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Diffstat (limited to 'Documentation/scheduler')
| -rw-r--r-- | Documentation/scheduler/completion.rst | 293 | ||||
| -rw-r--r-- | Documentation/scheduler/index.rst | 28 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-arch.rst | 76 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-bwc.rst | 174 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-capacity.rst | 439 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-deadline.rst | 888 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-design-CFS.rst | 249 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-domains.rst | 85 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-energy.rst | 422 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-nice-design.rst | 112 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-pelt.c | 109 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-rt-group.rst | 185 | ||||
| -rw-r--r-- | Documentation/scheduler/sched-stats.rst | 167 | ||||
| -rw-r--r-- | Documentation/scheduler/text_files.rst | 5 |
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diff --git a/Documentation/scheduler/completion.rst b/Documentation/scheduler/completion.rst new file mode 100644 index 000000000..9f039b4f4 --- /dev/null +++ b/Documentation/scheduler/completion.rst @@ -0,0 +1,293 @@ +================================================ +Completions - "wait for completion" barrier APIs +================================================ + +Introduction: +------------- + +If you have one or more threads that must wait for some kernel activity +to have reached a point or a specific state, completions can provide a +race-free solution to this problem. Semantically they are somewhat like a +pthread_barrier() and have similar use-cases. + +Completions are a code synchronization mechanism which is preferable to any +misuse of locks/semaphores and busy-loops. Any time you think of using +yield() or some quirky msleep(1) loop to allow something else to proceed, +you probably want to look into using one of the wait_for_completion*() +calls and complete() instead. + +The advantage of using completions is that they have a well defined, focused +purpose which makes it very easy to see the intent of the code, but they +also result in more efficient code as all threads can continue execution +until the result is actually needed, and both the waiting and the signalling +is highly efficient using low level scheduler sleep/wakeup facilities. + +Completions are built on top of the waitqueue and wakeup infrastructure of +the Linux scheduler. The event the threads on the waitqueue are waiting for +is reduced to a simple flag in 'struct completion', appropriately called "done". + +As completions are scheduling related, the code can be found in +kernel/sched/completion.c. + + +Usage: +------ + +There are three main parts to using completions: + + - the initialization of the 'struct completion' synchronization object + - the waiting part through a call to one of the variants of wait_for_completion(), + - the signaling side through a call to complete() or complete_all(). + +There are also some helper functions for checking the state of completions. +Note that while initialization must happen first, the waiting and signaling +part can happen in any order. I.e. it's entirely normal for a thread +to have marked a completion as 'done' before another thread checks whether +it has to wait for it. + +To use completions you need to #include <linux/completion.h> and +create a static or dynamic variable of type 'struct completion', +which has only two fields:: + + struct completion { + unsigned int done; + wait_queue_head_t wait; + }; + +This provides the ->wait waitqueue to place tasks on for waiting (if any), and +the ->done completion flag for indicating whether it's completed or not. + +Completions should be named to refer to the event that is being synchronized on. +A good example is:: + + wait_for_completion(&early_console_added); + + complete(&early_console_added); + +Good, intuitive naming (as always) helps code readability. Naming a completion +'complete' is not helpful unless the purpose is super obvious... + + +Initializing completions: +------------------------- + +Dynamically allocated completion objects should preferably be embedded in data +structures that are assured to be alive for the life-time of the function/driver, +to prevent races with asynchronous complete() calls from occurring. + +Particular care should be taken when using the _timeout() or _killable()/_interruptible() +variants of wait_for_completion(), as it must be assured that memory de-allocation +does not happen until all related activities (complete() or reinit_completion()) +have taken place, even if these wait functions return prematurely due to a timeout +or a signal triggering. + +Initializing of dynamically allocated completion objects is done via a call to +init_completion():: + + init_completion(&dynamic_object->done); + +In this call we initialize the waitqueue and set ->done to 0, i.e. "not completed" +or "not done". + +The re-initialization function, reinit_completion(), simply resets the +->done field to 0 ("not done"), without touching the waitqueue. +Callers of this function must make sure that there are no racy +wait_for_completion() calls going on in parallel. + +Calling init_completion() on the same completion object twice is +most likely a bug as it re-initializes the queue to an empty queue and +enqueued tasks could get "lost" - use reinit_completion() in that case, +but be aware of other races. + +For static declaration and initialization, macros are available. + +For static (or global) declarations in file scope you can use +DECLARE_COMPLETION():: + + static DECLARE_COMPLETION(setup_done); + DECLARE_COMPLETION(setup_done); + +Note that in this case the completion is boot time (or module load time) +initialized to 'not done' and doesn't require an init_completion() call. + +When a completion is declared as a local variable within a function, +then the initialization should always use DECLARE_COMPLETION_ONSTACK() +explicitly, not just to make lockdep happy, but also to make it clear +that limited scope had been considered and is intentional:: + + DECLARE_COMPLETION_ONSTACK(setup_done) + +Note that when using completion objects as local variables you must be +acutely aware of the short life time of the function stack: the function +must not return to a calling context until all activities (such as waiting +threads) have ceased and the completion object is completely unused. + +To emphasise this again: in particular when using some of the waiting API variants +with more complex outcomes, such as the timeout or signalling (_timeout(), +_killable() and _interruptible()) variants, the wait might complete +prematurely while the object might still be in use by another thread - and a return +from the wait_on_completion*() caller function will deallocate the function +stack and cause subtle data corruption if a complete() is done in some +other thread. Simple testing might not trigger these kinds of races. + +If unsure, use dynamically allocated completion objects, preferably embedded +in some other long lived object that has a boringly long life time which +exceeds the life time of any helper threads using the completion object, +or has a lock or other synchronization mechanism to make sure complete() +is not called on a freed object. + +A naive DECLARE_COMPLETION() on the stack triggers a lockdep warning. + +Waiting for completions: +------------------------ + +For a thread to wait for some concurrent activity to finish, it +calls wait_for_completion() on the initialized completion structure:: + + void wait_for_completion(struct completion *done) + +A typical usage scenario is:: + + CPU#1 CPU#2 + + struct completion setup_done; + + init_completion(&setup_done); + initialize_work(...,&setup_done,...); + + /* run non-dependent code */ /* do setup */ + + wait_for_completion(&setup_done); complete(setup_done); + +This is not implying any particular order between wait_for_completion() and +the call to complete() - if the call to complete() happened before the call +to wait_for_completion() then the waiting side simply will continue +immediately as all dependencies are satisfied; if not, it will block until +completion is signaled by complete(). + +Note that wait_for_completion() is calling spin_lock_irq()/spin_unlock_irq(), +so it can only be called safely when you know that interrupts are enabled. +Calling it from IRQs-off atomic contexts will result in hard-to-detect +spurious enabling of interrupts. + +The default behavior is to wait without a timeout and to mark the task as +uninterruptible. wait_for_completion() and its variants are only safe +in process context (as they can sleep) but not in atomic context, +interrupt context, with disabled IRQs, or preemption is disabled - see also +try_wait_for_completion() below for handling completion in atomic/interrupt +context. + +As all variants of wait_for_completion() can (obviously) block for a long +time depending on the nature of the activity they are waiting for, so in +most cases you probably don't want to call this with held mutexes. + + +wait_for_completion*() variants available: +------------------------------------------ + +The below variants all return status and this status should be checked in +most(/all) cases - in cases where the status is deliberately not checked you +probably want to make a note explaining this (e.g. see +arch/arm/kernel/smp.c:__cpu_up()). + +A common problem that occurs is to have unclean assignment of return types, +so take care to assign return-values to variables of the proper type. + +Checking for the specific meaning of return values also has been found +to be quite inaccurate, e.g. constructs like:: + + if (!wait_for_completion_interruptible_timeout(...)) + +... would execute the same code path for successful completion and for the +interrupted case - which is probably not what you want:: + + int wait_for_completion_interruptible(struct completion *done) + +This function marks the task TASK_INTERRUPTIBLE while it is waiting. +If a signal was received while waiting it will return -ERESTARTSYS; 0 otherwise:: + + unsigned long wait_for_completion_timeout(struct completion *done, unsigned long timeout) + +The task is marked as TASK_UNINTERRUPTIBLE and will wait at most 'timeout' +jiffies. If a timeout occurs it returns 0, else the remaining time in +jiffies (but at least 1). + +Timeouts are preferably calculated with msecs_to_jiffies() or usecs_to_jiffies(), +to make the code largely HZ-invariant. + +If the returned timeout value is deliberately ignored a comment should probably explain +why (e.g. see drivers/mfd/wm8350-core.c wm8350_read_auxadc()):: + + long wait_for_completion_interruptible_timeout(struct completion *done, unsigned long timeout) + +This function passes a timeout in jiffies and marks the task as +TASK_INTERRUPTIBLE. If a signal was received it will return -ERESTARTSYS; +otherwise it returns 0 if the completion timed out, or the remaining time in +jiffies if completion occurred. + +Further variants include _killable which uses TASK_KILLABLE as the +designated tasks state and will return -ERESTARTSYS if it is interrupted, +or 0 if completion was achieved. There is a _timeout variant as well:: + + long wait_for_completion_killable(struct completion *done) + long wait_for_completion_killable_timeout(struct completion *done, unsigned long timeout) + +The _io variants wait_for_completion_io() behave the same as the non-_io +variants, except for accounting waiting time as 'waiting on IO', which has +an impact on how the task is accounted in scheduling/IO stats:: + + void wait_for_completion_io(struct completion *done) + unsigned long wait_for_completion_io_timeout(struct completion *done, unsigned long timeout) + + +Signaling completions: +---------------------- + +A thread that wants to signal that the conditions for continuation have been +achieved calls complete() to signal exactly one of the waiters that it can +continue:: + + void complete(struct completion *done) + +... or calls complete_all() to signal all current and future waiters:: + + void complete_all(struct completion *done) + +The signaling will work as expected even if completions are signaled before +a thread starts waiting. This is achieved by the waiter "consuming" +(decrementing) the done field of 'struct completion'. Waiting threads +wakeup order is the same in which they were enqueued (FIFO order). + +If complete() is called multiple times then this will allow for that number +of waiters to continue - each call to complete() will simply increment the +done field. Calling complete_all() multiple times is a bug though. Both +complete() and complete_all() can be called in IRQ/atomic context safely. + +There can only be one thread calling complete() or complete_all() on a +particular 'struct completion' at any time - serialized through the wait +queue spinlock. Any such concurrent calls to complete() or complete_all() +probably are a design bug. + +Signaling completion from IRQ context is fine as it will appropriately +lock with spin_lock_irqsave()/spin_unlock_irqrestore() and it will never +sleep. + + +try_wait_for_completion()/completion_done(): +-------------------------------------------- + +The try_wait_for_completion() function will not put the thread on the wait +queue but rather returns false if it would need to enqueue (block) the thread, +else it consumes one posted completion and returns true:: + + bool try_wait_for_completion(struct completion *done) + +Finally, to check the state of a completion without changing it in any way, +call completion_done(), which returns false if there are no posted +completions that were not yet consumed by waiters (implying that there are +waiters) and true otherwise:: + + bool completion_done(struct completion *done) + +Both try_wait_for_completion() and completion_done() are safe to be called in +IRQ or atomic context. diff --git a/Documentation/scheduler/index.rst b/Documentation/scheduler/index.rst new file mode 100644 index 000000000..88900aabd --- /dev/null +++ b/Documentation/scheduler/index.rst @@ -0,0 +1,28 @@ +=============== +Linux Scheduler +=============== + +.. toctree:: + :maxdepth: 1 + + + completion + sched-arch + sched-bwc + sched-deadline + sched-design-CFS + sched-domains + sched-capacity + sched-energy + sched-nice-design + sched-rt-group + sched-stats + + text_files + +.. only:: subproject and html + + Indices + ======= + + * :ref:`genindex` diff --git a/Documentation/scheduler/sched-arch.rst b/Documentation/scheduler/sched-arch.rst new file mode 100644 index 000000000..0eaec6697 --- /dev/null +++ b/Documentation/scheduler/sched-arch.rst @@ -0,0 +1,76 @@ +================================================================= +CPU Scheduler implementation hints for architecture specific code +================================================================= + + Nick Piggin, 2005 + +Context switch +============== +1. Runqueue locking +By default, the switch_to arch function is called with the runqueue +locked. This is usually not a problem unless switch_to may need to +take the runqueue lock. This is usually due to a wake up operation in +the context switch. See arch/ia64/include/asm/switch_to.h for an example. + +To request the scheduler call switch_to with the runqueue unlocked, +you must `#define __ARCH_WANT_UNLOCKED_CTXSW` in a header file +(typically the one where switch_to is defined). + +Unlocked context switches introduce only a very minor performance +penalty to the core scheduler implementation in the CONFIG_SMP case. + +CPU idle +======== +Your cpu_idle routines need to obey the following rules: + +1. Preempt should now disabled over idle routines. Should only + be enabled to call schedule() then disabled again. + +2. need_resched/TIF_NEED_RESCHED is only ever set, and will never + be cleared until the running task has called schedule(). Idle + threads need only ever query need_resched, and may never set or + clear it. + +3. When cpu_idle finds (need_resched() == 'true'), it should call + schedule(). It should not call schedule() otherwise. + +4. The only time interrupts need to be disabled when checking + need_resched is if we are about to sleep the processor until + the next interrupt (this doesn't provide any protection of + need_resched, it prevents losing an interrupt): + + 4a. Common problem with this type of sleep appears to be:: + + local_irq_disable(); + if (!need_resched()) { + local_irq_enable(); + *** resched interrupt arrives here *** + __asm__("sleep until next interrupt"); + } + +5. TIF_POLLING_NRFLAG can be set by idle routines that do not + need an interrupt to wake them up when need_resched goes high. + In other words, they must be periodically polling need_resched, + although it may be reasonable to do some background work or enter + a low CPU priority. + + - 5a. If TIF_POLLING_NRFLAG is set, and we do decide to enter + an interrupt sleep, it needs to be cleared then a memory + barrier issued (followed by a test of need_resched with + interrupts disabled, as explained in 3). + +arch/x86/kernel/process.c has examples of both polling and +sleeping idle functions. + + +Possible arch/ problems +======================= + +Possible arch problems I found (and either tried to fix or didn't): + +ia64 - is safe_halt call racy vs interrupts? (does it sleep?) (See #4a) + +sh64 - Is sleeping racy vs interrupts? (See #4a) + +sparc - IRQs on at this point(?), change local_irq_save to _disable. + - TODO: needs secondary CPUs to disable preempt (See #1) diff --git a/Documentation/scheduler/sched-bwc.rst b/Documentation/scheduler/sched-bwc.rst new file mode 100644 index 000000000..9801d6b28 --- /dev/null +++ b/Documentation/scheduler/sched-bwc.rst @@ -0,0 +1,174 @@ +===================== +CFS Bandwidth Control +===================== + +[ This document only discusses CPU