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diff --git a/Documentation/scheduler/sched-deadline.rst b/Documentation/scheduler/sched-deadline.rst new file mode 100644 index 000000000000..14a2f7bf63fe --- /dev/null +++ b/Documentation/scheduler/sched-deadline.rst @@ -0,0 +1,888 @@ +======================== +Deadline Task Scheduling +======================== + +.. CONTENTS + + 0. WARNING + 1. Overview + 2. Scheduling algorithm + 2.1 Main algorithm + 2.2 Bandwidth reclaiming + 3. Scheduling Real-Time Tasks + 3.1 Definitions + 3.2 Schedulability Analysis for Uniprocessor Systems + 3.3 Schedulability Analysis for Multiprocessor Systems + 3.4 Relationship with SCHED_DEADLINE Parameters + 4. Bandwidth management + 4.1 System-wide settings + 4.2 Task interface + 4.3 Default behavior + 4.4 Behavior of sched_yield() + 5. Tasks CPU affinity + 5.1 SCHED_DEADLINE and cpusets HOWTO + 6. Future plans + A. Test suite + B. Minimal main() + + +0. WARNING +========== + + Fiddling with these settings can result in an unpredictable or even unstable + system behavior. As for -rt (group) scheduling, it is assumed that root users + know what they're doing. + + +1. Overview +=========== + + The SCHED_DEADLINE policy contained inside the sched_dl scheduling class is + basically an implementation of the Earliest Deadline First (EDF) scheduling + algorithm, augmented with a mechanism (called Constant Bandwidth Server, CBS) + that makes it possible to isolate the behavior of tasks between each other. + + +2. Scheduling algorithm +======================= + +2.1 Main algorithm +------------------ + + SCHED_DEADLINE [18] uses three parameters, named "runtime", "period", and + "deadline", to schedule tasks. A SCHED_DEADLINE task should receive + "runtime" microseconds of execution time every "period" microseconds, and + these "runtime" microseconds are available within "deadline" microseconds + from the beginning of the period. In order to implement this behavior, + every time the task wakes up, the scheduler computes a "scheduling deadline" + consistent with the guarantee (using the CBS[2,3] algorithm). Tasks are then + scheduled using EDF[1] on these scheduling deadlines (the task with the + earliest scheduling deadline is selected for execution). Notice that the + task actually receives "runtime" time units within "deadline" if a proper + "admission control" strategy (see Section "4. Bandwidth management") is used + (clearly, if the system is overloaded this guarantee cannot be respected). + + Summing up, the CBS[2,3] algorithm assigns scheduling deadlines to tasks so + that each task runs for at most its runtime every period, avoiding any + interference between different tasks (bandwidth isolation), while the EDF[1] + algorithm selects the task with the earliest scheduling deadline as the one + to be executed next. Thanks to this feature, tasks that do not strictly comply + with the "traditional" real-time task model (see Section 3) can effectively + use the new policy. + + In more details, the CBS algorithm assigns scheduling deadlines to + tasks in the following way: + + - Each SCHED_DEADLINE task is characterized by the "runtime", + "deadline", and "period" parameters; + + - The state of the task is described by a "scheduling deadline", and + a "remaining runtime". These two parameters are initially set to 0; + + - When a SCHED_DEADLINE task wakes up (becomes ready for execution), + the scheduler checks if:: + + remaining runtime runtime + ---------------------------------- > --------- + scheduling deadline - current time period + + then, if the scheduling deadline is smaller than the current time, or + this condition is verified, the scheduling deadline and the + remaining runtime are re-initialized as + + scheduling deadline = current time + deadline + remaining runtime = runtime + + otherwise, the scheduling deadline and the remaining runtime are + left unchanged; + + - When a SCHED_DEADLINE task executes for an amount of time t, its + remaining runtime is decreased as:: + + remaining runtime = remaining runtime - t + + (technically, the runtime is decreased at every tick, or when the + task is descheduled / preempted); + + - When the remaining runtime becomes less or equal than 0, the task is + said to be "throttled" (also known as "depleted" in real-time literature) + and cannot be scheduled until its scheduling deadline. The "replenishment + time" for this task (see next item) is set to be equal to the current + value of the scheduling deadline; + + - When the current time is equal to the replenishment time of a + throttled