Chapter 8. Bottom Halves and Deferring Work

Interrupt handlers, discussed in the previous chapter, can form only the first half of any interrupt processing solution, with the following limitations:

Operating systems need a quick, asynchronous, and simple mechanism for immediately responding to hardware and performing any time-critical actions. Interrupt handlers serve this function well. Less critical work can and should be deferred to a later point when interrupts are enabled.

Consequently, managing interrupts is divided into two parts, or halves.

  1. The first part, interrupt handlers (top halves), are executed by the kernel asynchronously in immediate response to a hardware interrupt, as discussed in the previous chapter.
  2. This chapter looks at the second part of the interrupt solution, bottom halves.

Bottom Halves

The job of bottom halves is to perform any interrupt-related work not performed by the interrupt handler. You want the interrupt handler to perform as little work as possible and in turn be as fast as possible. By offloading as much work as possible to the bottom half, the interrupt handler can return control of the system to whatever it interrupted as quickly as possible.

However, the interrupt handler must perform some work; if the work is timing-sensitive, it makes sense to perform it in the interrupt handler. For example, the interrupt handler almost assuredly needs to acknowledge to the hardware the receipt of the interrupt. It may need to copy data to or from the hardware.

Almost anything else can be performed in the bottom half. For example, if you copy data from the hardware into memory in the top half, it makes sense to process it in the bottom half.

No hard and fast rules exist about what work to perform where. The decision is left entirely up to the device-driver author. Although no arrangement is illegal, an arrangement can certainly be suboptimal.

Since interrupt handlers run asynchronously, with at least the current interrupt line disabled, minimizing their duration is important. The following useful tips help decide how to divide the work between the top and bottom half:

When attempting to write your own device driver, looking at other interrupt handlers and their corresponding bottom halves can help. When deciding how to divide your interrupt processing work between the top and bottom half, ask yourself what must be in the top half and what can be in the bottom half. Generally, the quicker the interrupt handler executes, the better.

Why Bottom Halves?

It is crucial to understand:

You want to limit the amount of work you perform in an interrupt handler because:

Thus, minimizing the time spent with interrupts disabled is important for system response and performance.

Besides, interrupt handlers run asynchronously with respect to other code, even other interrupt handlers. Therefore you should work to minimize how long interrupt handlers run. Processing incoming network traffic should not prevent the kernel's receipt of keystrokes.The solution is to defer some of the work until later.

When is "later"? *

Later is often simply not now. The point of a bottom half is not to do work at some specific point in the future, but simply to defer work until any point in the future when the system is less busy and interrupts are again enabled. Often, bottom halves run immediately after the interrupt returns. The key is that they run with all interrupts enabled.

Linux is not the only operating systems that separates the processing of hardware interrupts into two parts.

This design keeps system latency low by running with interrupts disabled for as little time as necessary.

A World of Bottom Halves

While the top half is implemented entirely via the interrupt handler, multiple mechanisms are available for implementing a bottom half. These mechanisms are different interfaces and subsystems that enable you to implement bottom halves. This chapter discusses multiple methods of implementing bottom halves. [p135]

The Original "Bottom Half"

In the beginning, Linux provided only the "bottom half" for implementing bottom halves. This name was logical because at the time that was the only means available for deferring work. The infrastructure was also known as BH to avoid confusion with the generic term bottom half.

The BH interface was simple.

This was simple and easy to use, but was also inflexible and a bottleneck.

Task Queues

The kernel developers later introduced task queues both as a method of deferring work and as a replacement for the BH mechanism.

The kernel defined a family of queues.

This worked fairly well, but it was still too inflexible to replace the BH interface entirely. It also was not lightweight enough for performance-critical subsystems, such as networking.

Softirqs and Tasklets

The softirqs and tasklets were introduced during the 2.3 development series, to completely replace the BH interface.

Tasklets are a good trade-off between performance and ease of use. For most bottom-half processing, the tasklet is sufficient. Softirqs are useful when performance is critical, such as with networking. Using softirqs requires more care, however, because two of the same softirq can run at the same time. In addition, softirqs must be registered statically at compile time. Conversely, code can dynamically register tasklets.

While developing the 2.5 kernel, all BH users were converted to the other bottom-half interfaces. Additionally, the task queue interface was replaced by the work queue interface. Work queues are a simple yet useful method of queuing work to later be performed in process context.

Consequently, the 2.6 kernel has three bottom-half mechanisms in the kernel:

Kernel Timers *

Kernel timers is another mechanism for deferring work. Unlike the mechanisms discussed in the chapter, timers defer work for a specified amount of time. That is, although the tools discussed in this chapter are useful to defer work to any time but now, you use timers to defer work until at least a specific time has elapsed.