bandwidth control for SCHED_NORMAL. + The SCHED_RT case is covered in Documentation/scheduler/sched-rt-group.rst ] + +CFS bandwidth control is a CONFIG_FAIR_GROUP_SCHED extension which allows the +specification of the maximum CPU bandwidth available to a group or hierarchy. + +The bandwidth allowed for a group is specified using a quota and period. Within +each given "period" (microseconds), a task group is allocated up to "quota" +microseconds of CPU time. That quota is assigned to per-cpu run queues in +slices as threads in the cgroup become runnable. Once all quota has been +assigned any additional requests for quota will result in those threads being +throttled. Throttled threads will not be able to run again until the next +period when the quota is replenished. + +A group's unassigned quota is globally tracked, being refreshed back to +cfs_quota units at each period boundary. As threads consume this bandwidth it +is transferred to cpu-local "silos" on a demand basis. The amount transferred +within each of these updates is tunable and described as the "slice". + +Management +---------- +Quota and period are managed within the cpu subsystem via cgroupfs. + +cpu.cfs_quota_us: the total available run-time within a period (in microseconds) +cpu.cfs_period_us: the length of a period (in microseconds) +cpu.stat: exports throttling statistics [explained further below] + +The default values are:: + + cpu.cfs_period_us=100ms + cpu.cfs_quota=-1 + +A value of -1 for cpu.cfs_quota_us indicates that the group does not have any +bandwidth restriction in place, such a group is described as an unconstrained +bandwidth group. This represents the traditional work-conserving behavior for +CFS. + +Writing any (valid) positive value(s) will enact the specified bandwidth limit. +The minimum quota allowed for the quota or period is 1ms. There is also an +upper bound on the period length of 1s. Additional restrictions exist when +bandwidth limits are used in a hierarchical fashion, these are explained in +more detail below. + +Writing any negative value to cpu.cfs_quota_us will remove the bandwidth limit +and return the group to an unconstrained state once more. + +Any updates to a group's bandwidth specification will result in it becoming +unthrottled if it is in a constrained state. + +System wide settings +-------------------- +For efficiency run-time is transferred between the global pool and CPU local +"silos" in a batch fashion. This greatly reduces global accounting pressure +on large systems. The amount transferred each time such an update is required +is described as the "slice". + +This is tunable via procfs:: + + /proc/sys/kernel/sched_cfs_bandwidth_slice_us (default=5ms) + +Larger slice values will reduce transfer overheads, while smaller values allow +for more fine-grained consumption. + +Statistics +---------- +A group's bandwidth statistics are exported via 3 fields in cpu.stat. + +cpu.stat: + +- nr_periods: Number of enforcement intervals that have elapsed. +- nr_throttled: Number of times the group has been throttled/limited. +- throttled_time: The total time duration (in nanoseconds) for which entities + of the group have been throttled. + +This interface is read-only. + +Hierarchical considerations +--------------------------- +The interface enforces that an individual entity's bandwidth is always +attainable, that is: max(c_i) <= C. However, over-subscription in the +aggregate case is explicitly allowed to enable work-conserving semantics +within a hierarchy: + + e.g. \Sum (c_i) may exceed C + +[ Where C is the parent's bandwidth, and c_i its children ] + + +There are two ways in which a group may become throttled: + + a. it fully consumes its own quota within a period + b. a parent's quota is fully consumed within its period + +In case b) above, even though the child may have runtime remaining it will not +be allowed to until the parent's runtime is refreshed. + +CFS Bandwidth Quota Caveats +--------------------------- +Once a slice is assigned to a cpu it does not expire. However all but 1ms of +the slice may be returned to the global pool if all threads on that cpu become +unrunnable. This is configured at compile time by the min_cfs_rq_runtime +variable. This is a performance tweak that helps prevent added contention on +the global lock. + +The fact that cpu-local slices do not expire results in some interesting corner +cases that should be understood. + +For cgroup cpu constrained applications that are cpu limited this is a +relatively moot point because they will naturally consume the entirety of their +quota as well as the entirety of each cpu-local slice in each period. As a +result it is expected that nr_periods roughly equal nr_throttled, and that +cpuacct.usage will increase roughly equal to cfs_quota_us in each period. + +For highly-threaded, non-cpu bound applications this non-expiration nuance +allows applications to briefly burst past their quota limits by the amount of +unused slice on each cpu that the task group is running on (typically at most +1ms per cpu or as defined by min_cfs_rq_runtime). This slight burst only +applies if quota had been assigned to a cpu and then not fully used or returned +in previous periods. This burst amount will not be transferred between cores. +As a result, this mechanism still strictly limits the task group to quota +average usage, albeit over a longer time window than a single period. This +also limits the burst ability to no more than 1ms per cpu. This provides +better more predictable user experience for highly threaded applications with +small quota limits on high core count machines. It also eliminates the +propensity to throttle these applications while simultanously using less than +quota amounts of cpu. Another way to say this, is that by allowing the unused +portion of a slice to remain valid across periods we have decreased the +possibility of wastefully expiring quota on cpu-local silos that don't need a +full slice's amount of cpu time. + +The interaction between cpu-bound and non-cpu-bound-interactive applications +should also be considered, especially when single core usage hits 100%. If you +gave each of these applications half of a cpu-core and they both got scheduled +on the same CPU it is theoretically possible that the non-cpu bound application +will use up to 1ms additional quota in some periods, thereby preventing the +cpu-bound application from fully using its quota by that same amount. In these +instances it will be up to the CFS algorithm (see sched-design-CFS.rst) to +decide which application is chosen to run, as they will both be runnable and +have remaining quota. This runtime discrepancy will be made up in the following +periods when the interactive application idles. + +Examples +-------- +1. Limit a group to 1 CPU worth of runtime:: + + If period is 250ms and quota is also 250ms, the group will get + 1 CPU worth of runtime every 250ms. + + # echo 250000 > cpu.cfs_quota_us /* quota = 250ms */ + # echo 250000 > cpu.cfs_period_us /* period = 250ms */ + +2. Limit a group to 2 CPUs worth of runtime on a multi-CPU machine + + With 500ms period and 1000ms quota, the group can get 2 CPUs worth of + runtime every 500ms:: + + # echo 1000000 > cpu.cfs_quota_us /* quota = 1000ms */ + # echo 500000 > cpu.cfs_period_us /* period = 500ms */ + + The larger period here allows for increased burst capacity. + +3. Limit a group to 20% of 1 CPU. + + With 50ms period, 10ms quota will be equivalent to 20% of 1 CPU:: + + # echo 10000 > cpu.cfs_quota_us /* quota = 10ms */ + # echo 50000 > cpu.cfs_period_us /* period = 50ms */ + + By using a small period here we are ensuring a consistent latency + response at the expense of burst capacity. diff --git a/Documentation/scheduler/sched-capacity.rst b/Documentation/scheduler/sched-capacity.rst new file mode 100644 index 000000000..9b7cbe43b --- /dev/null +++ b/Documentation/scheduler/sched-capacity.rst @@ -0,0 +1,439 @@ +========================= +Capacity Aware Scheduling +========================= + +1. CPU Capacity +=============== + +1.1 Introduction +---------------- + +Conventional, homogeneous SMP platforms are composed of purely identical +CPUs. Heterogeneous platforms on the other hand are composed of CPUs with +different performance characteristics - on such platforms, not all CPUs can be +considered equal. + +CPU capacity is a measure of the performance a CPU can reach, normalized against +the most performant CPU in the system. Heterogeneous systems are also called +asymmetric CPU capacity systems, as they contain CPUs of different capacities. + +Disparity in maximum attainable performance (IOW in maximum CPU capacity) stems +from two factors: + +- not all CPUs may have the same microarchitecture (µarch). +- with Dynamic Voltage and Frequency Scaling (DVFS), not all CPUs may be + physically able to attain the higher Operating Performance Points (OPP). + +Arm big.LITTLE systems are an example of both. The big CPUs are more +performance-oriented