task, the scheduling deadline and the remaining runtime are + updated as:: + + scheduling deadline = scheduling deadline + period + remaining runtime = remaining runtime + runtime + + The SCHED_FLAG_DL_OVERRUN flag in sched_attr's sched_flags field allows a task + to get informed about runtime overruns through the delivery of SIGXCPU + signals. + + +2.2 Bandwidth reclaiming +------------------------ + + Bandwidth reclaiming for deadline tasks is based on the GRUB (Greedy + Reclamation of Unused Bandwidth) algorithm [15, 16, 17] and it is enabled + when flag SCHED_FLAG_RECLAIM is set. + + The following diagram illustrates the state names for tasks handled by GRUB:: + + ------------ + (d) | Active | + ------------->| | + | | Contending | + | ------------ + | A | + ---------- | | + | | | | + | Inactive | |(b) | (a) + | | | | + ---------- | | + A | V + | ------------ + | | Active | + --------------| Non | + (c) | Contending | + ------------ + + A task can be in one of the following states: + + - ActiveContending: if it is ready for execution (or executing); + + - ActiveNonContending: if it just blocked and has not yet surpassed the 0-lag + time; + + - Inactive: if it is blocked and has surpassed the 0-lag time. + + State transitions: + + (a) When a task blocks, it does not become immediately inactive since its + bandwidth cannot be immediately reclaimed without breaking the + real-time guarantees. It therefore enters a transitional state called + ActiveNonContending. The scheduler arms the "inactive timer" to fire at + the 0-lag time, when the task's bandwidth can be reclaimed without + breaking the real-time guarantees. + + The 0-lag time for a task entering the ActiveNonContending state is + computed as:: + + (runtime * dl_period) + deadline - --------------------- + dl_runtime + + where runtime is the remaining runtime, while dl_runtime and dl_period + are the reservation parameters. + + (b) If the task wakes up before the inactive timer fires, the task re-enters + the ActiveContending state and the "inactive timer" is canceled. + In addition, if the task wakes up on a different runqueue, then + the task's utilization must be removed from the previous runqueue's active + utilization and must be added to the new runqueue's active utilization. + In order to avoid races between a task waking up on a runqueue while the + "inactive timer" is running on a different CPU, the "dl_non_contending" + flag is used to indicate that a task is not on a runqueue but is active + (so, the flag is set when the task blocks and is cleared when the + "inactive timer" fires or when the task wakes up). + + (c) When the "inactive timer" fires, the task enters the Inactive state and + its utilization is removed from the runqueue's active utilization. + + (d) When an inactive task wakes up, it enters the ActiveContending state and + its utilization is added to the active utilization of the runqueue where + it has been enqueued. + + For each runqueue, the algorithm GRUB keeps track of two different bandwidths: + + - Active bandwidth (running_bw): this is the sum of the bandwidths of all + tasks in active state (i.e., ActiveContending or ActiveNonContending); + + - Total bandwidth (this_bw): this is the sum of all tasks "belonging" to the + runqueue, including the tasks in Inactive state. + + + The algorithm reclaims the bandwidth of the tasks in Inactive state. + It does so by decrementing the runtime of the executing task Ti at a pace equal + to + + dq = -max{ Ui / Umax, (1 - Uinact - Uextra) } dt + + where: + + - Ui is the bandwidth of task Ti; + - Umax is the maximum reclaimable utilization (subjected to RT throttling + limits); + - Uinact is the (per runqueue) inactive utilization, computed as + (this_bq - running_bw); + - Uextra is the (per runqueue) extra reclaimable utilization + (subjected to RT throttling limits). + + + Let's now see a trivial example of two deadline tasks with runtime equal + to 4 and period equal to 8 (i.e., bandwidth equal to 0.5):: + + A Task T1 + | + | | + | | + |-------- |---- + | | V + |---|---|---|---|---|---|---|---|--------->t + 0 1 2 3 4 5 6 7 8 + + + A Task T2 + | + | | + | | + | ------------------------| + | | V + |---|---|---|---|---|---|---|---|--------->t + 0 1 2 3 4 5 6 7 8 + + + A running_bw + | + 1 ----------------- ------ + | | | + 0.5- ----------------- + | | + |---|---|---|---|---|---|---|---|--------->t + 0 1 2 3 4 5 6 7 8 + + + - Time t = 0: + + Both tasks are ready for execution and therefore in ActiveContending state. + Suppose Task T1 is the first task to start execution. + Since