Therefore, timers have different uses than the general mechanisms discussed in this chapter. A full discussion of timers is given in Chapter 11 Timers and Time Management.

Dispelling the Confusion

Bottom half is a generic operating system term referring to the deferred portion of interrupt processing. All the kernel's mechanisms for deferring work are "bottom halves". Some people also confusingly call all bottom halves "softirqs".

Bottom half also refers to the original deferred work mechanism in Linux. This mechanism is also known as a BH, so we call it by that name now and leave the former as a generic description. The BH mechanism was deprecated a while back and fully removed in the 2.5 development kernel series.

In the current three methods that exist for deferring work, tasklets are built on softirqs and work queues are their own subsystem. The following table presents a history of bottom halves.

Bottom Half Status
BH Removed in 2.5
Task queues Removed in 2.5
Softirq Available since 2.3
Tasklet Available since 2.3
Work queues Available since 2.5


Softirqs are rarely used directly; tasklets, which are built on softirqs are a much more common form of bottom half. The softirq code lives in the file kernel/softirq.c in the kernel source tree.

Implementing Softirqs

Softirqs are statically allocated at compile time. Unlike tasklets, you cannot dynamically register and destroy softirqs. Softirqs are represented by the softirq_action structure, which is defined in <linux/interrupt.h>:


struct softirq_action
    void    (*action)(struct softirq_action *);

A 32-entry array of this structure is declared in kernel/softirq.c:

static struct softirq_action softirq_vec[NR_SOFTIRQS];

Each registered softirq consumes one entry in the array. Consequently, there are NR_SOFTIRQS registered softirqs. The number of registered softirqs is statically determined at compile time and cannot be changed dynamically. The kernel enforces a limit of 32 registered softirqs. In the current kernel, only nine exist.

The Softirq Handler

The prototype of a softirq handler, action, looks like:

void softirq_handler(struct softirq_action *);

When the kernel runs a softirq handler, it executes this action function with a pointer to the corresponding softirq_action structure as its argument. For example, if my_softirq pointed to an entry in the softirq_vec array, the kernel would invoke the softirq handler function as:


Note that the kernel passes the entire structure to the softirq handler. This trick enables future additions to the structure without requiring a change in every softirq handler.

A softirq never preempts another softirq. The only event that can preempt a softirq is an interrupt handler. Another softirq (even the same one) can run on another processor, however.

Executing Softirqs

A registered softirq must be marked before it will execute. This is called raising the softirq. Usually, an interrupt handler marks its softirq for execution before returning. Then, at a suitable time, the softirq runs. Pending softirqs are checked for and executed in the following places:

Regardless of the method of invocation, softirq execution occurs in __do_softirq(), which is invoked by do_softirq(). The function is quite simple. If there are pending softirqs, __do_softirq() loops over each one, invoking its handler.

The following code is a simplified variant of the important part of __do_softirq():

u32 pending;

pending = local_softirq_pending();
if (pending) {
    struct softirq_action *h;

    /* reset the pending bitmask */

    h = softirq_vec;
    do {
        if (pending & 1)
        pending >>= 1;
    } while (pending);

It checks for, and executes, any pending softirqs. It specifically does the following:

  1. It sets the pending local variable to the value returned by the local_softirq_pending() macro.
    • This is a 32-bit mask of pending softirqs: if bit n is set, the nth softirq is pending.
  2. After the pending bitmask of softirqs is saved, it clears the actual bitmask.
  3. The pointer h is set to the first entry in the softirq_vec.
  4. If the first bit in pending is set, h->action(h) is called.
  5. The pointer h is incremented by one so that it now points to the second entry in the softirq_vec array.
  6. The bitmask pending is right-shifted by one.
    • This discards the first bit and moves all other bits one place to the right.
  7. The pointer h now points to the second entry in the array, and the pending bitmask now has the second bit as the first. Repeat the previous steps.
  8. Continue repeating until pending is zero, at which point there are no more pending softirqs and the work is done.
    • This check is sufficient to ensure h always points to a valid entry in softirq_vec because pending has at most 32 set bits and thus this loop executes at most 32 times.

Using Softirqs

Softirqs are reserved for the most timing-critical and important bottom-half processing on the system.

Currently, only two subsystems directly use softirqs:

Additionally, kernel timers and tasklets are built on top of softirqs.

If you add a new softirq, you normally want to ask yourself why using a tasklet is insufficient. Tasklets are dynamically created and are simpler to use because of their weaker locking requirements, and they still perform quite well. Nonetheless, for timing-critical applications that can do their own locking in an efficient way, softirqs might be the correct solution.

Assigning an Index

Softirqs are declared statically at compile time via an enum in <linux/interrupt.h>. The kernel uses this index, which starts at zero, as a relative priority. Softirqs with the lowest numerical priority execute before those with a higher numerical priority.