than the LITTLE ones (more pipeline stages, bigger caches, +smarter predictors, etc), and can usually reach higher OPPs than the LITTLE ones +can. + +CPU performance is usually expressed in Millions of Instructions Per Second +(MIPS), which can also be expressed as a given amount of instructions attainable +per Hz, leading to:: + + capacity(cpu) = work_per_hz(cpu) * max_freq(cpu) + +1.2 Scheduler terms +------------------- + +Two different capacity values are used within the scheduler. A CPU's +``capacity_orig`` is its maximum attainable capacity, i.e. its maximum +attainable performance level. A CPU's ``capacity`` is its ``capacity_orig`` to +which some loss of available performance (e.g. time spent handling IRQs) is +subtracted. + +Note that a CPU's ``capacity`` is solely intended to be used by the CFS class, +while ``capacity_orig`` is class-agnostic. The rest of this document will use +the term ``capacity`` interchangeably with ``capacity_orig`` for the sake of +brevity. + +1.3 Platform examples +--------------------- + +1.3.1 Identical OPPs +~~~~~~~~~~~~~~~~~~~~ + +Consider an hypothetical dual-core asymmetric CPU capacity system where + +- work_per_hz(CPU0) = W +- work_per_hz(CPU1) = W/2 +- all CPUs are running at the same fixed frequency + +By the above definition of capacity: + +- capacity(CPU0) = C +- capacity(CPU1) = C/2 + +To draw the parallel with Arm big.LITTLE, CPU0 would be a big while CPU1 would +be a LITTLE. + +With a workload that periodically does a fixed amount of work, you will get an +execution trace like so:: + + CPU0 work ^ + | ____ ____ ____ + | | | | | | | + +----+----+----+----+----+----+----+----+----+----+-> time + + CPU1 work ^ + | _________ _________ ____ + | | | | | | + +----+----+----+----+----+----+----+----+----+----+-> time + +CPU0 has the highest capacity in the system (C), and completes a fixed amount of +work W in T units of time. On the other hand, CPU1 has half the capacity of +CPU0, and thus only completes W/2 in T. + +1.3.2 Different max OPPs +~~~~~~~~~~~~~~~~~~~~~~~~ + +Usually, CPUs of different capacity values also have different maximum +OPPs. Consider the same CPUs as above (i.e. same work_per_hz()) with: + +- max_freq(CPU0) = F +- max_freq(CPU1) = 2/3 * F + +This yields: + +- capacity(CPU0) = C +- capacity(CPU1) = C/3 + +Executing the same workload as described in 1.3.1, which each CPU running at its +maximum frequency results in:: + + CPU0 work ^ + | ____ ____ ____ + | | | | | | | + +----+----+----+----+----+----+----+----+----+----+-> time + + workload on CPU1 + CPU1 work ^ + | ______________ ______________ ____ + | | | | | | + +----+----+----+----+----+----+----+----+----+----+-> time + +1.4 Representation caveat +------------------------- + +It should be noted that having a *single* value to represent differences in CPU +performance is somewhat of a contentious point. The relative performance +difference between two different µarchs could be X% on integer operations, Y% on +floating point operations, Z% on branches, and so on. Still, results using this +simple approach have been satisfactory for now. + +2. Task utilization +=================== + +2.1 Introduction +---------------- + +Capacity aware scheduling requires an expression of a task's requirements with +regards to CPU capacity. Each scheduler class can express this differently, and +while task utilization is specific to CFS, it is convenient to describe it here +in order to introduce more generic concepts. + +Task utilization is a percentage meant to represent the throughput requirements +of a task. A simple approximation of it is the task's duty cycle, i.e.:: + + task_util(p) = duty_cycle(p) + +On an SMP system with fixed frequencies, 100% utilization suggests the task is a +busy loop. Conversely, 10% utilization hints it is a small periodic task that +spends more time sleeping than executing. Variable CPU frequencies and +asymmetric CPU capacities complexify this somewhat; the following sections will +expand on these. + +2.2 Frequency invariance +------------------------ + +One issue that needs to be taken into account is that a workload's duty cycle is +directly impacted by the current OPP the CPU is running at. Consider running a +periodic workload at a given frequency F:: + + CPU work ^ + | ____ ____ ____ + | | | | | | | + +----+----+----+----+----+----+----+----+----+----+-> time + +This yields duty_cycle(p) == 25%. + +Now, consider running the *same* workload at frequency F/2:: + + CPU work ^ + | _________ _________ ____ + | | | | | | + +----+----+----+----+----+----+----+----+----+----+-> time + +This yields duty_cycle(p) == 50%, despite the task having the exact same +behaviour (i.e. executing the same amount of work) in both executions. + +The task utilization signal can be made frequency invariant using the following +formula:: + + task_util_freq_inv(p) = duty_cycle(p) * (curr_frequency(cpu) / max_frequency(cpu)) + +Applying this formula to the two examples above yields a frequency invariant +task utilization of 25%. + +2.3 CPU invariance +------------------ + +CPU capacity has a similar effect on task utilization in that running an +identical workload on CPUs of different capacity values will yield different +duty cycles. + +Consider the system described in 1.3.2., i.e.:: + +- capacity(CPU0) = C +- capacity(CPU1) = C/3 + +Executing a given periodic workload on each CPU at their maximum frequency would +result in:: + + CPU0 work ^ + | ____ ____ ____ + | | | | | | | + +----+----+----+----+----+----+----+----+----+----+-> time + + CPU1 work ^ + | ______________ ______________ ____ + | | | | | | + +----+----+----+----+----+----+----+----+----+----+-> time + +IOW, + +- duty_cycle(p) == 25% if p runs on CPU0 at its maximum frequency +- duty_cycle(p) == 75% if p runs on CPU1 at its maximum frequency + +The task utilization signal can be made CPU invariant using the following +formula:: + + task_util_cpu_inv(p) = duty_cycle(p) * (capacity(cpu) / max_capacity) + +with ``max_capacity`` being the highest CPU capacity value in the +system. Applying this formula to the above example above yields a CPU +invariant task utilization of 25%. + +2.4 Invariant task utilization +------------------------------ + +Both frequency and CPU invariance need to be applied to task utilization in +order to obtain a truly invariant signal. The pseudo-formula for a task +utilization that is both CPU and frequency invariant is thus, for a given +task p:: + + curr_frequency(cpu) capacity(cpu) + task_util_inv(p) = duty_cycle(p) * ------------------- * ------------- + max_frequency(cpu) max_capacity + +In other words, invariant task utilization describes the behaviour of a task as +if it were running on the highest-capacity CPU in the system, running at its +maximum frequency. + +Any mention of task utilization in the following sections will imply its +invariant form. + +2.5 Utilization estimation +-------------------------- + +Without a crystal ball, task behaviour (and thus task utilization) cannot +accurately be predicted the moment a task first becomes runnable. The CFS class +maintains a handful of CPU and task signals based on the Per-Entity Load +Tracking (PELT) mechanism, one of those yielding an *average* utilization (as +opposed to instantaneous). + +This means that while the capacity aware scheduling criteria will be written +considering a "true" task utilization (using a crystal ball), the implementation +will only ever be able to use an estimator thereof. + +3. Capacity aware scheduling requirements +========================================= + +3.1 CPU capacity +---------------- + +Linux cannot currently figure out CPU capacity on its own, this information thus +needs to be handed to it. Architectures must define arch_scale_cpu_capacity() +for that purpose. + +The arm and arm64 architectures directly map this to the arch_topology driver +CPU scaling data, which is derived from the capacity-dmips-mhz CPU binding; see +Documentation/devicetree/bindings/arm/cpu-capacity.txt. + +3.2 Frequency invariance +------------------------ + +As stated in 2.2, capacity-aware scheduling requires a frequency-invariant task +utilization. Architectures must define arch_scale_freq_capacity(cpu) for that +purpose. + +Implementing this function requires figuring out at which frequency each CPU +have been running at. One way to implement this is to leverage hardware counters +whose increment rate scale with a CPU's current frequency (APERF/MPERF on x86, +AMU on arm64). Another is to directly hook into cpufreq frequency transitions, +when the kernel is aware of the switched-to frequency (also employed by +arm/arm64). + +4. Scheduler topology +===================== + +During the construction of the sched domains, the scheduler will figure out +whether the system exhibits asymmetric CPU capacities. Should that be the +case: + +- The sched_asym_cpucapacity static key will be enabled. +- The SD_ASYM_CPUCAPACITY flag will be set at the lowest sched_domain level that + spans all unique CPU capacity values. + +The sched_asym_cpucapacity static key is intended to guard sections of code that +cater to asymmetric CPU capacity systems. Do note however that said key is +*system-wide*. Imagine the following setup using cpusets:: + + capacity C/2 C + ________ ________ + / \ / \ + CPUs 0 1 2 3 4 5 6 7 + \__/ \______________/ + cpusets cs0 cs1 + +Which could be created via: + +.. code-block:: sh + + mkdir /sys/fs/cgroup/cpuset/cs0 + echo 0-1 > /sys/fs/cgroup/cpuset/cs0/cpuset.cpus + echo 0 > /sys/fs/cgroup/cpuset/cs0/cpuset.mems + + mkdir /sys/fs/cgroup/cpuset/cs1 + echo 2-7 > /sys/fs/cgroup/cpuset/cs1/cpuset.cpus + echo 0 > /sys/fs/cgroup/cpuset/cs1/cpuset.mems + + echo 0 > /sys/fs/cgroup/cpuset/cpuset.sched_load_balance + +Since there *is* CPU capacity asymmetry in the system, the +sched_asym_cpucapacity static key will be enabled. However, the sched_domain +hierarchy of CPUs 0-1 spans a single capacity value: SD_ASYM_CPUCAPACITY isn't +set in that hierarchy, it describes an SMP island and should be treated as such. + +Therefore, the 'canonical' pattern for protecting codepaths that cater to +asymmetric CPU capacities is to: + +- Check the sched_asym_cpucapacity static key +- If it is enabled, then also check for the presence of SD_ASYM_CPUCAPACITY in + the sched_domain hierarchy (if relevant, i.e. the codepath targets a specific + CPU or group thereof) + +5. Capacity aware scheduling implementation +=========================================== + +5.1 CFS +------- + +5.1.1 Capacity fitness +~~~~~~~~~~~~~~~~~~~~~~ + +The main capacity scheduling criterion of CFS is:: + + task_util(p) < capacity(task_cpu(p)) + +This is commonly called the capacity fitness criterion, i.e. CFS must ensure a +task "fits" on its CPU. If it is violated, the task will need to achieve more +work than what its CPU can provide: it will be CPU-bound. + +Furthermore, uclamp lets userspace specify a minimum and a maximum utilization +value for a task, either via sched_setattr() or via the cgroup interface (see +Documentation/admin-guide/cgroup-v2.rst). As its name imply, this can be used to +clamp task_util() in the previous criterion. + +5.1.2 Wakeup CPU selection +~~~~~~~~~~~~~~~~~~~~~~~~~~ + +CFS task wakeup CPU selection follows the capacity fitness criterion described +above. On top of that, uclamp is used to clamp the task utilization values, +which lets userspace have more leverage over the CPU selection of CFS +tasks. IOW, CFS wakeup CPU selection searches for a CPU that satisfies:: + + clamp(task_util(p), task_uclamp_min(p), task_uclamp_max(p)) < capacity(cpu) + +By using uclamp, userspace can e.g. allow a busy loop (100% utilization) to run +on any CPU by giving it a low uclamp.max value. Conversely, it can force a small +periodic task (e.g. 10% utilization) to run on the highest-performance CPUs by +giving it a high uclamp.min value. + +.. note:: + + Wakeup CPU selection in CFS can be eclipsed by Energy Aware Scheduling + (EAS), which is described in Documentation/scheduler/sched-energy.rst. + +5.1.3 Load balancing +~~~~~~~~~~~~~~~~~~~~ + +A pathological case in the wakeup CPU selection occurs when a task rarely +sleeps, if at all - it thus rarely wakes up, if at all. Consider:: + + w == wakeup event + + capacity(CPU0) = C + capacity(CPU1) = C / 3 + + workload on CPU0 + CPU work ^ + | _________ _________ ____ + | | | | | | + +----+----+----+----+----+----+----+----+----+----+-> time + w w w + + workload on CPU1 + CPU work ^ + | ____________________________________________ + | | + +----+----+----+----+----+----+----+----+----+----+-> + w + +This workload should run on CPU0, but if the task either: + +- was improperly scheduled from the start (inaccurate initial + utilization estimation) +- was properly scheduled from the start, but suddenly needs more + processing power + +then it might become CPU-bound, IOW ``task_util(p) > capacity(task_cpu(p))``; +the CPU capacity scheduling criterion is violated, and there may not be any more +wakeup event to fix this up via wakeup CPU selection. + +Tasks that are in this situation are dubbed "misfit" tasks, and the mechanism +put in place to handle this shares the same name. Misfit task migration +leverages the CFS load balancer, more specifically the active load balance part +(which caters to migrating currently running tasks). When load balance happens, +a misfit active load balance will be triggered if a misfit task can be migrated +to a CPU with more capacity than its current one. + +5.2 RT +------ + +5.2.1 Wakeup CPU selection +~~~~~~~~~~~~~~~~~~~~~~~~~~ + +RT task wakeup CPU selection searches for a CPU that satisfies:: + + task_uclamp_min(p) <= capacity(task_cpu(cpu)) + +while still following the usual priority constraints. If none of the candidate +CPUs can satisfy this capacity criterion, then strict priority based scheduling +is followed and CPU capacities are ignored. + +5.3 DL +------ + +5.3.1 Wakeup CPU selection +~~~~~~~~~~~~~~~~~~~~~~~~~~ + +DL task wakeup CPU selection searches for a CPU that satisfies:: + + task_bandwidth(p) < capacity(task_cpu(p)) + +while still respecting the usual bandwidth and deadline constraints. If +none of the candidate CPUs can satisfy this capacity criterion, then the +task will remain on its current CPU. diff --git a/Documentation/scheduler/sched-deadline.rst b/Documentation/scheduler/sched-deadline.rst new file mode 100644 index 000000000..14a2f7bf6 --- /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; + } diff --git a/Documentation/scheduler/sched-design-CFS.rst b/Documentation/scheduler/sched-design-CFS.rst new file mode 100644 index 000000000..a96c72651 --- /dev/null +++ b/Documentation/scheduler/sched-design-CFS.rst @@ -0,0 +1,249 @@ +============= +CFS Scheduler +============= + + +1. OVERVIEW +============ + +CFS stands for "Completely Fair Scheduler," and is the new "desktop" process +scheduler implemented by Ingo Molnar and merged in Linux 2.6.23. It is the +replacement for the previous vanilla scheduler's SCHED_OTHER interactivity +code. + +80% of CFS's design can be summed up in a single sentence: CFS basically models +an "ideal, precise multi-tasking CPU" on real hardware. + +"Ideal multi-tasking CPU" is a (non-existent :-)) CPU that has 100% physical +power and which can run each task at precise equal speed, in parallel, each at +1/nr_running speed. For example: if there are 2 tasks running, then it runs +each at 50% physical power --- i.e., actually in parallel. + +On real hardware, we can run only a single task at once, so we have to +introduce the concept of "virtual runtime." The virtual runtime of a task +specifies when its next timeslice would start execution on the ideal +multi-tasking CPU described above. In practice, the virtual runtime of a task +is its actual runtime normalized to the total number of running tasks. + + + +2. FEW IMPLEMENTATION DETAILS +============================== + +In CFS the virtual runtime is expressed and tracked via the per-task +p->se.vruntime (nanosec-unit) value. This way, it's possible to accurately +timestamp and measure the "expected CPU time" a task should have gotten. + +[ small detail: on "ideal" hardware, at any time all tasks would have the same + p->se.vruntime value --- i.e., tasks would execute simultaneously and no task + would ever get "out of balance" from the "ideal" share of CPU time. ] + +CFS's task picking logic is based on this p->se.vruntime value and it is thus +very simple: it always tries to run the task with the smallest p->se.vruntime +value (i.e., the task which executed least so far). CFS always tries to split +up CPU time between runnable tasks as close to "ideal multitasking hardware" as +possible. + +Most of the rest of CFS's design just falls out of this really simple concept, +with a few add-on embellishments like nice levels, multiprocessing and various +algorithm variants to recognize sleepers. + + + +3. THE RBTREE +============== + +CFS's design is quite radical: it does not use the old data structures for the +runqueues, but it uses a time-ordered rbtree to build a "timeline" of future +task execution, and thus has no "array switch" artifacts (by which both the +previous vanilla scheduler and RSDL/SD are affected). + +CFS also maintains the rq->cfs.min_vruntime value, which is a monotonic +increasing value tracking the smallest vruntime among all tasks in the +runqueue. The total amount of work done by the system is tracked using +min_vruntime; that value is used to place newly activated entities on the left +side of the tree as much as possible. + +The total number of running tasks in the runqueue is accounted through the +rq->cfs.load value, which is the sum of the weights of the tasks queued on the +runqueue. + +CFS maintains a time-ordered rbtree, where all runnable tasks are sorted by the +p->se.vruntime key. CFS picks the "leftmost" task from this tree and sticks to it. +As the system progresses forwards, the executed tasks are put into the tree +more and more