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Understanding the Linux Kernel (Part 2 of 3)

Part 2 of 3

This article consists of three parts. Part One introduced the choices made by Linux to schedule processes in the abstract. In this part we discusses the data structures used to implement scheduling and the corresponding algorithm.

The Linux scheduling algorithm works by dividing the CPU time into epochs. In a single epoch, every process has a specified time quantum whose duration is computed when the epoch begins. In general, different processes have different time quantum durations. The time quantum value is the maximum CPU time portion assigned to the process in that epoch. When a process has exhausted its time quantum, it is preempted and replaced by another runnable process. Of course, a process can be selected several times from the scheduler in the same epoch, as long as its quantum has not been exhausted--for instance, if it suspends itself to wait for I/O, it preserves some of its time quantum and can be selected again during the same epoch. The epoch ends when all runnable processes have exhausted their quantum; in this case, the scheduler algorithm recomputes the time-quantum durations of all processes and a new epoch begins.

Each process has a base time quantum: it is the time-quantum value assigned by the scheduler to the process if it has exhausted its quantum in the previous epoch. The users can change the base time quantum of their processes by using the nice( ) and setpriority( ) system calls. A new process always inherits the base time quantum of its parent.

The INIT_TASK macro sets the value of the base time quantum of process 0 (swapper) to DEF_PRIORITY; that macro is defined as follows:

#define DEF_PRIORITY (20*HZ/100)

Since HZ, which denotes the frequency of timer interrupts, is set to 100 for IBM PCs, the value of DEF_PRIORITY is 20 ticks, that is, about 210 ms.

Users rarely change the base time quantum of their processes, so DEF_PRIORITY also denotes the base time quantum of most processes in the system.

In order to select a process to run, the Linux scheduler must consider the priority of each process. Actually, there are two kinds of priority:

Static priority
This kind is assigned by the users to real-time processes and ranges from 1 to 99. It is never changed by the scheduler.

Dynamic priority
This kind applies only to conventional processes; it is essentially the sum of the base time quantum (which is therefore also called the base priority of the process) and of the number of ticks of CPU time left to the process before its quantum expires in the current epoch.

Of course, the static priority of a real-time process is always higher than the dynamic priority of a conventional one: the scheduler will start running conventional processes only when there is no real-time process in a TASK_RUNNING state.

Data Structures Used by the Scheduler
The process list links together all process descriptors, while the runqueue list links together the process descriptors of all runnable processes--that is, of those in a TASK_RUNNING state. In both cases, the init_task process descriptor plays the role of list header.

Each process descriptor includes several fields related to scheduling:

need_resched
A flag checked by ret_from_intr( ) to decide whether to invoke the schedule( ) function.

policy
The scheduling class. The values permitted are:

SCHED_FIFO
A First-In, First-Out real-time process. When the scheduler assigns the CPU to the process, it leaves the process descriptor in its current position in the runqueue list. If no other higher-priority real-time process is runnable, the process will continue to use the CPU as long as it wishes, even if other real-time processes having the same priority are runnable.

SCHED_RR
A Round Robin real-time process. When the scheduler assigns the CPU to the process, it puts the process descriptor at the end of the runqueue list. This policy ensures a fair assignment of CPU time to all SCHED_RR real-time processes that have the same priority.

SCHED_OTHER
A conventional, time-shared process.

The policy field also encodes a SCHED_YIELD binary flag. This flag is set when the process invokes the sched_ yield( ) system call (a way of voluntarily relinquishing the processor without the need to start an I/O operation or go to sleep). The scheduler puts the process descriptor at the bottom of the runqueue list.

rt_priority
The static priority of a real-time process. Conventional processes do not make use of this field.

priority
The base time quantum (or base priority) of the process.

counter
The number of ticks of CPU time left to the process before its quantum expires; when a new epoch begins, this field contains the time-quantum duration of the process. Recall that the update_process_times( ) function decrements the counter field of the current process by 1 at every tick.