Creating a new softirq includes adding a new entry to this enum. When adding a new softirq, you might not want to simply add your entry to the end of the list; instead, you need to insert the new entry depending on the priority you want to give it. By convention, HI_SOFTIRQ is always the first and RCU_SOFTIRQ is always the last entry. A new entry likely belongs in between BLOCK_SOFTIRQ and TASKLET_SOFTIRQ.

The following table contains a list of the existing softirq types.

Tasklet Priority Softirq Description
HI_SOFTIRQ 0 High-priority tasklets
NET_TX_SOFTIRQ 2 Send network packets
NET_RX_SOFTIRQ 3 Receive network packets
BLOCK_SOFTIRQ 4 Block devices
TASKLET_SOFTIRQ 5 Normal priority tasklets
HRTIMER_SOFTIRQ 7 High-resolution timers
Registering Your Handler

Next, the softirq handler is registered at run-time via open_softirq(), which takes two parameters: the softirq's index and its handler function. For example, the networking subsystem registers its softirqs like this, in net/core/dev.c:


open_softirq(NET_TX_SOFTIRQ, net_tx_action);
open_softirq(NET_RX_SOFTIRQ, net_rx_action);

Here, locking is an important point, and it is the reason tasklets are usually preferred. Simply preventing your softirqs from running concurrently is not ideal. If a softirq obtained a lock to prevent another instance of itself from running simultaneously, there would be no reason to use a softirq. Consequently, most softirq handlers resort to per-processor data (data unique to each processor and thus not requiring locking) and other tricks to avoid explicit locking and provide excellent scalability.

The reason for using softirqs is scalability. If you do not need to scale to infinitely many processors, then use a tasklet. Tasklets are essentially softirqs in which multiple instances of the same handler cannot run concurrently on multiple processors.

Raising Your Softirq

After a handler is added to the enum list and registered via open_softirq(), it is ready to run. To mark it pending, so it is run at the next invocation of do_softirq(), call raise_softirq(). For example, the networking subsystem would call:


This raises the NET_TX_SOFTIRQ softirq. Its handler, net_tx_action(), runs the next time the kernel executes softirqs. This function (raise_softirq()) disables interrupts prior to actually raising the softirq and then restores them to their previous state. If interrupts are already off, the function raise_softirq_irqoff() can be used as a small optimization. For example:

 * interrupts must already be off!

Softirqs are most often raised from within interrupt handlers. In the case of interrupt handlers, the interrupt handler performs the basic hardware-related work, raises the softirq, and then exits. When processing interrupts, the kernel invokes do_softirq(). The softirq then runs and picks up where the interrupt handler left off. In this example, the "top half" and "bottom half" naming should make sense.


Tasklets are a bottom-half mechanism built on top of softirqs. As mentioned, they have nothing to do with tasks. Tasklets are similar in nature and behavior to softirqs, but have a simpler interface and relaxed locking rules.

When writing a device driver, you almost always want to use tasklets. Softirqs are required only for high-frequency and highly threaded uses. Tasklets, on the other hand, work fine for the vast majority of cases and are very easy to use.

Implementing Tasklets

Because tasklets are implemented on top of softirqs, they are softirqs. As discussed, tasklets are represented by two softirqs:

The only difference in these types is that the HI_SOFTIRQ-based tasklets run prior to the TASKLET_SOFTIRQ-based tasklets.

The Tasklet Structure

Tasklets are represented by the tasklet_struct structure. Each structure represents a unique tasklet. The structure is declared in <linux/interrupt.h>:


struct tasklet_struct {
    struct tasklet_struct *next;  /* next tasklet in the list */
    unsigned long state;          /* state of the tasklet */
    atomic_t count;               /* reference counter */
    void (*func)(unsigned long);  /* tasklet handler function */
    unsigned long data;           /* argument to the tasklet function */
Scheduling Tasklets

Scheduled tasklets (the equivalent of raised softirqs) are stored in two per-processor structures:

Both of these structures are linked lists of tasklet_struct structures. Each tasklet_struct structure in the list represents a different tasklet.