to the right --- slowly but surely giving a chance for every task +to become the "leftmost task" and thus get on the CPU within a deterministic +amount of time. + +Summing up, CFS works like this: it runs a task a bit, and when the task +schedules (or a scheduler tick happens) the task's CPU usage is "accounted +for": the (small) time it just spent using the physical CPU is added to +p->se.vruntime. Once p->se.vruntime gets high enough so that another task +becomes the "leftmost task" of the time-ordered rbtree it maintains (plus a +small amount of "granularity" distance relative to the leftmost task so that we +do not over-schedule tasks and trash the cache), then the new leftmost task is +picked and the current task is preempted. + + + +4. SOME FEATURES OF CFS +======================== + +CFS uses nanosecond granularity accounting and does not rely on any jiffies or +other HZ detail. Thus the CFS scheduler has no notion of "timeslices" in the +way the previous scheduler had, and has no heuristics whatsoever. There is +only one central tunable (you have to switch on CONFIG_SCHED_DEBUG): + + /proc/sys/kernel/sched_min_granularity_ns + +which can be used to tune the scheduler from "desktop" (i.e., low latencies) to +"server" (i.e., good batching) workloads. It defaults to a setting suitable +for desktop workloads. SCHED_BATCH is handled by the CFS scheduler module too. + +Due to its design, the CFS scheduler is not prone to any of the "attacks" that +exist today against the heuristics of the stock scheduler: fiftyp.c, thud.c, +chew.c, ring-test.c, massive_intr.c all work fine and do not impact +interactivity and produce the expected behavior. + +The CFS scheduler has a much stronger handling of nice levels and SCHED_BATCH +than the previous vanilla scheduler: both types of workloads are isolated much +more aggressively. + +SMP load-balancing has been reworked/sanitized: the runqueue-walking +assumptions are gone from the load-balancing code now, and iterators of the +scheduling modules are used. The balancing code got quite a bit simpler as a +result. + + + +5. Scheduling policies +====================== + +CFS implements three scheduling policies: + + - SCHED_NORMAL (traditionally called SCHED_OTHER): The scheduling + policy that is used for regular tasks. + + - SCHED_BATCH: Does not preempt nearly as often as regular tasks + would, thereby allowing tasks to run longer and make better use of + caches but at the cost of interactivity. This is well suited for + batch jobs. + + - SCHED_IDLE: This is even weaker than nice 19, but its not a true + idle timer scheduler in order to avoid to get into priority + inversion problems which would deadlock the machine. + +SCHED_FIFO/_RR are implemented in sched/rt.c and are as specified by +POSIX. + +The command chrt from util-linux-ng 2.13.1.1 can set all of these except +SCHED_IDLE. + + + +6. SCHEDULING CLASSES +====================== + +The new CFS scheduler has been designed in such a way to introduce "Scheduling +Classes," an extensible hierarchy of scheduler modules. These modules +encapsulate scheduling policy details and are handled by the scheduler core +without the core code assuming too much about them. + +sched/fair.c implements the CFS scheduler described above. + +sched/rt.c implements SCHED_FIFO and SCHED_RR semantics, in a simpler way than +the previous vanilla scheduler did. It uses 100 runqueues (for all 100 RT +priority levels, instead of 140 in the previous scheduler) and it needs no +expired array. + +Scheduling classes are implemented through the sched_class structure, which +contains hooks to functions that must be called whenever an interesting event +occurs. + +This is the (partial) list of the hooks: + + - enqueue_task(...) + + Called when a task enters a runnable state. + It puts the scheduling entity (task) into the red-black tree and + increments the nr_running variable. + + - dequeue_task(...) + + When a task is no longer runnable, this function is called to keep the + corresponding scheduling entity out of the red-black tree. It decrements + the nr_running variable. + + - yield_task(...) + + This function is basically just a dequeue followed by an enqueue, unless the + compat_yield sysctl is turned on; in that case, it places the scheduling + entity at the right-most end of the red-black tree. + + - check_preempt_curr(...) + + This function checks if a task that entered the runnable state should + preempt the currently running task. + + - pick_next_task(...) + + This function chooses the most appropriate task eligible to run next. + + - set_curr_task(...) + + This function is called when a task changes its scheduling class or changes + its task group. + + - task_tick(...) + + This function is mostly called from time tick functions; it might lead to + process switch. This drives the running preemption. + + + + +7. GROUP SCHEDULER EXTENSIONS TO CFS +===================================== + +Normally, the scheduler operates on individual tasks and strives to provide +fair CPU time to each task. Sometimes, it may be desirable to group tasks and +provide fair CPU time to each such task group. For example, it may be +desirable to first provide fair CPU time to each user on the system and then to +each task belonging to a user. + +CONFIG_CGROUP_SCHED strives to achieve exactly that. It lets tasks to be +grouped and divides CPU time fairly among such groups. + +CONFIG_RT_GROUP_SCHED permits to group real-time (i.e., SCHED_FIFO and +SCHED_RR) tasks. + +CONFIG_FAIR_GROUP_SCHED permits to group CFS (i.e., SCHED_NORMAL and +SCHED_BATCH) tasks. + + These options need CONFIG_CGROUPS to be defined, and let the administrator + create arbitrary groups of tasks, using the "cgroup" pseudo filesystem. See + Documentation/admin-guide/cgroup-v1/cgroups.rst for more information about this filesystem. + +When CONFIG_FAIR_GROUP_SCHED is defined, a "cpu.shares" file is created for each +group created using the pseudo filesystem. See example steps below to create +task groups and modify their CPU share using the "cgroups" pseudo filesystem:: + + # mount -t tmpfs cgroup_root /sys/fs/cgroup + # mkdir /sys/fs/cgroup/cpu + # mount -t cgroup -ocpu none /sys/fs/cgroup/cpu + # cd /sys/fs/cgroup/cpu + + # mkdir multimedia # create "multimedia" group of tasks + # mkdir browser # create "browser" group of tasks + + # #Configure the multimedia group to receive twice the CPU bandwidth + # #that of browser group + + # echo 2048 > multimedia/cpu.shares + # echo 1024 > browser/cpu.shares + + # firefox & # Launch firefox and move it to "browser" group + # echo <firefox_pid> > browser/tasks + + # #Launch gmplayer (or your favourite movie player) + # echo <movie_player_pid> > multimedia/tasks diff --git a/Documentation/scheduler/sched-domains.rst b/Documentation/scheduler/sched-domains.rst new file mode 100644 index 000000000..5c4b7f4f0 --- /dev/null +++ b/Documentation/scheduler/sched-domains.rst @@ -0,0 +1,85 @@ +================= +Scheduler Domains +================= + +Each CPU has a "base" scheduling domain (struct sched_domain). The domain +hierarchy is built from these base domains via the ->parent pointer. ->parent +MUST be NULL terminated, and domain structures should be per-CPU as they are +locklessly updated. + +Each scheduling domain spans a number of CPUs (stored in the ->span field). +A domain's span MUST be a superset of it child's span (this restriction could +be relaxed if the need arises), and a base domain for CPU i MUST span at least +i. The top domain for each CPU will generally span all CPUs in the system +although strictly it doesn't have to, but this could lead to a case where some +CPUs will never be given tasks to run unless the CPUs allowed mask is +explicitly set. A sched domain's span means "balance process load among these +CPUs". + +Each scheduling domain must have one or more CPU groups (struct sched_group) +which are organised as a circular one way linked list from the ->groups +pointer. The union of cpumasks of these groups MUST be the same as the +domain's span. The group pointed to by the ->groups pointer MUST contain the CPU +to which the domain belongs. Groups may be shared among CPUs as they contain +read only data after they have been set up. The intersection of cpumasks from +any two of these groups may be non empty. If this is the case the SD_OVERLAP +flag is set on the corresponding scheduling domain and its groups may not be +shared between CPUs. + +Balancing within a sched domain occurs between groups. That is, each group +is treated as one entity. The load of a group is defined as the sum of the +load of each of its member CPUs, and only when the load of a group becomes +out of balance are tasks moved between groups. + +In kernel/sched/core.c, trigger_load_balance() is run periodically on each CPU +through scheduler_tick(). It raises a softirq after the next regularly scheduled +rebalancing event for the current runqueue has arrived. The actual load +balancing