there are no inactive tasks, its runtime is decreased as dq = -1 dt. + + - Time t = 2: + + Suppose that task T1 blocks + Task T1 therefore enters the ActiveNonContending state. Since its remaining + runtime is equal to 2, its 0-lag time is equal to t = 4. + Task T2 start execution, with runtime still decreased as dq = -1 dt since + there are no inactive tasks. + + - Time t = 4: + + This is the 0-lag time for Task T1. Since it didn't woken up in the + meantime, it enters the Inactive state. Its bandwidth is removed from + running_bw. + Task T2 continues its execution. However, its runtime is now decreased as + dq = - 0.5 dt because Uinact = 0.5. + Task T2 therefore reclaims the bandwidth unused by Task T1. + + - Time t = 8: + + Task T1 wakes up. It enters the ActiveContending state again, and the + running_bw is incremented. + + +2.3 Energy-aware scheduling +--------------------------- + + When cpufreq's schedutil governor is selected, SCHED_DEADLINE implements the + GRUB-PA [19] algorithm, reducing the CPU operating frequency to the minimum + value that still allows to meet the deadlines. This behavior is currently + implemented only for ARM architectures. + + A particular care must be taken in case the time needed for changing frequency + is of the same order of magnitude of the reservation period. In such cases, + setting a fixed CPU frequency results in a lower amount of deadline misses. + + +3. Scheduling Real-Time Tasks +============================= + + + + .. BIG FAT WARNING ****************************************************** + + .. warning:: + + This section contains a (not-thorough) summary on classical deadline + scheduling theory, and how it applies to SCHED_DEADLINE. + The reader can "safely" skip to Section 4 if only interested in seeing + how the scheduling policy can be used. Anyway, we strongly recommend + to come back here and continue reading (once the urge for testing is + satisfied :P) to be sure of fully understanding all technical details. + + .. ************************************************************************ + + There are no limitations on what kind of task can exploit this new + scheduling discipline, even if it must be said that it is particularly + suited for periodic or sporadic real-time tasks that need guarantees on their + timing behavior, e.g., multimedia, streaming, control applications, etc. + +3.1 Definitions +------------------------ + + A typical real-time task is composed of a repetition of computation phases + (task instances, or jobs) which are activated on a periodic or sporadic + fashion. + Each job J_j (where J_j is the j^th job of the task) is characterized by an + arrival time r_j (the time when the job starts), an amount of computation + time c_j needed to finish the job, and a job absolute deadline d_j, which + is the time within which the job should be finished. The maximum execution + time max{c_j} is called "Worst Case Execution Time" (WCET) for the task. + A real-time task can be periodic with period P if r_{j+1} = r_j + P, or + sporadic with minimum inter-arrival time P is r_{j+1} >= r_j + P. Finally, + d_j = r_j + D, where D is the task's relative deadline. + Summing up, a real-time task can be described as + + Task = (WCET, D, P) + + The utilization of a real-time task is defined as the ratio between its + WCET and its period (or minimum inter-arrival time), and represents + the fraction of CPU time needed to execute the task. + + If the total utilization U=sum(WCET_i/P_i) is larger than M (with M equal + to the number of CPUs), then the scheduler is unable to respect all the + deadlines. + Note that total utilization is defined as the sum of the utilizations + WCET_i/P_i over all the real-time tasks in the system. When considering + multiple real-time tasks, the parameters of the i-th task are indicated + with the "_i" suffix. + Moreover, if the total utilization is larger than M, then we risk starving + non- real-time tasks by real-time tasks. + If, instead, the total utilization is smaller than M, then non real-time + tasks will not be starved and the system might be able to respect all the + deadlines. + As a matter of fact, in this case it is possible to provide an upper bound + for tardiness (defined as the maximum between 0 and the difference + between the finishing time of a job and its absolute deadline). + More precisely, it can be proven that using a global EDF scheduler the + maximum tardiness of each task is smaller or equal than + + ((M − 1) · WCET_max − WCET_min)/(M − (M − 2) · U_max) + WCET_max + + where WCET_max = max{WCET_i} is the