Tasklets are scheduled via the following functions:

Either of them receives a pointer to the tasklet's tasklet_struct as its lone argument. Each function ensures that the provided tasklet is not yet scheduled and then calls __tasklet_schedule() and __tasklet_hi_schedule() as appropriate. The two functions are similar. The difference is that one uses TASKLET_SOFTIRQ and one uses HI_SOFTIRQ.

tasklet_schedule() undertakes the following steps:

  1. Check whether the tasklet's state is TASKLET_STATE_SCHED. If it is, the tasklet is already scheduled to run and the function can immediately return.
  2. Call __tasklet_schedule().
  3. Save the state of the interrupt system, and then disable local interrupts. This ensures that nothing on this processor will mess with the tasklet code while tasklet_schedule() is manipulating the tasklets.
  4. Add the tasklet to be scheduled to the head of the tasklet_vec or tasklet_hi_vec linked list, which is unique to each processor in the system.
  5. Raise the TASKLET_SOFTIRQ or HI_SOFTIRQ softirq, so do_softirq() executes this tasklet in the near future.
  6. Restore interrupts to their previous state and return.

do_softirq() is run at the next earliest convenience, (as discussed in the previous section). Because most tasklets and softirqs are marked pending in interrupt handlers, do_softirq() most likely runs when the last interrupt returns. Because TASKLET_SOFTIRQ or HI_SOFTIRQ is now raised, do_softirq() executes the associated handlers. These handlers, tasklet_action() and tasklet_hi_action(), are the heart of tasklet processing; they perform the following steps:

  1. Disable local interrupt delivery (there is no need to first save their state because the code here is always called as a softirq handler and interrupts are always enabled) and retrieve the tasklet_vec or tasklet_hi_vec list for this processor.
  2. Clear the list for this processor by setting it equal to NULL.
  3. Enable local interrupt delivery. Again, there is no need to restore them to their previous state because this function knows that they were always originally enabled.
  4. Loop over each pending tasklet in the retrieved list.
  5. If this is a multiprocessing machine, check whether the tasklet is running on another processor by checking the TASKLET_STATE_RUN flag. If it is currently running, do not execute it now and skip to the next pending tasklet. Recall that only one tasklet of a given type may run concurrently.
  6. If the tasklet is not currently running, set the TASKLET_STATE_RUN flag, so another processor will not run it.
  7. Check for a zero count value, to ensure that the tasklet is not disabled. If the tasklet is disabled, skip it and go to the next pending tasklet.
  8. Run the tasklet handler after ensuring the following:
    • The tasklet is not running elsewhere
    • The tasklet is marked as running so it will not start running elsewhere
    • The tasklet has a zero count value.
  9. After the tasklet runs, clear the TASKLET_STATE_RUN flag in the tasklet's state field.
  10. Repeat for the next pending tasklet, until there are no more scheduled tasklets waiting to run.

The implementation of tasklets is simple, but rather clever:

  1. All tasklets are multiplexed on top of two softirqs, HI_SOFTIRQ and TASKLET_SOFTIRQ.
  2. When a tasklet is scheduled, the kernel raises one of these softirqs.
  3. These softirqs, in turn, are handled by special functions that then run any scheduled tasklets.
  4. The special functions ensure that only one tasklet of a given type runs at the same time. However, other tasklets can run simultaneously.

All this complexity is then hidden behind a clean and simple interface.

Using Tasklets

In most cases, tasklets are the preferred mechanism with which to implement your bottom half for a normal hardware device. Tasklets are dynamically created, easy to use, and quick.

Declaring Your Tasklet

You can create tasklets statically or dynamically, depending on whether you have (or want) a direct or indirect reference to the tasklet. If you are going to statically create the tasklet (and thus have a direct reference to it), use one of two macros in <linux/interrupt.h>:

DECLARE_TASKLET(name, func, data)

Both these macros statically create a struct tasklet_struct with the given name. When the tasklet is scheduled, the given function func is executed and passed the argument data. The difference between the two macros is the initial reference count. The first macro creates the tasklet with a count of zero, and the tasklet is enabled. The second macro sets count to one, and the tasklet is disabled.

For example:

DECLARE_TASKLET(my_tasklet, my_tasklet_handler, dev);

This line is equivalent to

struct tasklet_struct my_tasklet = { NULL, 0, ATOMIC_INIT(0),
                                     my_tasklet_handler, dev };

This creates a tasklet named my_tasklet enabled with my_tasklet_handler as its handler. The value of dev is passed to the handler when it is executed.

To initialize a tasklet given an indirect reference (a pointer) to a dynamically created struct tasklet_struct named t, call tasklet_init() like this:

tasklet_init(t, tasklet_handler, dev); /* dynamically as opposed to statically */
Writing Your Tasklet Handler

The tasklet handler must match the correct prototype:

void tasklet_handler(unsigned long data)
Scheduling Your Tasklet

To schedule a tasklet for execution, tasklet_schedule() is called and passed a pointer to the relevant tasklet_struct:

tasklet_schedule(&my_tasklet); /* mark my_tasklet as pending */

After a tasklet is scheduled, it runs once at some time in the near future. If the same tasklet is scheduled again, before it has had a chance to run, it still runs only once. If it is already running, for example on another processor, the tasklet is rescheduled and runs again. As an optimization, a tasklet always runs on the processor that scheduled it, making better use of the processor's cache.