workhorse, run_rebalance_domains()->rebalance_domains(), is then run +in softirq context (SCHED_SOFTIRQ). + +The latter function takes two arguments: the current CPU and whether it was idle +at the time the scheduler_tick() happened and iterates over all sched domains +our CPU is on, starting from its base domain and going up the ->parent chain. +While doing that, it checks to see if the current domain has exhausted its +rebalance interval. If so, it runs load_balance() on that domain. It then checks +the parent sched_domain (if it exists), and the parent of the parent and so +forth. + +Initially, load_balance() finds the busiest group in the current sched domain. +If it succeeds, it looks for the busiest runqueue of all the CPUs' runqueues in +that group. If it manages to find such a runqueue, it locks both our initial +CPU's runqueue and the newly found busiest one and starts moving tasks from it +to our runqueue. The exact number of tasks amounts to an imbalance previously +computed while iterating over this sched domain's groups. + +Implementing sched domains +========================== + +The "base" domain will "span" the first level of the hierarchy. In the case +of SMT, you'll span all siblings of the physical CPU, with each group being +a single virtual CPU. + +In SMP, the parent of the base domain will span all physical CPUs in the +node. Each group being a single physical CPU. Then with NUMA, the parent +of the SMP domain will span the entire machine, with each group having the +cpumask of a node. Or, you could do multi-level NUMA or Opteron, for example, +might have just one domain covering its one NUMA level. + +The implementor should read comments in include/linux/sched.h: +struct sched_domain fields, SD_FLAG_*, SD_*_INIT to get an idea of +the specifics and what to tune. + +Architectures may retain the regular override the default SD_*_INIT flags +while using the generic domain builder in kernel/sched/core.c if they wish to +retain the traditional SMT->SMP->NUMA topology (or some subset of that). This +can be done by #define'ing ARCH_HASH_SCHED_TUNE. + +Alternatively, the architecture may completely override the generic domain +builder by #define'ing ARCH_HASH_SCHED_DOMAIN, and exporting your +arch_init_sched_domains function. This function will attach domains to all +CPUs using cpu_attach_domain. + +The sched-domains debugging infrastructure can be enabled by enabling +CONFIG_SCHED_DEBUG. This enables an error checking parse of the sched domains +which should catch most possible errors (described above). It also prints out +the domain structure in a visual format. diff --git a/Documentation/scheduler/sched-energy.rst b/Documentation/scheduler/sched-energy.rst new file mode 100644 index 000000000..001e09c95 --- /dev/null +++ b/Documentation/scheduler/sched-energy.rst @@ -0,0 +1,422 @@ +======================= +Energy Aware Scheduling +======================= + +1. Introduction +--------------- + +Energy Aware Scheduling (or EAS) gives the scheduler the ability to predict +the impact of its decisions on the energy consumed by CPUs. EAS relies on an +Energy Model (EM) of the CPUs to select an energy efficient CPU for each task, +with a minimal impact on throughput. This document aims at providing an +introduction on how EAS works, what are the main design decisions behind it, and +details what is needed to get it to run. + +Before going any further, please note that at the time of writing:: + + /!\ EAS does not support platforms with symmetric CPU topologies /!\ + +EAS operates only on heterogeneous CPU topologies (such as Arm big.LITTLE) +because this is where the potential for saving energy through scheduling is +the highest. + +The actual EM used by EAS is _not_ maintained by the scheduler, but by a +dedicated framework. For details about this framework and what it provides, +please refer to its documentation (see Documentation/power/energy-model.rst). + + +2. Background and Terminology +----------------------------- + +To make it clear from the start: + - energy = [joule] (resource like a battery on powered devices) + - power = energy/time = [joule/second] = [watt] + +The goal of EAS is to minimize energy, while still getting the job done. That +is, we want to maximize:: + + performance [inst/s] + -------------------- + power [W] + +which is equivalent to minimizing:: + + energy [J] + ----------- + instruction + +while still getting 'good' performance. It is essentially an alternative +optimization objective to the current performance-only objective for the +scheduler. This alternative considers two objectives: energy-efficiency and +performance. + +The idea behind introducing an EM is to allow the scheduler to evaluate the +implications of its decisions rather than blindly applying energy-saving +techniques that may have positive effects only on some platforms. At the same +time, the EM must be as simple as possible to minimize the scheduler latency +impact. + +In short, EAS changes the way CFS tasks are assigned to CPUs. When it is time +for the scheduler to decide where a task should run (during wake-up), the EM +is used to break the tie between several good CPU candidates and pick the one +that is predicted to yield the best energy consumption without harming the +system's throughput. The predictions made by EAS rely on specific elements of +knowledge about the platform's topology, which include the 'capacity' of CPUs, +and their respective energy costs. + + +3. Topology information +----------------------- + +EAS (as well as the rest of the scheduler) uses the notion of 'capacity' to +differentiate CPUs with different computing throughput. The 'capacity' of a CPU +represents the amount of work it can absorb when running at its highest +frequency compared to the most capable CPU of the system. Capacity values are +normalized in a 1024 range, and are comparable with the utilization signals of +tasks and CPUs computed by the Per-Entity Load Tracking (PELT) mechanism. Thanks +to capacity and utilization values, EAS is able to estimate how big/busy a +task/CPU is, and to take this into consideration when evaluating performance vs +energy trade-offs. The capacity of CPUs is provided via arch-specific code +through the arch_scale_cpu_capacity() callback. + +The rest of platform knowledge used by EAS is directly read from the Energy +Model (EM) framework. The EM of a platform is composed of a power cost table +per 'performance domain' in the system (see Documentation/power/energy-model.rst +for futher details about performance domains). + +The scheduler manages references to the EM objects in the topology code when the +scheduling domains are built, or re-built. For each root domain (rd), the +scheduler maintains a singly linked list of all performance domains intersecting +the current rd->span. Each node in the list contains a pointer to a struct +em_perf_domain as provided by the EM framework. + +The lists are attached to the root domains in order to cope with exclusive +cpuset configurations. Since the boundaries of exclusive cpusets do not +necessarily match those of performance domains, the lists of different root +domains can contain duplicate elements. + +Example 1. + Let us consider a platform with 12 CPUs, split in 3 performance domains + (pd0, pd4 and pd8), organized as follows:: + + CPUs: 0 1 2 3 4 5 6 7 8 9 10 11 + PDs: |--pd0--|--pd4--|---pd8---| + RDs: |----rd1----|-----rd2-----| + + Now, consider that userspace decided to split the system with two + exclusive cpusets, hence creating two independent root domains, each + containing 6 CPUs. The two root domains are denoted rd1 and rd2 in the + above figure. Since pd4 intersects with both rd1 and rd2, it will be + present in the linked list '->pd' attached to each of them: + + * rd1->pd: pd0 -> pd4 + * rd2->pd: pd4 -> pd8 + + Please note that the scheduler will create two duplicate list nodes for + pd4 (one for each list). However, both just hold a pointer to the same + shared data structure of the EM framework. + +Since the access to these lists can happen concurrently with hotplug and other +things, they are protected by RCU, like the rest of topology structures +manipulated by the scheduler. + +EAS also maintains a static key (sched_energy_present) which is enabled when at +least one root domain meets all conditions for EAS to start. Those conditions +are summarized in Section 6. + + +4. Energy-Aware task placement +------------------------------ + +EAS overrides the CFS task wake-up balancing code. It uses the EM of the +platform and the PELT signals to choose an energy-efficient target CPU during +wake-up balance. When EAS is enabled, select_task_rq_fair() calls +find_energy_efficient_cpu() to do the placement decision. This function looks +for the CPU with the highest spare capacity (CPU capacity - CPU utilization) in +each performance domain since it is the one which will allow us to keep the +frequency the lowest. Then, the function checks if placing the task there could +save energy