849 lines
35 KiB
Plaintext
849 lines
35 KiB
Plaintext
Deadline Task Scheduling
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------------------------
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CONTENTS
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========
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0. WARNING
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1. Overview
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2. Scheduling algorithm
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2.1 Main algorithm
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2.2 Bandwidth reclaiming
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3. Scheduling Real-Time Tasks
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3.1 Definitions
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3.2 Schedulability Analysis for Uniprocessor Systems
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3.3 Schedulability Analysis for Multiprocessor Systems
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3.4 Relationship with SCHED_DEADLINE Parameters
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4. Bandwidth management
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4.1 System-wide settings
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4.2 Task interface
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4.3 Default behavior
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4.4 Behavior of sched_yield()
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5. Tasks CPU affinity
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5.1 SCHED_DEADLINE and cpusets HOWTO
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6. Future plans
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A. Test suite
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B. Minimal main()
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0. WARNING
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==========
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Fiddling with these settings can result in an unpredictable or even unstable
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system behavior. As for -rt (group) scheduling, it is assumed that root users
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know what they're doing.
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1. Overview
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===========
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The SCHED_DEADLINE policy contained inside the sched_dl scheduling class is
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basically an implementation of the Earliest Deadline First (EDF) scheduling
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algorithm, augmented with a mechanism (called Constant Bandwidth Server, CBS)
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that makes it possible to isolate the behavior of tasks between each other.
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2. Scheduling algorithm
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==================
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2.1 Main algorithm
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------------------
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SCHED_DEADLINE uses three parameters, named "runtime", "period", and
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"deadline", to schedule tasks. A SCHED_DEADLINE task should receive
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"runtime" microseconds of execution time every "period" microseconds, and
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these "runtime" microseconds are available within "deadline" microseconds
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from the beginning of the period. In order to implement this behavior,
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every time the task wakes up, the scheduler computes a "scheduling deadline"
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consistent with the guarantee (using the CBS[2,3] algorithm). Tasks are then
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scheduled using EDF[1] on these scheduling deadlines (the task with the
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earliest scheduling deadline is selected for execution). Notice that the
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task actually receives "runtime" time units within "deadline" if a proper
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"admission control" strategy (see Section "4. Bandwidth management") is used
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(clearly, if the system is overloaded this guarantee cannot be respected).
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Summing up, the CBS[2,3] algorithm assigns scheduling deadlines to tasks so
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that each task runs for at most its runtime every period, avoiding any
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interference between different tasks (bandwidth isolation), while the EDF[1]
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algorithm selects the task with the earliest scheduling deadline as the one
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to be executed next. Thanks to this feature, tasks that do not strictly comply
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with the "traditional" real-time task model (see Section 3) can effectively
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use the new policy.
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In more details, the CBS algorithm assigns scheduling deadlines to
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tasks in the following way:
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- Each SCHED_DEADLINE task is characterized by the "runtime",
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"deadline", and "period" parameters;
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- The state of the task is described by a "scheduling deadline", and
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a "remaining runtime". These two parameters are initially set to 0;
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- When a SCHED_DEADLINE task wakes up (becomes ready for execution),
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the scheduler checks if
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remaining runtime runtime
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---------------------------------- > ---------
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scheduling deadline - current time period
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then, if the scheduling deadline is smaller than the current time, or
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this condition is verified, the scheduling deadline and the
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remaining runtime are re-initialized as
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scheduling deadline = current time + deadline
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remaining runtime = runtime
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otherwise, the scheduling deadline and the remaining runtime are
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left unchanged;
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- When a SCHED_DEADLINE task executes for an amount of time t, its
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remaining runtime is decreased as
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remaining runtime = remaining runtime - t
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(technically, the runtime is decreased at every tick, or when the
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task is descheduled / preempted);
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- When the remaining runtime becomes less or equal than 0, the task is
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said to be "throttled" (also known as "depleted" in real-time literature)
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and cannot be scheduled until its scheduling deadline. The "replenishment
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time" for this task (see next item) is set to be equal to the current
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value of the scheduling deadline;
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- When the current time is equal to the replenishment time of a
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throttled task, the scheduling deadline and the remaining runtime are
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updated as
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scheduling deadline = scheduling deadline + period
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remaining runtime = remaining runtime + runtime
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2.2 Bandwidth reclaiming
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------------------------
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Bandwidth reclaiming for deadline tasks is based on the GRUB (Greedy
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Reclamation of Unused Bandwidth) algorithm [15, 16, 17] and it is enabled
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when flag SCHED_FLAG_RECLAIM is set.