For example:

tasklet_disable(&my_tasklet); /* tasklet is now disabled */

/* we can now do stuff knowing that the tasklet cannot run .. */

tasklet_enable(&my_tasklet); /* tasklet is now enabled */

You can remove a tasklet from the pending queue via tasklet_kill(). This function receives a pointer as a lone argument to the tasklet's tasklet_struct. Removing a scheduled tasklet from the queue is useful when dealing with a tasklet that often reschedules itself. This function first waits for the tasklet to finish executing and then it removes the tasklet from the queue. Nothing stops some other code from rescheduling the tasklet. This function must not be used from interrupt context because it sleeps.


Softirq processing is aided by a set of per-processor kernel threads. These kernel threads help in the processing of softirqs when the system is overwhelmed with softirqs. Because tasklets are implemented using softirqs, the following discussion applies equally to softirqs and tasklets.

As described, the kernel processes softirqs in a number of places, most commonly on return from handling an interrupt. There are two characteristics with softirqs:

The combination of these two can result in user-space programs being starved of processor time.

Not processing the reactivated softirqs in a timely manner is unacceptable. When softirqs were first designed, this caused a dilemma that needed fixing, and neither of the two obvious solution was a good one, as discussed below:

First solution: keep processing *

The first solution is simply to keep processing softirqs as they come in and to recheck and reprocess any pending softirqs before returning. This ensures that the kernel processes softirqs in a timely manner and, most important, that any reactivated softirqs are also immediately processed. The problem lies in high load environments, in which many softirqs continually reactivate themselves. The kernel might continually service softirqs without accomplishing much else. User-space is neglected; nothing but softirqs and interrupt handlers run. This approach might work fine if the system is never under intense load; if the system experiences moderate interrupt levels, this solution is not acceptable. User-space cannot be starved for significant periods.

Second solution: not handle reactivated softirqs *

The second solution is not to handle reactivated softirqs. On return from interrupt, the kernel merely looks at all pending softirqs and executes them as normal. If any softirqs reactivate themselves, however, they will not run until the next time the kernel handles pending softirqs. This is most likely not until the next interrupt occurs, which can equate to a lengthy amount of time before any new (or reactivated) softirqs are executed. Worse, on an otherwise idle system, it is beneficial to process the softirqs right away. Unfortunately, this approach is oblivious to which processes are runnable. Therefore, although this method prevents starving user-space, it does starve the softirqs and does not take good advantage of an idle system.

Final solution: compromise *

The solution ultimately implemented in the kernel is to not immediately process reactivated softirqs. Instead, if the number of softirqs grows excessive, the kernel wakes up a family of kernel threads to handle the load. The kernel threads run with the lowest possible priority (nice value of 19), which ensures they do not run in lieu of anything important. The advantage it brings are:

There is one thread per processor, each named ksoftirqd/n where n is the processor number. On a two-processor system, they are ksoftirqd/0 and ksoftirqd/1. Having a thread on each processor ensures an idle processor, if available, can always service softirqs. After the threads are initialized, they run a tight loop similar to this:

for (;;) {
    if (!softirq_pending(cpu))


    while (softirq_pending(cpu)) {
        if (need_resched())


The above code does this:

The Old BH Mechanism

(skipped) [p148]

Work Queues

Work queues are a different form of deferring work. Work queues defer work into a kernel thread; this bottom half always runs in process context:

Normally, it is easy to decide between using work queues and softirqs/tasklets:

Indeed, the usual alternative to work queues is kernel threads. Because the kernel developers frown upon creating a new kernel thread, work queues are strongly preferred. They are also easy to use.

If you need a schedulable entity to perform your bottom-half processing, you need work queues. They are the only bottom-half mechanisms that run in process context, and thus the only ones that can sleep. This means they are useful for situations in which you need to allocate a lot of memory, obtain a semaphore, or perform block I/O. If you do not need a kernel thread to handle your deferred work, consider a tasklet instead.

Implementing Work Queues

In its most basic form, the work queue subsystem is an interface for creating kernel threads to handle work queued from elsewhere. These kernel threads are called worker threads. Work queues enables your driver to create a special worker thread to handle deferred work. The work queue subsystem, however, implements and provides a default worker thread for handling work. Therefore, in its most common form, a work queue is a simple interface for deferring work to a generic kernel thread.

The default worker threads are called events/n where n is the processor number; there is one per processor. For example, on a uniprocessor system there is one thread, events/0. A dual processor system would additionally have an events/1 thread. The default worker thread handles deferred work from multiple locations. Many drivers in the kernel defer their bottom-half work to the default thread. Unless a driver or subsystem has a strong requirement for creating its own thread, the default thread is preferred.