maximum WCET, WCET_min=min{WCET_i} + is the minimum WCET, and U_max = max{WCET_i/P_i} is the maximum + utilization[12]. + +3.2 Schedulability Analysis for Uniprocessor Systems +---------------------------------------------------- + + If M=1 (uniprocessor system), or in case of partitioned scheduling (each + real-time task is statically assigned to one and only one CPU), it is + possible to formally check if all the deadlines are respected. + If D_i = P_i for all tasks, then EDF is able to respect all the deadlines + of all the tasks executing on a CPU if and only if the total utilization + of the tasks running on such a CPU is smaller or equal than 1. + If D_i != P_i for some task, then it is possible to define the density of + a task as WCET_i/min{D_i,P_i}, and EDF is able to respect all the deadlines + of all the tasks running on a CPU if the sum of the densities of the tasks + running on such a CPU is smaller or equal than 1: + + sum(WCET_i / min{D_i, P_i}) <= 1 + + It is important to notice that this condition is only sufficient, and not + necessary: there are task sets that are schedulable, but do not respect the + condition. For example, consider the task set {Task_1,Task_2} composed by + Task_1=(50ms,50ms,100ms) and Task_2=(10ms,100ms,100ms). + EDF is clearly able to schedule the two tasks without missing any deadline + (Task_1 is scheduled as soon as it is released, and finishes just in time + to respect its deadline; Task_2 is scheduled immediately after Task_1, hence + its response time cannot be larger than 50ms + 10ms = 60ms) even if + + 50 / min{50,100} + 10 / min{100, 100} = 50 / 50 + 10 / 100 = 1.1 + + Of course it is possible to test the exact schedulability of tasks with + D_i != P_i (checking a condition that is both sufficient and necessary), + but this cannot be done by comparing the total utilization or density with + a constant. Instead, the so called "processor demand" approach can be used, + computing the total amount of CPU time h(t) needed by all the tasks to + respect all of their deadlines in a time interval of size t, and comparing + such a time with the interval size t. If h(t) is smaller than t (that is, + the amount of time needed by the tasks in a time interval of size t is + smaller than the size of the interval) for all the possible values of t, then + EDF is able to schedule the tasks respecting all of their deadlines. Since + performing this check for all possible values of t is impossible, it has been + proven[4,5,6] that it is sufficient to perform the test for values of t + between 0 and a maximum value L. The cited papers contain all of the + mathematical details and explain how to compute h(t) and L. + In any case, this kind of analysis is too complex as well as too + time-consuming to be performed on-line. Hence, as explained in Section + 4 Linux uses an admission test based on the tasks' utilizations. + +3.3 Schedulability Analysis for Multiprocessor Systems +------------------------------------------------------ + + On multiprocessor systems with global EDF scheduling (non partitioned + systems), a sufficient test for schedulability can not be based on the + utilizations or densities: it can be shown that even if D_i = P_i task + sets with utilizations slightly larger than 1 can miss deadlines regardless + of the number of CPUs. + + Consider a set {Task_1,...Task_{M+1}} of M+1 tasks on a system with M + CPUs, with the first task Task_1=(P,P,P) having period, relative deadline + and WCET equal to P. The remaining M tasks Task_i=(e,P-1,P-1) have an + arbitrarily small worst case execution time (indicated as "e" here) and a + period smaller than the one of the first task. Hence, if all the tasks + activate at the same time t, global EDF schedules these M tasks first + (because their absolute deadlines are equal to t + P - 1, hence they are + smaller than the absolute deadline of Task_1, which is t + P). As a + result, Task_1 can be scheduled only at time t + e, and will finish at + time t + e + P, after its absolute deadline. The total utilization of the + task set is U = M · e / (P - 1) + P / P = M · e / (P - 1) + 1, and for small + values of e this can become very close to 1. This is known as "Dhall's + effect"[7]. Note: the example in the original paper by Dhall has been + slightly simplified here (for example, Dhall more correctly computed + lim_{e->0}U). + + More complex schedulability tests for global EDF have been developed in + real-time literature[8,9], but they are not based on a simple comparison + between total utilization (or density) and a fixed constant. If