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The following diagram illustrates the state names for tasks handled by GRUB:
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------------
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(d) | Active |
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------------->| |
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| | Contending |
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| ------------
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| A |
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---------- | |
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| | | |
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| Inactive | |(b) | (a)
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| | | |
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---------- | |
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A | V
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| ------------
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| | Active |
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--------------| Non |
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(c) | Contending |
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------------
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A task can be in one of the following states:
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- ActiveContending: if it is ready for execution (or executing);
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- ActiveNonContending: if it just blocked and has not yet surpassed the 0-lag
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time;
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- Inactive: if it is blocked and has surpassed the 0-lag time.
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State transitions:
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(a) When a task blocks, it does not become immediately inactive since its
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bandwidth cannot be immediately reclaimed without breaking the
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real-time guarantees. It therefore enters a transitional state called
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ActiveNonContending. The scheduler arms the "inactive timer" to fire at
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the 0-lag time, when the task's bandwidth can be reclaimed without
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breaking the real-time guarantees.
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The 0-lag time for a task entering the ActiveNonContending state is
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computed as
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(runtime * dl_period)
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deadline - ---------------------
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dl_runtime
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where runtime is the remaining runtime, while dl_runtime and dl_period
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are the reservation parameters.
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(b) If the task wakes up before the inactive timer fires, the task re-enters
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the ActiveContending state and the "inactive timer" is canceled.
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In addition, if the task wakes up on a different runqueue, then
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the task's utilization must be removed from the previous runqueue's active
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utilization and must be added to the new runqueue's active utilization.
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In order to avoid races between a task waking up on a runqueue while the
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"inactive timer" is running on a different CPU, the "dl_non_contending"
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flag is used to indicate that a task is not on a runqueue but is active
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(so, the flag is set when the task blocks and is cleared when the
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"inactive timer" fires or when the task wakes up).
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(c) When the "inactive timer" fires, the task enters the Inactive state and
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its utilization is removed from the runqueue's active utilization.
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(d) When an inactive task wakes up, it enters the ActiveContending state and
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its utilization is added to the active utilization of the runqueue where
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it has been enqueued.
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For each runqueue, the algorithm GRUB keeps track of two different bandwidths:
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- Active bandwidth (running_bw): this is the sum of the bandwidths of all
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tasks in active state (i.e., ActiveContending or ActiveNonContending);
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- Total bandwidth (this_bw): this is the sum of all tasks "belonging" to the
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runqueue, including the tasks in Inactive state.
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The algorithm reclaims the bandwidth of the tasks in Inactive state.
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It does so by decrementing the runtime of the executing task Ti at a pace equal
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to
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dq = -max{ Ui / Umax, (1 - Uinact - Uextra) } dt
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where:
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- Ui is the bandwidth of task Ti;
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- Umax is the maximum reclaimable utilization (subjected to RT throttling
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limits);
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- Uinact is the (per runqueue) inactive utilization, computed as
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(this_bq - running_bw);
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- Uextra is the (per runqueue) extra reclaimable utilization
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(subjected to RT throttling limits).
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Let's now see a trivial example of two deadline tasks with runtime equal
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to 4 and period equal to 8 (i.e., bandwidth equal to 0.5):
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A Task T1
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| |
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| |
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|-------- |----
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| | V
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|---|---|---|---|---|---|---|---|--------->t
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0 1 2 3 4 5 6 7 8
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A Task T2
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| |
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| |
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| ------------------------|
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| | V
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|---|---|---|---|---|---|---|---|--------->t
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0 1 2 3 4 5 6 7 8
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A running_bw
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1 ----------------- ------
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0.5- -----------------
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| |
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|---|---|---|---|---|---|---|---|--------->t
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0 1 2 3 4 5 6 7 8
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- Time t = 0:
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Both tasks are ready for execution and therefore in ActiveContending state.
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Suppose Task T1 is the first task to start execution.
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Since there are no inactive tasks, its runtime is decreased as dq = -1 dt.
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- Time t = 2:
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Suppose that task T1 blocks
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Task T1 therefore enters the ActiveNonContending state. Since its remaining
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runtime is equal to 2, its 0-lag time is equal to t = 4.
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Task T2 start execution, with runtime still decreased as dq = -1 dt since
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there are no inactive tasks.
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- Time t = 4:
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This is the 0-lag time for Task T1. Since it didn't woken up in the
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meantime, it enters the Inactive state. Its bandwidth is removed from
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running_bw.
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Task T2 continues its execution. However, its runtime is now decreased as
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dq = - 0.5 dt because Uinact = 0.5.
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Task T2 therefore reclaims the bandwidth unused by Task T1.
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- Time t = 8:
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Task T1 wakes up. It enters the ActiveContending state again, and the
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running_bw is incremented.
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3. Scheduling Real-Time Tasks
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=============================
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* BIG FAT WARNING ******************************************************
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*
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* This section contains a (not-thorough) summary on classical deadline
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* scheduling theory, and how it applies to SCHED_DEADLINE.