Creating your own worker thread might be advantageous if you perform large amounts of processing in the worker thread. Processor-intense and performance-critical work might benefit from its own thread. This also lightens the load on the default threads, which prevents starving the rest of the queued work.

Data Structures Representing the Threads

The worker threads are represented by the workqueue_struct structure:


 * The externally visible workqueue abstraction is an array of
 * per-CPU workqueues:
struct workqueue_struct {
    struct cpu_workqueue_struct cpu_wq[NR_CPUS];
    struct list_head list;
    const char *name;
    int singlethread;
    int freezeable;
    int rt;

This structure, defined in kernel/workqueue.c, contains an array of struct cpu_workqueue_struct, one per processor on the system. Because the worker threads exist on each processor in the system, there is one of these structures per worker thread, per processor, on a given machine. The cpu_workqueue_struct is the core data structure and is also defined in kernel/workqueue.c:

struct cpu_workqueue_struct {
    spinlock_t lock;             /* lock protecting this structure */
    struct list_head worklist;   /* list of work */
    wait_queue_head_t more_work;
    struct work_struct *current_struct;
    struct workqueue_struct *wq; /* associated workqueue_struct */
    task_t *thread;              /* associated thread */

Note that each type of worker thread has one workqueue_struct associated to it. Inside, there is one cpu_workqueue_struct for every thread and, thus, every processor, because there is one worker thread on each processor.

Data Structures Representing the Work

All worker threads are implemented as normal kernel threads running the worker_thread() function. After initial setup, this function enters an infinite loop and goes to sleep. When work is queued, the thread is awakened and processes the work. When there is no work left to process, it goes back to sleep.

The work is represented by the work_struct structure, defined in <linux/workqueue.h>:


struct work_struct {
    atomic_long_t data;
    struct list_head entry;
    work_func_t func;

These structures are strung into a linked list, one for each type of queue on each processor. For example, there is one list of deferred work for the generic thread, per processor.

The core of worker_thread is simplified as follows:


for (;;) {
    prepare_to_wait(&cwq->more_work, &wait, TASK_INTERRUPTIBLE);
    if (list_empty(&cwq->worklist))
    finish_wait(&cwq->more_work, &wait);

This function performs the following functions, in an infinite loop:

  1. The thread marks itself sleeping (the task's state is set to TASK_INTERRUPTIBLE) and adds itself to a wait queue.
  2. If the linked list of work is empty, the thread calls schedule() and goes to sleep.
  3. If the list is not empty, the thread does not go to sleep. Instead, it marks itself TASK_RUNNING and removes itself from the wait queue.
  4. If the list is nonempty, the thread calls run_workqueue() to perform the deferred work.

The function run_workqueue() actually performs the deferred work:


while (!list_empty(&cwq->worklist)) {
    struct work_struct *work;
    work_func_t f;
    void *data;
    work = list_entry(cwq->, struct work_struct, entry);
    f = work->func;

This function loops over each entry in the linked list of pending work and executes the func member of the workqueue_struct for each entry in the linked list:

  1. While the list is not empty, it grabs the next entry in the list.
  2. It retrieves the function that should be called, func, and its argument, data.
  3. It removes this entry from the list and clears the pending bit in the structure itself.
  4. It invokes the function.
  5. Repeat.
Work Queue Implementation Summary

The relationship between the different data structures is illustrated in the following graph:

Figure 8.1 The relationship between work, work queues, and the worker threads.

Worker threads are at the highest level.

For example, assume that in addition to the generic events worker type, you also create a falcon worker type. If you have a four-processor computer, then there are four events threads (four cpu_workqueue_struct structures) and four falcon threads (four cpu_workqueue_struct structures). There is one workqueue_struct for the events type and one for the falcon type.

The driver creates work, which it wants to defer to later. The work_struct structure represents this work. This structure contains a pointer to the function that handles the deferred work. The work is submitted to a specific worker thread (in this case, a specific falcon thread). The worker thread then wakes up and performs the queued work.

Most drivers use the existing default worker threads, named events. They are easy and simple. Some more serious situations, however, demand their own worker threads. The XFS filesystem, for example, creates two new types of worker threads.

Using Work Queues

Using work queues is easy. This section covers the default events queue and creating new worker threads.

Creating Work

The first step is actually creating some work to defer. To create the structure statically at runtime, use DECLARE_WORK:

DECLARE_WORK(name, void (*func)(void *), void *data);

This statically creates a work_struct structure named name with handler function func and argument data.

Alternatively, you can create work at runtime via a pointer:

INIT_WORK(struct work_struct *work, void (*func)(void *), void *data);

This dynamically initializes the work queue pointed to by work with handler function func and argument data.