When a new process is created, do_fork( ) sets the counter field of both current (the parent) and p (the child) processes in the following way:

current->counter >>= 1;
p->counter = current->counter;

In other words, the number of ticks left to the parent is split in two halves, one for the parent and one for the child. This is done to prevent users from getting an unlimited amount of CPU time by using the following method: the parent process creates a child process that runs the same code and then kills itself; by properly adjusting the creation rate, the child process would always get a fresh quantum before the quantum of its parent expires. This programming trick does not work since the kernel does not reward forks. Similarly, a user cannot hog an unfair share of the processor by starting lots of background processes in a shell or by opening a lot of windows on a graphical desktop. More generally speaking, a process cannot hog resources (unless it has privileges to give itself a real-time policy) by forking multiple descendents.

Notice that the priority and counter fields play different roles for the various kinds of processes. For conventional processes, they are used both to implement time-sharing and to compute the process dynamic priority. For SCHED_RR real-time processes, they are used only to implement time-sharing. Finally, for SCHED_FIFO real-time processes, they are not used at all, because the scheduling algorithm regards the quantum duration as unlimited.

The schedule( ) Function
schedule( ) implements the scheduler. Its objective is to find a process in the runqueue list and then assign the CPU to it. It is invoked, directly or in a lazy way, by several kernel routines.

Direct invocation
The scheduler is invoked directly when the current process must be blocked right away because the resource it needs is not available. In this case, the kernel routine that wants to block it proceeds as follows:

  1. Inserts current in the proper wait queue
  2. Changes the state of current either to TASK_INTERRUPTIBLE or to TASK_UNINTERRUPTIBLE
  3. Invokes schedule( )
  4. Checks if the resource is available; if not, goes to step 2
  5. Once the resource is available, removes current from the wait queue
As can be seen, the kernel routine checks repeatedly whether the resource needed by the process is available; if not, it yields the CPU to some other process by invoking schedule( ). Later, when the scheduler once again grants the CPU to the process, the availability of the resource is again checked.

The scheduler is also directly invoked by many device drivers that execute long iterative tasks. At each iteration cycle, the driver checks the value of the need_resched field and, if necessary, invokes schedule( ) to voluntarily relinquish the CPU.

Lazy invocation
The scheduler can also be invoked in a lazy way by setting the need_resched field of current to 1. Since a check on the value of this field is always made before resuming the execution of a User Mode process, schedule( ) will definitely be invoked at some close future time.

Lazy invocation of the scheduler is performed in the following cases:

  • When current has used up its quantum of CPU time; this is done by the update_process_times( ) function.

  • When a process is woken up and its priority is higher than that of the current process; this task is performed by the reschedule_idle( ) function, which is invoked by the wake_up_process( ) function:

    if (goodness(current, p) > goodness(current, current)) current->need_resched = 1;

  • When a sched_setscheduler( ) or sched_ yield( ) system call is issued.

    Actions performed by schedule( )
    Before actually scheduling a process, the schedule( ) function starts by running the functions left by other kernel control paths in various queues. The function invokes run_task_queue( ) on the tq _scheduler task queue. Linux puts a function in that task queue when it must defer its execution until the next schedule( ) invocation:

    run_task_queue(&tq_scheduler);

    The function then executes all active unmasked bottom halves. These are usually present to perform tasks requested by device drivers:

    if (bh_active & bh_mask)
    do_bottom_half( );

    Now comes the actual scheduling, and therefore a potential process switch.

    The value of current is saved in the prev local variable and the need_resched field of prev is set to 0. The key outcome of the function is to set another local variable called next so that it points to the descriptor of the process selected to replace prev.

    First, a check is made to determine whether prev is a Round Robin real-time process (policy field set to SCHED_RR) that has exhausted its quantum. If so, schedule( ) assigns a new quantum to prev and puts it at the bottom of the runqueue list:

    if (!prev->counter && prev->policy == SCHED_RR) {
    prev->counter = prev->priority;
    move_last_runqueue(prev);
    }

    Now schedule( ) examines the state of prev. If it has nonblocked pending signals and its state is TASK_INTERRUPTIBLE, the function wakes up the process as follows. This action is not the same as assigning the processor to prev; it just gives prev a chance to be selected for execution:

    if (prev->state == TASK_INTERRUPTIBLE &&
    signal_pending(prev))
    prev->state = TASK_RUNNING;