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* The reader can "safely" skip to Section 4 if only interested in seeing
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* how the scheduling policy can be used. Anyway, we strongly recommend
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* to come back here and continue reading (once the urge for testing is
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* satisfied :P) to be sure of fully understanding all technical details.
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************************************************************************
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There are no limitations on what kind of task can exploit this new
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scheduling discipline, even if it must be said that it is particularly
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suited for periodic or sporadic real-time tasks that need guarantees on their
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timing behavior, e.g., multimedia, streaming, control applications, etc.
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3.1 Definitions
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------------------------
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A typical real-time task is composed of a repetition of computation phases
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(task instances, or jobs) which are activated on a periodic or sporadic
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fashion.
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Each job J_j (where J_j is the j^th job of the task) is characterized by an
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arrival time r_j (the time when the job starts), an amount of computation
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time c_j needed to finish the job, and a job absolute deadline d_j, which
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is the time within which the job should be finished. The maximum execution
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time max{c_j} is called "Worst Case Execution Time" (WCET) for the task.
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A real-time task can be periodic with period P if r_{j+1} = r_j + P, or
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sporadic with minimum inter-arrival time P is r_{j+1} >= r_j + P. Finally,
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d_j = r_j + D, where D is the task's relative deadline.
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Summing up, a real-time task can be described as
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Task = (WCET, D, P)
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The utilization of a real-time task is defined as the ratio between its
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WCET and its period (or minimum inter-arrival time), and represents
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the fraction of CPU time needed to execute the task.
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If the total utilization U=sum(WCET_i/P_i) is larger than M (with M equal
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to the number of CPUs), then the scheduler is unable to respect all the
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deadlines.
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Note that total utilization is defined as the sum of the utilizations
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WCET_i/P_i over all the real-time tasks in the system. When considering
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multiple real-time tasks, the parameters of the i-th task are indicated
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with the "_i" suffix.
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Moreover, if the total utilization is larger than M, then we risk starving
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non- real-time tasks by real-time tasks.
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If, instead, the total utilization is smaller than M, then non real-time
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tasks will not be starved and the system might be able to respect all the
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deadlines.
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As a matter of fact, in this case it is possible to provide an upper bound
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for tardiness (defined as the maximum between 0 and the difference
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between the finishing time of a job and its absolute deadline).
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More precisely, it can be proven that using a global EDF scheduler the
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maximum tardiness of each task is smaller or equal than
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((M − 1) · WCET_max − WCET_min)/(M − (M − 2) · U_max) + WCET_max
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where WCET_max = max{WCET_i} is the maximum WCET, WCET_min=min{WCET_i}
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is the minimum WCET, and U_max = max{WCET_i/P_i} is the maximum
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utilization[12].
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3.2 Schedulability Analysis for Uniprocessor Systems
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------------------------
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If M=1 (uniprocessor system), or in case of partitioned scheduling (each
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real-time task is statically assigned to one and only one CPU), it is
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possible to formally check if all the deadlines are respected.
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If D_i = P_i for all tasks, then EDF is able to respect all the deadlines
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of all the tasks executing on a CPU if and only if the total utilization
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of the tasks running on such a CPU is smaller or equal than 1.
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If D_i != P_i for some task, then it is possible to define the density of
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a task as WCET_i/min{D_i,P_i}, and EDF is able to respect all the deadlines
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of all the tasks running on a CPU if the sum of the densities of the tasks
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running on such a CPU is smaller or equal than 1:
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sum(WCET_i / min{D_i, P_i}) <= 1
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It is important to notice that this condition is only sufficient, and not
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necessary: there are task sets that are schedulable, but do not respect the
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condition. For example, consider the task set {Task_1,Task_2} composed by
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Task_1=(50ms,50ms,100ms) and Task_2=(10ms,100ms,100ms).
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EDF is clearly able to schedule the two tasks without missing any deadline
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(Task_1 is scheduled as soon as it is released, and finishes just in time
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to respect its deadline; Task_2 is scheduled immediately after Task_1, hence
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its response time cannot be larger than 50ms + 10ms = 60ms) even if
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50 / min{50,100} + 10 / min{100, 100} = 50 / 50 + 10 / 100 = 1.1
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Of course it is possible to test the exact schedulability of tasks with
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D_i != P_i (checking a condition that is both sufficient and necessary),
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but this cannot be done by comparing the total utilization or density with
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a constant. Instead, the so called "processor demand" approach can be used,
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computing the total amount of CPU time h(t) needed by all the tasks to
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respect all of their deadlines in a time interval of size t, and comparing
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such a time with the interval size t. If h(t) is smaller than t (that is,
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the amount of time needed by the tasks in a time interval of size t is
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smaller than the size of the interval) for all the possible values of t, then
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EDF is able to schedule the tasks respecting all of their deadlines. Since
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performing this check for all possible values of t is impossible, it has been
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proven[4,5,6] that it is sufficient to perform the test for values of t
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between 0 and a maximum value L. The cited papers contain all of the
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mathematical details and explain how to compute h(t) and L.