Your Work Queue Handler

The prototype for the work queue handler is:

void work_handler(void *data)

A worker thread executes this function, so the function runs in process context. By default, interrupts are enabled and no locks are held. If needed, the function can sleep.

Despite running in process context, the work handlers cannot access user-space memory because there is no associated user-space memory map for kernel threads. The kernel can access user memory only when running on behalf of a user-space process, such as when executing a system call. Only then is user memory mapped in.

Locking between work queues or other parts of the kernel is handled just as with any other process context code. This makes writing work handlers much easier.

Scheduling Work

Now that the work is created, we can schedule it. To queue a given work's handler function with the default events worker threads, call:


The work is scheduled immediately and is run as soon as the events worker thread on the current processor wakes up.

Sometimes you do not want the work to execute immediately, but instead after some delay. You can schedule work to execute at a given time in the future:

schedule_delayed_work(&work, delay);

In this case, the work_struct represented by &work will not execute for at least delay timer ticks into the future. Using ticks as a unit of time is covered in Chapter 10.

Flushing Work

Queued work is executed when the worker thread next wakes up. Sometimes, you need to ensure that a given batch of work has completed before continuing. This is especially important because:

For these needs, there is a function to flush a given work queue:

void flush_scheduled_work(void);

Note that this function does not cancel any delayed work. Any work that was scheduled via schedule_delayed_work(), and whose delay is not yet up, is not flushed via flush_scheduled_work(). To cancel delayed work, call:

int cancel_delayed_work(struct work_struct *work);

This function cancels the pending work, if any, associated with the given work_struct.

Creating New Work Queues

If the default queue is insufficient for your needs, you can create a new work queue and corresponding worker threads. Because this creates one worker thread per processor, you should create unique work queues only if your code needs the performance of a unique set of threads.

You create a new work queue and the associated worker threads via a simple function:

struct workqueue_struct *create_workqueue(const char *name);

The parameter name is used to name the kernel threads. For example, the default events queue is created via:

struct workqueue_struct *keventd_wq;
keventd_wq = create_workqueue("events");

This function creates all the worker threads (one for each processor in the system) and prepares them to handle work.

Creating work is handled in the same manner regardless of the queue type. After the work is created, the following functions are analogous to schedule_work() and schedule_delayed_work(), except that they work on the given work queue and not the default events queue.

int queue_work(struct workqueue_struct *wq, struct work_struct *work)

int queue_delayed_work(struct workqueue_struct *wq,
                       struct work_struct *work,
                       unsigned long delay)

Finally, you can flush a wait queue via a call to the function:

flush_workqueue(struct workqueue_struct *wq)

As previously discussed, this function works identically to flush_scheduled_work(), except that it waits for the given queue to empty before returning.

The Old Task Queue Mechanism

(skipped) [p155-156]

Which Bottom Half Should I Use?

The decision over which bottom half to use is important. The current 2.6 kernel provides three choices: softirqs, tasklets, and work queues. Tasklets are built on softirqs and, therefore, both are similar. The work queue mechanism is an entirely different creature and is built on kernel threads.

Softirqs: least serialization, for highly threaded code *

Softirqs, by design, provide the least serialization. This requires softirq handlers to go through extra steps to ensure that shared data is safe because two or more softirqs of the same type may run concurrently on different processors. If the code in question is already highly threaded, such as in a networking subsystem that is chest-deep in per-processor variables, softirqs make a good choice. They are certainly the fastest alternative for timing-critical and high-frequency uses.

Tasklets: simple interface, for less threaded code *

Tasklets make more sense if the code is not finely threaded. They have a simpler interface and, because two tasklets of the same type might not run concurrently, they are easier to implement. Tasklets are effectively softirqs that do not run concurrently. A driver developer should always choose tasklets over softirqs, unless prepared to utilize per-processor variables or similar magic to ensure that the softirq can safely run concurrently on multiple processors.

Work queues: process context *

If the deferred work needs to run in process context, the only choice of the three is work queues. If process context is not a requirements (specifically, if you have no need to sleep), softirqs or tasklets are perhaps better suited. Work queues involve the highest overhead because they involve kernel threads and, therefore, context switching. This doesn't mean they are inefficient, but in light of thousands of interrupts hitting per second (as the networking subsystem might experience), other methods make more sense. However, work queues are sufficient for most situations.

Softirqs vs. tasklets vs. work queues

In terms of ease of use, work queues wins. Using the default events queue is easy. Next come tasklets, which also have a simple interface. Coming in last are softirqs, which need to be statically created and require careful thinking with their implementation.

The following table is a comparison between the three bottom-half interfaces.

