature of the O/S and if you protected everything it relies on, an interrupt becomes more of a suggestion than the imperative implied by its name.
So what stops the scehduler from putting interrupt context to sleep and taking next schedulable process and passing it the control?
Scheduling happens on timer interrupts. The basic rule is that only one interrupt can be open at a time, so if you go to sleep in the "got data from device X" interrupt, the timer interrupt cannot run to schedule it out.
Interrupts also happen many times and overlap. If you put the "got data" interrupt to sleep, and then get more data, what happens? It's confusing (and fragile) enough that the catch-all rule is: no sleeping in interrupts. You will do it wrong.
Disallowing an interrupt handler to block is a design choice. When some data is on the device, the interrupt handler intercepts the current process, prepares the transfer of the data and enables the interrupt; before the handler enables the current interrupt, the device has to hang. We want keep our I/O busy and our system responsive, then we had better not block the interrupt handler.
I don't think the "unstable states" are an essential reason. Processes, no matter they are in user-mode or kernel-mode, should be aware that they may be interrupted by interrupts. If some kernel-mode data structure will be accessed by both interrupt handler and the current process, and race condition exists, then the current process should disable local interrupts, and moreover for multi-processor architectures, spinlocks should be used to during the critical sections.
I also don't think if the interrupt handler were blocked, it cannot be waken up. When we say "block", basically it means that the blocked process is waiting for some event/resource, so it links itself into some wait-queue for that event/resource. Whenever the resource is released, the releasing process is responsible for waking up the waiting process(es).
However, the really annoying thing is that the blocked process can do nothing during the blocking time; it did nothing wrong for this punishment, which is unfair. And nobody could surely predict the blocking time, so the innocent process has to wait for unclear reason and for unlimited time.
Even if you could put an ISR to sleep, you wouldn't want to do it. You want your ISRs to be as fast as possible to reduce the risk of missing subsequent interrupts.
The linux kernel has two ways to allocate interrupt stack. One is on the kernel stack of the interrupted process, the other is a dedicated interrupt stack per CPU. If the interrupt context is saved on the dedicated interrupt stack per CPU, then indeed the interrupt context is completely not associated with any process. The "current" macro will produce an invalid pointer to current running process, since the "current" macro with some architecture are computed with the stack pointer. The stack pointer in the interrupt context may point to the dedicated interrupt stack, not the kernel stack of some process.
By nature, the question is whether in interrupt handler you can get a valid "current" (address to the current process task_structure), if yes, it's possible to modify the content there accordingly to make it into "sleep" state, which can be back by scheduler later if the state get changed somehow. The answer may be hardware-dependent.
But in ARM, it's impossible since 'current' is irrelevant to process under interrupt mode. See the code below:
#linux/arch/arm/include/asm/thread_info.h
94 static inline struct thread_info *current_thread_info(void)
95 {
96 register unsigned long sp asm ("sp");
97 return (struct thread_info *)(sp & ~(THREAD_SIZE - 1));
98 }
sp in USER mode and SVC mode are the "same" ("same" here not mean they're equal, instead, user mode's sp point to user space stack, while svc mode's sp r13_svc point to the kernel stack, where the user process's task_structure was updated at previous task switch, When a system call occurs, the process enter kernel space again, when the sp (sp_svc) is still not changed, these 2 sp are associated with each other, in this sense, they're 'same'), So under SVC mode, kernel code can get the valid 'current'. But under other privileged modes, say interrupt mode, sp is 'different', point to dedicated address defined in cpu_init(). The 'current' calculated under these mode will be irrelevant to the interrupted process, accessing it will result in unexpected behaviors. That's why it's always said that system call can sleep but interrupt handler can't, system call works on process context but interrupt not.
High-level interrupt handlers mask the operations of all lower-priority interrupts, including those of the system timer interrupt. Consequently, the interrupt handler must avoid involving itself in an activity that might cause it to sleep. If the handler sleeps, then the system may hang because the timer is masked and incapable of scheduling the sleeping thread.
Does this make sense?
If a higher-level interrupt routine gets to the point where the next thing it must do has to happen after a period of time, then it needs to put a request into the timer queue, asking that another interrupt routine be run (at lower priority level) some time later.
When that interrupt routine runs, it would then raise priority level back to the level of the original interrupt routine, and continue execution. This has the same effect as a sleep.
It is just a design/implementation choices in Linux OS. The advantage of this design is simple, but it may not be good for real time OS requirements.
Other OSes have other designs/implementations.
For example, in Solaris, the interrupts could have different priorities, that allows most of devices interrupts are invoked in interrupt threads. The interrupt threads allows sleep because each of interrupt threads has separate stack in the context of the thread.
The interrupt threads design is good for real time threads which should have higher priorities than interrupts.