all tasks + have D_i = P_i, a sufficient schedulability condition can be expressed in + a simple way: + + sum(WCET_i / P_i) <= M - (M - 1) · U_max + + where U_max = max{WCET_i / P_i}[10]. Notice that for U_max = 1, + M - (M - 1) · U_max becomes M - M + 1 = 1 and this schedulability condition + just confirms the Dhall's effect. A more complete survey of the literature + about schedulability tests for multi-processor real-time scheduling can be + found in [11]. + + As seen, enforcing that the total utilization is smaller than M does not + guarantee that global EDF schedules the tasks without missing any deadline + (in other words, global EDF is not an optimal scheduling algorithm). However, + a total utilization smaller than M is enough to guarantee that non real-time + tasks are not starved and that the tardiness of real-time tasks has an upper + bound[12] (as previously noted). Different bounds on the maximum tardiness + experienced by real-time tasks have been developed in various papers[13,14], + but the theoretical result that is important for SCHED_DEADLINE is that if + the total utilization is smaller or equal than M then the response times of + the tasks are limited. + +3.4 Relationship with SCHED_DEADLINE Parameters +----------------------------------------------- + + Finally, it is important to understand the relationship between the + SCHED_DEADLINE scheduling parameters described in Section 2 (runtime, + deadline and period) and the real-time task parameters (WCET, D, P) + described in this section. Note that the tasks' temporal constraints are + represented by its absolute deadlines d_j = r_j + D described above, while + SCHED_DEADLINE schedules the tasks according to scheduling deadlines (see + Section 2). + If an admission test is used to guarantee that the scheduling deadlines + are respected, then SCHED_DEADLINE can be used to schedule real-time tasks + guaranteeing that all the jobs' deadlines of a task are respected. + In order to do this, a task must be scheduled by setting: + + - runtime >= WCET + - deadline = D + - period <= P + + IOW, if runtime >= WCET and if period is <= P, then the scheduling deadlines + and the absolute deadlines (d_j) coincide, so a proper admission control + allows to respect the jobs' absolute deadlines for this task (this is what is + called "hard schedulability property" and is an extension of Lemma 1 of [2]). + Notice that if runtime > deadline the admission control will surely reject + this task, as it is not possible to respect its temporal constraints. + + References: + + 1 - C. L. Liu and J. W. Layland. Scheduling algorithms for multiprogram- + ming in a hard-real-time environment. Journal of the Association for + Computing Machinery, 20(1), 1973. + 2 - L. Abeni , G. Buttazzo. Integrating Multimedia Applications in Hard + Real-Time Systems. Proceedings of the 19th IEEE Real-time Systems + Symposium, 1998. http://retis.sssup.it/~giorgio/paps/1998/rtss98-cbs.pdf + 3 - L. Abeni. Server Mechanisms for Multimedia Applications. ReTiS Lab + Technical Report. http://disi.unitn.it/~abeni/tr-98-01.pdf + 4 - J. Y. Leung and M.L. Merril. A Note on Preemptive Scheduling of + Periodic, Real-Time Tasks. Information Processing Letters, vol. 11, + no. 3, pp. 115-118, 1980. + 5 - S. K. Baruah, A. K. Mok and L. E. Rosier. Preemptively Scheduling + Hard-Real-Time Sporadic Tasks on One Processor. Proceedings of the + 11th IEEE Real-time Systems Symposium, 1990. + 6 - S. K. Baruah, L. E. Rosier and R. R. Howell. Algorithms and Complexity + Concerning the Preemptive Scheduling of Periodic Real-Time tasks on + One Processor. Real-Time Systems Journal, vol. 4, no. 2, pp 301-324, + 1990. + 7 - S. J. Dhall and C. L. Liu. On a real-time scheduling problem. Operations + research, vol. 26, no. 1, pp 127-140, 1978. + 8 - T. Baker. Multiprocessor EDF and Deadline Monotonic Schedulability + Analysis. Proceedings of the 24th IEEE Real-Time Systems Symposium, 2003. + 9 - T. Baker. An Analysis of EDF Schedulability on a Multiprocessor. + IEEE Transactions on Parallel and Distributed Systems, vol. 16, no. 8, + pp 760-768, 2005. + 10 - J. Goossens, S. Funk and S. Baruah, Priority-Driven Scheduling of + Periodic Task Systems on Multiprocessors. Real-Time Systems Journal, + vol. 25, no. 2–3, pp. 187–205, 2003. + 11 - R. Davis and A. Burns. A Survey of Hard Real-Time Scheduling for + Multiprocessor Systems. ACM Computing Surveys, vol. 43, no. 4, 2011. + http://www-users.cs.york.ac.uk/~robdavis/papers/MPSurveyv5.0.pdf + 12 - U. C. Devi and J. H. Anderson. Tardiness Bounds under Global EDF + Scheduling on a Multiprocessor. Real-Time Systems Journal, vol. 32, + no. 2, pp 133-189, 2008. + 13 - P. Valente and G. Lipari. An Upper Bound to the Lateness of Soft + Real-Time Tasks Scheduled by EDF on Multiprocessors. Proceedings of + the 26th IEEE Real-Time Systems Symposium, 2005. + 14 - J. Erickson, U. Devi and S. Baruah. Improved tardiness bounds for + Global EDF. Proceedings of the 22nd Euromicro Conference on + Real-Time Systems, 2010. + 15 - G. Lipari, S. Baruah, Greedy reclamation of unused bandwidth in + constant-bandwidth servers, 12th IEEE Euromicro Conference on Real-Time + Systems, 2000. + 16 - L. Abeni, J. Lelli, C. Scordino, L. Palopoli, Greedy CPU reclaiming for + SCHED DEADLINE. In Proceedings of the Real-Time Linux Workshop (RTLWS), + Dusseldorf, Germany, 2014. + 17 - L. Abeni, G. Lipari, A. Parri, Y. Sun, Multicore CPU reclaiming: parallel + or sequential?. In Proceedings of the 31st Annual ACM Symposium on Applied + Computing, 2016. + 18 - J. Lelli, C. Scordino, L. Abeni, D. Faggioli, Deadline scheduling in the + Linux kernel, Software: Practice and Experience, 46(6): 821-839, June + 2016. + 19 - C. Scordino, L. Abeni, J. Lelli, Energy-Aware Real-Time Scheduling in + the Linux Kernel, 33rd ACM/SIGAPP Symposium On Applied Computing (SAC + 2018), Pau, France, April 2018. + + +4. Bandwidth management +======================= + + As previously mentioned, in order for -deadline scheduling to be + effective and useful (that is, to be able to provide "runtime" time units + within "deadline"), it is important to have some method to keep the allocation + of the available fractions of CPU time to the various tasks under control. + This is usually called "admission control" and if it is not performed, then + no guarantee can be given on the actual scheduling of the -deadline tasks. + + As already stated in Section 3, a necessary condition to be respected to + correctly schedule a set of real-time tasks is that the total utilization + is smaller than M. When talking about -deadline tasks, this requires that + the sum of the ratio between runtime and period for all tasks is smaller + than M. Notice that the ratio runtime/period is equivalent to the utilization + of a "traditional" real-time task, and is also often referred to as + "bandwidth". + The interface used to control the CPU bandwidth that can be allocated + to -deadline tasks is similar to the one already used for -rt + tasks with real-time group scheduling (a.k.a. RT-throttling - see + Documentation/scheduler/sched-rt-group.rst), and is based on readable/ + writable control files located in procfs (for system wide settings). + Notice that per-group settings (controlled through cgroupfs) are still not + defined for -deadline tasks, because more discussion is needed in order to + figure out how we want to manage SCHED_DEADLINE bandwidth at the task group + level. + + A main difference between deadline bandwidth management and RT-throttling + is that -deadline tasks have bandwidth on their own (while -rt ones don't!), + and thus we don't need a higher level throttling mechanism to enforce the + desired bandwidth. In other words, this means that interface parameters are + only used at admission control time (i.e., when the user calls + sched_setattr()). Scheduling is then performed considering actual tasks' + parameters, so that CPU bandwidth is allocated to SCHED_DEADLINE tasks + respecting their needs in terms of granularity. Therefore, using this simple + interface we can put a cap on total utilization of -deadline tasks (i.e., + \Sum (runtime_i / period_i) < global_dl_utilization_cap). + +4.1 System wide settings +------------------------ + + The system wide settings are configured under the /proc virtual file system. + + For now the -rt knobs are used for -deadline admission control and the + -deadline runtime is accounted against the -rt runtime. We realize that this + isn't entirely desirable; however, it is better to have a small interface for + now, and be able to change it easily later. The ideal situation (see 5.) is to + run -rt tasks from a -deadline server; in which case the -rt bandwidth is a + direct subset of dl_bw. + + This means that, for a root_domain comprising M CPUs, -deadline tasks + can be created while the sum of their bandwidths stays below: + + M * (sched_rt_runtime_us / sched_rt_period_us) + + It is also possible to disable this bandwidth management logic, and + be thus free of oversubscribing the system up to any arbitrary level. + This is done by writing -1 in /proc/sys/kernel/sched_rt_runtime_us. + + +4.2 Task