Bottom Half Context Inherent Serialization
Softirq Interrupt None
Tasklet Interrupt Against the same tasklet
Work queues Process None (scheduled as process context)

In short, normal driver writers have two choices:

If you need a schedulable entity to perform your deferred work, and if fundamentally, you need to sleep for any reason, then work queues are your only option. Otherwise, tasklets are preferred. Only if scalability becomes a concern do you investigate softirqs.

Locking Between the Bottom Halves

Locking is not discussed yet, which is an expansive topic in the next two chapters. It is crucial to protect shared data from concurrent access while using bottom halves, even on a single processor machine. A bottom half can run at virtually any moment.

One benefit of tasklets is that they are serialized with respect to themselves. The same tasklet will not run concurrently, even on two different processors. This means you do not have to worry about intra-tasklet concurrency issues. Inter-tasklet concurrency (when two different tasklets share the same data) requires proper locking.

Because softirqs provide no serialization, (even two instances of the same softirq might run simultaneously), all shared data needs an appropriate lock.

If process context code and a bottom half share data, you need to do both of the following before accessing the data:

This ensures both ensures local and SMP protection and prevents a deadlock.

If interrupt context code and a bottom half share data, you need to do both of the following before accessing the data:

This also ensures both local and SMP protection and prevents a deadlock.

Any shared data in a work queue requires locking.

The locking issues are no different from normal kernel code because work queues run in process context.

Chapter 9 provides a background on the issues surrounding concurrency, and Chapter 10 covers the kernel locking primitives. These chapters cover how to protect data that bottom halves use.

Disabling Bottom Halves

It is normally not sufficient to only disable bottom halves. To safely protect shared data, you usually need to obtain a lock and disable bottom halves. You might use such methods in a driver, which are covered in Chapter 10. If you are writing core kernel code, however, you might need to disable just the bottom halves.

The function is misnamed; no one bothered to change the name when the BH interface gave way to softirqs. The following table is a summary of these functions.

Method Description
void local_bh_disable() Disables softirq and tasklet processing on the local processor
void local_bh_enable() Enables softirq and tasklet processing on the local processor

The calls can be nested; only the final call to local_bh_enable() actually enables bottom halves. For example, the first time local_bh_disable() is called, local softirq processing is disabled. If local_bh_disable() is called three more times, local processing remains disabled. Processing is not reenabled until the fourth call to local_bh_enable().

The functions accomplish this by maintaining a per-task counter via the preempt_count, the same counter used by kernel preemption. This counter is used both by the interrupt and bottom-half subsystems. Thus, in Linux, a single pertask counter represents the atomicity of a task. This has proved useful for work such as debugging sleeping-while-atomic bugs. When the counter reaches zero, bottom-half processing is possible. Because bottom halves were disabled, local_bh_enable() also checks for any pending bottom halves and executes them.

The functions are unique to each supported architecture and are usually written as complicated macros in <asm/softirq.h>. The following are close C representations:

* disable local bottom halves by incrementing the preempt_count
void local_bh_disable(void)
    struct thread_info *t = current_thread_info();
    t->preempt_count += SOFTIRQ_OFFSET;

* decrement the preempt_count - this will ‘automatically’ enable
* bottom halves if the count returns to zero
* optionally run any bottom halves that are pending
void local_bh_enable(void)
    struct thread_info *t = current_thread_info();
    t->preempt_count -= SOFTIRQ_OFFSET;

    * is preempt_count zero and are any bottom halves pending?
    * if so, run them
    if (unlikely(!t->preempt_count && softirq_pending(smp_processor_id())))

These calls do not disable the execution of work queues.

Because work queues run in process context, there are no issues with asynchronous execution, and thus, there is no need to disable them. However, because softirqs and tasklets can occur asynchronously (for example, on return from handling an interrupt), kernel code may need to disable them. With work queues, on the other hand, protecting shared data is the same as in any process context. Chapters 8 and Chapter 9 give the details.


This chapter covers the three mechanisms used to defer work in the Linux kernel: softirqs, tasklets, and work queues and goes over their design and implementation. We also discussed a lot about synchronization and concurrency because such topics apply quite a bit to bottom halves. We even wrapped up the chapter with a discussion on disabling bottom halves for reasons of concurrency protection.

Chapter 9 discusses kernel synchronization and concurrency in the abstract, providing a foundation for understanding the issues at the heart of the problem. Chapter 10 discusses the specific interfaces provided by the kernel to solve these problems.

Doubts and Solution


p141 on softirq

This function (raise_softirq()) disables interrupts prior to actually raising the softirq and then restores them to their previous state.

Question: Why would raise_softirq() disable interrupt?

p145 on scheduling tasklets

After a tasklet is scheduled, it runs once at some time in the near future. If the same tasklet is scheduled again, before it has had a chance to run, it still runs only once. If it is already running, for example on another processor, the tasklet is rescheduled and runs again

Question: I'm confused. What does it mean?