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In any case, this kind of analysis is too complex as well as too
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time-consuming to be performed on-line. Hence, as explained in Section
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4 Linux uses an admission test based on the tasks' utilizations.
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3.3 Schedulability Analysis for Multiprocessor Systems
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------------------------
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On multiprocessor systems with global EDF scheduling (non partitioned
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systems), a sufficient test for schedulability can not be based on the
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utilizations or densities: it can be shown that even if D_i = P_i task
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sets with utilizations slightly larger than 1 can miss deadlines regardless
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of the number of CPUs.
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Consider a set {Task_1,...Task_{M+1}} of M+1 tasks on a system with M
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CPUs, with the first task Task_1=(P,P,P) having period, relative deadline
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and WCET equal to P. The remaining M tasks Task_i=(e,P-1,P-1) have an
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arbitrarily small worst case execution time (indicated as "e" here) and a
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period smaller than the one of the first task. Hence, if all the tasks
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activate at the same time t, global EDF schedules these M tasks first
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(because their absolute deadlines are equal to t + P - 1, hence they are
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smaller than the absolute deadline of Task_1, which is t + P). As a
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result, Task_1 can be scheduled only at time t + e, and will finish at
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time t + e + P, after its absolute deadline. The total utilization of the
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task set is U = M · e / (P - 1) + P / P = M · e / (P - 1) + 1, and for small
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values of e this can become very close to 1. This is known as "Dhall's
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effect"[7]. Note: the example in the original paper by Dhall has been
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slightly simplified here (for example, Dhall more correctly computed
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lim_{e->0}U).
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More complex schedulability tests for global EDF have been developed in
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real-time literature[8,9], but they are not based on a simple comparison
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between total utilization (or density) and a fixed constant. If all tasks
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have D_i = P_i, a sufficient schedulability condition can be expressed in
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a simple way:
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sum(WCET_i / P_i) <= M - (M - 1) · U_max
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where U_max = max{WCET_i / P_i}[10]. Notice that for U_max = 1,
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M - (M - 1) · U_max becomes M - M + 1 = 1 and this schedulability condition
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just confirms the Dhall's effect. A more complete survey of the literature
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about schedulability tests for multi-processor real-time scheduling can be
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found in [11].
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As seen, enforcing that the total utilization is smaller than M does not
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guarantee that global EDF schedules the tasks without missing any deadline
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(in other words, global EDF is not an optimal scheduling algorithm). However,
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a total utilization smaller than M is enough to guarantee that non real-time
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tasks are not starved and that the tardiness of real-time tasks has an upper
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bound[12] (as previously noted). Different bounds on the maximum tardiness
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experienced by real-time tasks have been developed in various papers[13,14],
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but the theoretical result that is important for SCHED_DEADLINE is that if
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the total utilization is smaller or equal than M then the response times of
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the tasks are limited.
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3.4 Relationship with SCHED_DEADLINE Parameters
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------------------------
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Finally, it is important to understand the relationship between the
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SCHED_DEADLINE scheduling parameters described in Section 2 (runtime,
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deadline and period) and the real-time task parameters (WCET, D, P)
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described in this section. Note that the tasks' temporal constraints are
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represented by its absolute deadlines d_j = r_j + D described above, while
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SCHED_DEADLINE schedules the tasks according to scheduling deadlines (see
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Section 2).
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If an admission test is used to guarantee that the scheduling deadlines
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are respected, then SCHED_DEADLINE can be used to schedule real-time tasks
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guaranteeing that all the jobs' deadlines of a task are respected.
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In order to do this, a task must be scheduled by setting:
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- runtime >= WCET
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- deadline = D
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- period <= P
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IOW, if runtime >= WCET and if period is <= P, then the scheduling deadlines
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and the absolute deadlines (d_j) coincide, so a proper admission control
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allows to respect the jobs' absolute deadlines for this task (this is what is
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called "hard schedulability property" and is an extension of Lemma 1 of [2]).
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Notice that if runtime > deadline the admission control will surely reject
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this task, as it is not possible to respect its temporal constraints.
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References:
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1 - C. L. Liu and J. W. Layland. Scheduling algorithms for multiprogram-
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ming in a hard-real-time environment. Journal of the Association for
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Computing Machinery, 20(1), 1973.
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2 - L. Abeni , G. Buttazzo. Integrating Multimedia Applications in Hard
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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.
|
||
|
||
|
||
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.txt), 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/cgroup-v1/cpusets.txt).
|
||
|
||
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;
|
||
}
|