interface +------------------ + + Specifying a periodic/sporadic task that executes for a given amount of + runtime at each instance, and that is scheduled according to the urgency of + its own timing constraints needs, in general, a way of declaring: + + - a (maximum/typical) instance execution time, + - a minimum interval between consecutive instances, + - a time constraint by which each instance must be completed. + + Therefore: + + * a new struct sched_attr, containing all the necessary fields is + provided; + * the new scheduling related syscalls that manipulate it, i.e., + sched_setattr() and sched_getattr() are implemented. + + For debugging purposes, the leftover runtime and absolute deadline of a + SCHED_DEADLINE task can be retrieved through /proc/<pid>/sched (entries + dl.runtime and dl.deadline, both values in ns). A programmatic way to + retrieve these values from production code is under discussion. + + +4.3 Default behavior +--------------------- + + The default value for SCHED_DEADLINE bandwidth is to have rt_runtime equal to + 950000. With rt_period equal to 1000000, by default, it means that -deadline + tasks can use at most 95%, multiplied by the number of CPUs that compose the + root_domain, for each root_domain. + This means that non -deadline tasks will receive at least 5% of the CPU time, + and that -deadline tasks will receive their runtime with a guaranteed + worst-case delay respect to the "deadline" parameter. If "deadline" = "period" + and the cpuset mechanism is used to implement partitioned scheduling (see + Section 5), then this simple setting of the bandwidth management is able to + deterministically guarantee that -deadline tasks will receive their runtime + in a period. + + Finally, notice that in order not to jeopardize the admission control a + -deadline task cannot fork. + + +4.4 Behavior of sched_yield() +----------------------------- + + When a SCHED_DEADLINE task calls sched_yield(), it gives up its + remaining runtime and is immediately throttled, until the next + period, when its runtime will be replenished (a special flag + dl_yielded is set and used to handle correctly throttling and runtime + replenishment after a call to sched_yield()). + + This behavior of sched_yield() allows the task to wake-up exactly at + the beginning of the next period. Also, this may be useful in the + future with bandwidth reclaiming mechanisms, where sched_yield() will + make the leftoever runtime available for reclamation by other + SCHED_DEADLINE tasks. + + +5. Tasks CPU affinity +===================== + + -deadline tasks cannot have an affinity mask smaller that the entire + root_domain they are created on. However, affinities can be specified + through the cpuset facility (Documentation/admin-guide/cgroup-v1/cpusets.rst). + +5.1 SCHED_DEADLINE and cpusets HOWTO +------------------------------------ + + An example of a simple configuration (pin a -deadline task to CPU0) + follows (rt-app is used to create a -deadline task):: + + mkdir /dev/cpuset + mount -t cgroup -o cpuset cpuset /dev/cpuset + cd /dev/cpuset + mkdir cpu0 + echo 0 > cpu0/cpuset.cpus + echo 0 > cpu0/cpuset.mems + echo 1 > cpuset.cpu_exclusive + echo 0 > cpuset.sched_load_balance + echo 1 > cpu0/cpuset.cpu_exclusive + echo 1 > cpu0/cpuset.mem_exclusive + echo $$ > cpu0/tasks + rt-app -t 100000:10000:d:0 -D5 # it is now actually superfluous to specify + # task affinity + +6. Future plans +=============== + + Still missing: + + - programmatic way to retrieve current runtime and absolute deadline + - refinements to deadline inheritance, especially regarding the possibility + of retaining bandwidth isolation among non-interacting tasks. This is + being studied from both theoretical and practical points of view, and + hopefully we should be able to produce some demonstrative code soon; + - (c)group based bandwidth management, and maybe scheduling; + - access control for non-root users (and related security concerns to + address), which is the best way to allow unprivileged use of the mechanisms + and how to prevent non-root users "cheat" the system? + + As already discussed, we are planning also to merge this work with the EDF + throttling patches [https://lkml.org/lkml/2010/2/23/239] but we still are in + the preliminary phases of the merge and we really seek feedback that would + help us decide on the direction it should take. + +Appendix A. Test suite +====================== + + The SCHED_DEADLINE policy can be easily tested using two applications that + are part of a wider Linux Scheduler validation suite. The suite is + available