    If prev is not in the TASK_RUNNING state, schedule( ) was directly invoked by the process itself because it had to wait on some external resource; therefore, prev must be removed from the runqueue list:

    if (prev->state != TASK_RUNNING)
    del_from_runqueue(prev);

    Next, schedule( ) must select the process to be executed in the next time quantum. To that end, the function scans the runqueue list. It starts from the process referenced by the next_run field of init_task, which is the descriptor of process 0 (swapper). The objective is to store in next the process descriptor pointer of the highest priority process. In order to do this, next is initialized to the first runnable process to be checked, and c is initialized to its "goodness":

    if (prev->state == TASK_RUNNING) {
    next = prev;
    if (prev->policy & SCHED_YIELD) {
    prev->policy &= ~SCHED_YIELD;
    c = 0;
    } else
    c = goodness(prev, prev);
    } else {
    c = -1000;
    next = &init_task;
    }
    If the SCHED_YIELD flag of prev->policy is set, prev has voluntarily relinquished the CPU by issuing a sched_ yield( ) system call. In this case, the function assigns a zero goodness to it.

    Now schedule( ) repeatedly invokes the goodness( ) function on the runnable processes to determine the best candidate:

    p = init_task.next_run;
    while (p != &init_task) {
    weight = goodness(prev, p);
    if (weight > c) {
    c = weight;
    next = p;
    }
    p = p->next_run;
    }
    The while loop selects the first process in the runqueue having maximum weight. If the previous process is runnable, it is preferred with respect to other runnable processes having the same weight.

    Notice that if the runqueue list is empty (no runnable process exists except for swapper), the cycle is not entered and next points to init_task. Moreover, if all processes in the runqueue list have a priority lesser than or equal to the priority of prev, no process switch will take place and the old process will continue to be executed.

    A further check must be made at the exit of the loop to determine whether c is 0. This occurs only when all the processes in the runqueue list have exhausted their quantum, that is, all of them have a zero counter field. When this happens, a new epoch begins, therefore schedule( ) assigns to all existing processes (not only to the TASK_RUNNING ones) a fresh quantum, whose duration is the sum of the priority value plus half the counter value:

    if (!c) {
    for_each_task(p)
    p->counter = (p->counter >> 1) + p->priority;
    }

    In this way, suspended or stopped processes have their dynamic priorities periodically increased. As stated earlier, the rationale for increasing the counter value of suspended or stopped processes is to give preference to I/O-bound processes. However, even after an infinite number of increases, the value of counter can never become larger than twice[3] the priority value.

    Now comes the concluding part of schedule( ): if a process other than prev has been selected, a process switch must take place. Before performing it, however, the context_swtch field of kstat is increased by 1 to update the statistics maintained by the kernel:

    if (prev != next) {
    kstat.context_swtch++;
    switch_to(prev,next);
    }
    return;

    Notice that the return statement that exits from schedule( ) will not be performed right away by the next process but at a later time by the prev one when the scheduler selects it again for execution.

    How Good Is a Runnable Process?
    The heart of the scheduling algorithm includes identifying the best candidate among all processes in the runqueue list. This is what the goodness( ) function does. It receives as input parameters prev (the descriptor pointer of the previously running process) and p (the descriptor pointer of the process to evaluate). The integer value c returned by goodness( ) measures the "goodness" of p and has the following meanings:

    c = -1000
    p must never be selected; this value is returned when the runqueue list contains only init_task.

    c = 0
    p has exhausted its quantum. Unless p is the first process in the runqueue list and all runnable processes have also exhausted their quantum, it will not be selected for execution.

    0 < c < 1000
    p is a conventional process that has not exhausted its quantum; a higher value of c denotes a higher level of goodness.

    c >= 1000
    p is a real-time process; a higher value of c denotes a higher level of goodness.

    The goodness( ) function is equivalent to:

    if (p->policy != SCHED_OTHER)
    return 1000 + p->rt_priority;
    if (p->counter == 0)
    return 0;
    if (p->mm == prev->mm)
    return p->counter + p->priority + 1;
    return p->counter + p->priority;
    If the process is real-time, its goodness is set to at least 1000. If it is a conventional process that has exhausted its quantum, its goodness is set to 0; otherwise, it is set to p->counter + p->priority.

    A small bonus is given to p if it shares the address space with prev (i.e., if their process descriptors' mm fields point to the same memory descriptor). The rationale for this bonus is that if p runs right after prev, it will use the same page tables, hence the same memory; some of the valuable data may still be in the hardware cache.

    The Linux/SMP Scheduler
    The Linux scheduler must be slightly modified in order to support the symmetric multiprocessor (SMP) architecture. Actually, each processor runs the schedule( ) function on its own, but processors must exchange information in order to boost system performance.

    When the scheduler computes the goodness of a runnable process, it should consider whether that process was previously running on the same CPU or on another one. A process that was running on the same CPU is always preferred, since the hardware cache of the CPU could still include useful data. This rule helps in reducing the number of cache misses.

    Let us suppose, however, that CPU 1 is running a process when a second, higher-priority process that was last running on CPU 2 becomes runnable. Now the kernel is faced with an interesting dilemma: should it immediately execute the higher-priority process on CPU 1, or should it defer that process's execution until CPU 2 becomes available? In the former case, hardware caches contents are discarded; in the latter case, parallelism of the SMP architecture may not be fully exploited when CPU 2 is running the idle process (swapper).

    In order to achieve good system performance, Linux/SMP adopts an empirical rule to solve the dilemma. The adopted choice is always a compromise, and the trade-off mainly depends on the size of the hardware caches integrated into each CPU: the larger the CPU cache is, the more convenient it is to keep a process bound on that CPU.

    Linux/SMP scheduler data structures
    An aligned_data table includes one data structure for each processor, which is used mainly to obtain the descriptors of current processes quickly. Each element is filled by every invocation of the schedule( ) function and has the following structure:

    struct schedule_data {
    struct task_struct * curr;
    unsigned long last_schedule;
    };

    The curr field points to the descriptor of the process running on the corresponding CPU, while last_schedule specifies when schedule( ) selected curr as the running process.

    Several SMP-related fields are included in the process descriptor. In particular, the avg_slice field keeps track of the average quantum duration of the process, and the processor field stores the logical identifier of the last CPU that executed it.

    The cacheflush_time variable contains a rough estimate of the minimal number of CPU cycles it takes to entirely overwrite the hardware cache content. It is initialized by the smp_tune_scheduling( ) function.

    Intel Pentium processors have a hardware cache of 8 KB, so their cacheflush_time is initialized to a few hundred CPU cycles, that is, a few microseconds. Recent Intel processors have larger hardware caches, and therefore the minimal cache flush time could range from 50 to 100 microseconds.

    If cacheflush_time is greater than the average time slice of some currently running process, no process preemption is performed because it is convenient in this case to bind processes to the processors that last executed them.

    The schedule( ) function
    When the schedule( ) function is executed on an SMP system, it carries out the following operations:

    1. Performs the initial part of schedule( ) as usual.

    2. Stores the logical identifier of the executing processor in the this_cpu local variable; such value is read from the processor field of prev (that is, of the process to be replaced).

    3. Initializes the sched_data local variable so that it points to the schedule_data structure of the this_cpu CPU.

    4. Invokes goodness( ) repeatedly to select the new process to be executed; this function also examines the processor field of the processes and gives a consistent bonus (PROC_CHANGE_PENALTY, usually 15) to the process that was last executed on the this_cpu CPU.

    5. If needed, recomputes process dynamic priorities as usual.

    6. Sets sched_data->curr to next.

    7. Sets next->has_cpu to 1 and next->processor to this_cpu.

    8. Stores the current Time Stamp Counter value in the t local variable.

    9. Stores the last time slice duration of prev in the this_slice local variable; this value is the difference between t and sched_data->last_schedule.

    10. Sets sched_data->last_schedule to t.

    11. Sets the avg_slice field of prev to (prev->avg_slice+this_slice)/2; in other words, updates the average.

    12. Performs the context switch.

    13. When the kernel returns here, the original previous process has been selected again by the scheduler; the prev local variable now refers to the process that has just been replaced. If prev is still runnable and it is not the idle task of this CPU, invokes the reschedule_idle( ) function on it (see the next section).