as a GitHub repository: https://github.com/scheduler-tools. + + The first testing application is called rt-app and can be used to + start multiple threads with specific parameters. rt-app supports + SCHED_{OTHER,FIFO,RR,DEADLINE} scheduling policies and their related + parameters (e.g., niceness, priority, runtime/deadline/period). rt-app + is a valuable tool, as it can be used to synthetically recreate certain + workloads (maybe mimicking real use-cases) and evaluate how the scheduler + behaves under such workloads. In this way, results are easily reproducible. + rt-app is available at: https://github.com/scheduler-tools/rt-app. + + Thread parameters can be specified from the command line, with something like + this:: + + # rt-app -t 100000:10000:d -t 150000:20000:f:10 -D5 + + The above creates 2 threads. The first one, scheduled by SCHED_DEADLINE, + executes for 10ms every 100ms. The second one, scheduled at SCHED_FIFO + priority 10, executes for 20ms every 150ms. The test will run for a total + of 5 seconds. + + More interestingly, configurations can be described with a json file that + can be passed as input to rt-app with something like this:: + + # rt-app my_config.json + + The parameters that can be specified with the second method are a superset + of the command line options. Please refer to rt-app documentation for more + details (`<rt-app-sources>/doc/*.json`). + + The second testing application is a modification of schedtool, called + schedtool-dl, which can be used to setup SCHED_DEADLINE parameters for a + certain pid/application. schedtool-dl is available at: + https://github.com/scheduler-tools/schedtool-dl.git. + + The usage is straightforward:: + + # schedtool -E -t 10000000:100000000 -e ./my_cpuhog_app + + With this, my_cpuhog_app is put to run inside a SCHED_DEADLINE reservation + of 10ms every 100ms (note that parameters are expressed in microseconds). + You can also use schedtool to create a reservation for an already running + application, given that you know its pid:: + + # schedtool -E -t 10000000:100000000 my_app_pid + +Appendix B. Minimal main() +========================== + + We provide in what follows a simple (ugly) self-contained code snippet + showing how SCHED_DEADLINE reservations can be created by a real-time + application developer:: + + #define _GNU_SOURCE + #include <unistd.h> + #include <stdio.h> + #include <stdlib.h> + #include <string.h> + #include <time.h> + #include <linux/unistd.h> + #include <linux/kernel.h> + #include <linux/types.h> + #include <sys/syscall.h> + #include <pthread.h> + + #define gettid() syscall(__NR_gettid) + + #define SCHED_DEADLINE 6 + + /* XXX use the proper syscall numbers */ + #ifdef __x86_64__ + #define __NR_sched_setattr 314 + #define __NR_sched_getattr 315 + #endif + + #ifdef __i386__ + #define __NR_sched_setattr 351 + #define __NR_sched_getattr 352 + #endif + + #ifdef __arm__ + #define __NR_sched_setattr 380 + #define __NR_sched_getattr 381 + #endif + + static volatile int done; + + struct sched_attr { + __u32 size; + + __u32 sched_policy; + __u64 sched_flags; + + /* SCHED_NORMAL, SCHED_BATCH */ + __s32 sched_nice; + + /* SCHED_FIFO, SCHED_RR */ + __u32 sched_priority; + + /* SCHED_DEADLINE (nsec) */ + __u64 sched_runtime; + __u64 sched_deadline; + __u64 sched_period; + }; + + int sched_setattr(pid_t pid, + const struct sched_attr *attr, + unsigned int flags) + { + return syscall(__NR_sched_setattr, pid, attr, flags); + } + + int sched_getattr(pid_t pid, + struct sched_attr *attr, + unsigned int size, + unsigned int flags) + { + return syscall(__NR_sched_getattr, pid, attr, size, flags); + } + + void *run_deadline(void *data) + { + struct sched_attr attr; + int x = 0; + int ret; + unsigned int flags = 0; + + printf("deadline thread started [%ld]\n", gettid()); + + attr.size = sizeof(attr); + attr.sched_flags = 0; + attr.sched_nice = 0; + attr.sched_priority = 0; + + /* This creates a 10ms/30ms reservation */ + attr.sched_policy = SCHED_DEADLINE; + attr.sched_runtime = 10 * 1000 * 1000; + attr.sched_period = attr.sched_deadline = 30 * 1000 * 1000; + + ret = sched_setattr(0, &attr, flags); + if (ret < 0) { + done = 0; + perror("sched_setattr"); + exit(-1); + } + + while (!done) { + x++; + } + + printf("deadline thread dies [%ld]\n", gettid()); + return NULL; + } + + int main (int argc, char **argv) + { + pthread_t thread; + + printf("main thread [%ld]\n", gettid()); + + pthread_create(&thread, NULL, run_deadline, NULL); + + sleep(10); + + done = 1; + pthread_join(thread, NULL); + + printf("main dies [%ld]\n", gettid()); + return 0; + } |