    14. Sets the has_cpu field of prev to 0.

    The reschedule_idle( ) function
    The reschedule_idle( ) function is invoked when a process p becomes runnable. On an SMP system, the function determines whether the process should preempt the current process of some CPU. It performs the following operations:

    1. If p is a real-time process, always attempts to perform preemption: go to step 3.

    2. Returns immediately (does not attempt to preempt) if there is a CPU whose current process satisfies both of the following conditions:[4]

  • cacheflush_time is greater than the average time slice of the current process. If this is true, the process is not dirtying the cache significantly.
  • Both p and the current process need the global kernel lock in order to access some critical kernel data structure. This check is performed because replacing a process holding the lock with another one that needs it is not fruitful.

    3. If the p->processor CPU (the one on which p was last running) is idle, selects it.

    4. Otherwise, computes the difference:

    goodness(tsk, p) - goodness(tsk, tsk)

    for each task tsk running on some CPU and selects the CPU for which the difference is greatest, provided it is a positive value.

    5. If CPU has been selected, sets the need_resched field of the corresponding running process and sends a "reschedule" message to that processor.

    Performance of the Scheduling Algorithm
    The scheduling algorithm of Linux is both self-contained and relatively easy to follow. For that reason, many kernel hackers love to try to make improvements. However, the scheduler is a rather mysterious component of the kernel. While you can change its performance significantly by modifying just a few key parameters, there is usually no theoretical support to justify the results obtained. Furthermore, you can't be sure that the positive (or negative) results obtained will continue to hold when the mix of requests submitted by the users (real-time, interactive, I/O-bound, background, etc.) varies significantly. Actually, for almost every proposed scheduling strategy, it is possible to derive an artificial mix of requests that yields poor system performances.

    Let us try to outline some pitfalls of the Linux scheduler. As it will turn out, some of these limitations become significant on large systems with many users. On a single workstation that is running a few tens of processes at a time, the Linux scheduler is quite efficient. Since Linux was born on an Intel 80386 and continues to be most popular in the PC world, we consider the current Linux scheduler quite appropriate.

    The algorithm does not scale well
    If the number of existing processes is very large, it is inefficient to recompute all dynamic priorities at once.

    In old traditional Unix kernels, the dynamic priorities were recomputed every second, thus the problem was even worse. Linux tries instead to minimize the overhead of the scheduler. Priorities are recomputed only when all runnable processes have exhausted their time quantum. Therefore, when the number of processes is large, the recomputation phase is more expensive but is executed less frequently.

    This simple approach has the disadvantage that when the number of runnable processes is very large, I/O-bound processes are seldom boosted, and therefore interactive applications have a longer response time.

    The predefined quantum is too large for high system loads
    The system responsiveness experienced by users depends heavily on the system load, which is the average number of processes that are runnable, and hence waiting for CPU time.[5]

    As mentioned before, system responsiveness depends also on the average time-quantum duration of the runnable processes. In Linux, the predefined time quantum appears to be too large for high-end machines having a very high expected system load.

    I/O-bound process boosting strategy is not optimal
    The preference for I/O-bound processes is a good strategy to ensure a short response time for interactive programs, but it is not perfect. Indeed, some batch programs with almost no user interaction are I/O-bound. For instance, consider a database search engine that must typically read lots of data from the hard disk or a network application that must collect data from a remote host on a slow link. Even if these kinds of processes do not need a short response time, they are boosted by the scheduling algorithm. On the other hand, interactive programs that are also CPU-bound may appear unresponsive to the users, since the increment of dynamic priority due to I/O blocking operations may not compensate for the decrement due to CPU usage.

    Support for real-time applications is weak
    Nonpreemptive kernels are not well suited for real-time applications, since processes may spend several milliseconds in Kernel Mode while handling an interrupt or exception. During this time, a real-time process that becomes runnable cannot be resumed. This is unacceptable for real-time applications, which require predictable and low response times.

    Future versions of Linux will likely address this problem, either by implementing SVR4's "fixed preemption points" or by making the kernel fully preemptive.

    However, kernel preemption is just one of several necessary conditions for implementing an effective real-time scheduler. Several other issues must be considered. For instance, real-time processes often must use resources also needed by conventional processes. A real-time process may thus end up waiting until a lower-priority process releases some resource. This phenomenon is called priority inversion. Moreover, a real-time process could require a kernel service that is granted on behalf of another lower-priority process (for example, a kernel thread). This phenomenon is called hidden scheduling. An effective real-time scheduler should address and resolve such problems.

    In the next article we'll describe the system calls that affect process scheduling.

  • More Stories By Daniel Bovet

    Daniel P. Bovet got a Ph.D. in computer science at UCLA in 1968 and is now full Professor at the University of Rome, "Tor Vergata," Italy.

    More Stories By Marco Cesati

    Marco Cesati got a degree in mathematics in 1992 and a Ph.D. in computer science (University of Rome, "La Sapienza") in 1995. He is now a research assistant in the computer science department of the School of Engineering (University of Rome, "Tor Vergata").

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    SYS-CON Events announced today that Tufin, the market-leading provider of Security Policy Orchestration Solutions, will exhibit at SYS-CON's 16th International Cloud Expo®, which will take place on June 9-11, 2015, at the Javits Center in New York City, NY. As the market leader of Security Policy Orchestration, Tufin automates and accelerates network configuration changes while maintaining security and compliance. Tufin's award-winning Orchestration Suite™ gives IT organizations the power and agility to enforce security policy across complex, multi-vendor enterprise networks. With more than 1...
    As enterprises move to all-IP networks and cloud-based applications, communications service providers (CSPs) – facing increased competition from over-the-top providers delivering content via the Internet and independently of CSPs – must be able to offer seamless cloud-based communication and collaboration solutions that can scale for small, midsize, and large enterprises, as well as public sector organizations, in order to keep and grow market share. The latest version of Oracle Communications Unified Communications Suite gives CSPs the capability to do just that. In addition, its integration ...
    SYS-CON Events announced today that Cloudian, Inc., the leading provider of hybrid cloud storage solutions, will exhibit at SYS-CON's 16th International Cloud Expo®, which will take place on June 9-11, 2015, at the Javits Center in New York City, NY. Cloudian, Inc., is a Foster City, California - based software company specializing in cloud storage software. The main product is Cloudian, an Amazon S3-compliant cloud object storage platform, the bedrock of cloud computing systems, that enables cloud service providers and enterprises to build reliable, affordable and scalable cloud storage solu...
    SYS-CON Events announced today that Gridstore™, the leader in hyper-converged infrastructure purpose-built to optimize Microsoft workloads, will exhibit at SYS-CON's 16th International Cloud Expo®, which will take place on June 9-11, 2015, at the Javits Center in New York City, NY. Gridstore™ is the leader in hyper-converged infrastructure purpose-built for Microsoft workloads and designed to accelerate applications in virtualized environments. Gridstore’s hyper-converged infrastructure is the industry’s first all flash version of HyperConverged Appliances that include both compute and storag...
    SYS-CON Events announced today that IDenticard will exhibit at SYS-CON's 16th International Cloud Expo®, which will take place on June 9-11, 2015, at the Javits Center in New York City, NY. IDenticard™ is the security division of Brady Corp (NYSE: BRC), a $1.5 billion manufacturer of identification products. We have small-company values with the strength and stability of a major corporation. IDenticard offers local sales, support and service to our customers across the United States and Canada. Our partner network encompasses some 300 of the world's leading systems integrators and security s...
    BroadSoft on Tuesday announced that it is a recipient of the 2014 Frost & Sullivan Market Leadership Award in the Hosted/Cloud Internet Protocol (IP) Telephony market for Latin America. According to Frost & Sullivan market research, the Latin America (LATAM) hosted/cloud Internet Protocol (IP) telephony market, including integrated unified communications and collaboration (UC&C) applications, is currently experiencing a rapid growth trajectory and is expected to exhibit a tenfold rise in annual revenues in the 2013-2020 period. With more than 600 cloud deployments internationally, BroadSoft w...