compared to leaving it on prev_cpu, i.e. the CPU where the task ran +in its previous activation. + +find_energy_efficient_cpu() uses compute_energy() to estimate what will be the +energy consumed by the system if the waking task was migrated. compute_energy() +looks at the current utilization landscape of the CPUs and adjusts it to +'simulate' the task migration. The EM framework provides the em_pd_energy() API +which computes the expected energy consumption of each performance domain for +the given utilization landscape. + +An example of energy-optimized task placement decision is detailed below. + +Example 2. + Let us consider a (fake) platform with 2 independent performance domains + composed of two CPUs each. CPU0 and CPU1 are little CPUs; CPU2 and CPU3 + are big. + + The scheduler must decide where to place a task P whose util_avg = 200 + and prev_cpu = 0. + + The current utilization landscape of the CPUs is depicted on the graph + below. CPUs 0-3 have a util_avg of 400, 100, 600 and 500 respectively + Each performance domain has three Operating Performance Points (OPPs). + The CPU capacity and power cost associated with each OPP is listed in + the Energy Model table. The util_avg of P is shown on the figures + below as 'PP':: + + CPU util. + 1024 - - - - - - - Energy Model + +-----------+-------------+ + | Little | Big | + 768 ============= +-----+-----+------+------+ + | Cap | Pwr | Cap | Pwr | + +-----+-----+------+------+ + 512 =========== - ##- - - - - | 170 | 50 | 512 | 400 | + ## ## | 341 | 150 | 768 | 800 | + 341 -PP - - - - ## ## | 512 | 300 | 1024 | 1700 | + PP ## ## +-----+-----+------+------+ + 170 -## - - - - ## ## + ## ## ## ## + ------------ ------------- + CPU0 CPU1 CPU2 CPU3 + + Current OPP: ===== Other OPP: - - - util_avg (100 each): ## + + + find_energy_efficient_cpu() will first look for the CPUs with the + maximum spare capacity in the two performance domains. In this example, + CPU1 and CPU3. Then it will estimate the energy of the system if P was + placed on either of them, and check if that would save some energy + compared to leaving P on CPU0. EAS assumes that OPPs follow utilization + (which is coherent with the behaviour of the schedutil CPUFreq + governor, see Section 6. for more details on this topic). + + **Case 1. P is migrated to CPU1**:: + + 1024 - - - - - - - + + Energy calculation: + 768 ============= * CPU0: 200 / 341 * 150 = 88 + * CPU1: 300 / 341 * 150 = 131 + * CPU2: 600 / 768 * 800 = 625 + 512 - - - - - - - ##- - - - - * CPU3: 500 / 768 * 800 = 520 + ## ## => total_energy = 1364 + 341 =========== ## ## + PP ## ## + 170 -## - - PP- ## ## + ## ## ## ## + ------------ ------------- + CPU0 CPU1 CPU2 CPU3 + + + **Case 2. P is migrated to CPU3**:: + + 1024 - - - - - - - + + Energy calculation: + 768 ============= * CPU0: 200 / 341 * 150 = 88 + * CPU1: 100 / 341 * 150 = 43 + PP * CPU2: 600 / 768 * 800 = 625 + 512 - - - - - - - ##- - -PP - * CPU3: 700 / 768 * 800 = 729 + ## ## => total_energy = 1485 + 341 =========== ## ## + ## ## + 170 -## - - - - ## ## + ## ## ## ## + ------------ ------------- + CPU0 CPU1 CPU2 CPU3 + + + **Case 3. P stays on prev_cpu / CPU 0**:: + + 1024 - - - - - - - + + Energy calculation: + 768 ============= * CPU0: 400 / 512 * 300 = 234 + * CPU1: 100 / 512 * 300 = 58 + * CPU2: 600 / 768 * 800 = 625 + 512 =========== - ##- - - - - * CPU3: 500 / 768 * 800 = 520 + ## ## => total_energy = 1437 + 341 -PP - - - - ## ## + PP ## ## + 170 -## - - - - ## ## + ## ## ## ## + ------------ ------------- + CPU0 CPU1 CPU2 CPU3 + + + From these calculations, the Case 1 has the lowest total energy. So CPU 1 + is be the best candidate from an energy-efficiency standpoint. + +Big CPUs are generally more power hungry than the little ones and are thus used +mainly when a task doesn't fit the littles. However, little CPUs aren't always +necessarily more energy-efficient than big CPUs. For some systems, the high OPPs +of the little CPUs can be less energy-efficient than the lowest OPPs of the +bigs, for example. So, if the little CPUs happen to have enough utilization at +a specific point in time, a small task waking up at that moment could be better +of executing on the big side in order to save energy, even though it would fit +on the little side. + +And even in the case where all OPPs of the big CPUs are less energy-efficient +than those of the little, using the big CPUs for a small task might still, under +specific conditions, save energy. Indeed, placing a task on a little CPU can +result in raising the OPP of the entire performance domain, and that will +increase the cost of the tasks already running there. If the waking task is +placed on a big CPU, its own execution cost might be higher than if it was +running on a little, but it won't impact the other tasks of the little CPUs +which will keep running at a lower OPP. So, when considering the total energy +consumed by CPUs, the extra cost of running that one task on a big core can be +smaller than the cost of raising the OPP on the little CPUs for all the other +tasks. + +The examples above would be nearly impossible to get right in a generic way, and +for all platforms, without knowing the cost of running at different OPPs on all +CPUs of the system. Thanks to its EM-based design, EAS should cope with them +correctly without too many troubles. However, in order to ensure a minimal +impact on throughput for high-utilization scenarios, EAS also implements another +mechanism called 'over-utilization'. + + +5. Over-utilization +------------------- + +From a general standpoint, the use-cases where EAS can help the most are those +involving a light/medium CPU utilization. Whenever long CPU-bound tasks are +being run, they will require all of the available CPU capacity, and there isn't +much that can be done by the scheduler to save energy without severly harming +throughput. In order to avoid hurting performance with EAS, CPUs are flagged as +'over-utilized' as soon as they are used at more than 80% of their compute +capacity. As long as no CPUs are over-utilized in a root domain, load balancing +is disabled and EAS overridess the wake-up balancing code. EAS is likely to load +the most energy efficient CPUs of the system more than the others if that can be +done without harming throughput. So, the load-balancer is disabled to prevent +it from breaking the energy-efficient task placement found by EAS. It is safe to +do so when the system isn't overutilized since being below the 80% tipping point +implies that: + + a. there is some idle time on all CPUs, so the utilization signals used by + EAS are likely to accurately represent the 'size' of the various tasks + in the system; + b. all tasks should already be provided with enough CPU capacity, + regardless of their nice values; + c. since there is spare capacity all tasks must be blocking/sleeping + regularly and balancing at wake-up is sufficient. + +As soon as one CPU goes above the 80% tipping point, at least one of the three +assumptions above becomes incorrect. In this scenario, the 'overutilized' flag +is raised for the entire root domain, EAS is disabled, and the load-balancer is +re-enabled. By doing so, the scheduler falls back onto load-based algorithms for +wake-up and load balance under CPU-bound conditions. This provides a better +respect of the nice values of tasks. + +Since the notion of overutilization largely relies on detecting whether or not +there is some idle time in the system, the CPU capacity 'stolen' by higher +(than CFS) scheduling classes (as well as IRQ) must be taken into account. As +such, the detection of overutilization accounts for the capacity used not only +by CFS tasks, but also by the other scheduling classes and IRQ. + + +6. Dependencies and requirements for EAS +---------------------------------------- + +Energy Aware Scheduling depends on the CPUs of the system having specific +hardware properties and on other features of the kernel being enabled. This +section lists these dependencies and provides hints as to how they can be met. + + +6.1 - Asymmetric CPU topology +^^^^^^^^^^^^^^^^^^^^^^^^^^^^^ + + +As mentioned in the introduction, EAS is only supported on platforms with +asymmetric CPU topologies for now. This requirement is checked at run-time by +looking for the presence of the SD_ASYM_CPUCAPACITY flag when the scheduling +domains are built. + +See Documentation/scheduler/sched-capacity.rst for requirements to be met for this +flag to be set in the sched_domain hierarchy. + +Please note that EAS is not fundamentally incompatible with SMP, but no +significant savings on SMP platforms have been observed yet. This restriction +could be amended in the future if proven otherwise. + + +6.2 - Energy Model presence +^^^^^^^^^^^^^^^^^^^^^^^^^^^ + +EAS uses the EM of a platform to estimate the impact of scheduling decisions on +energy. So, your platform must provide power cost tables to the EM framework in +order to make EAS start. To do so, please refer to documentation of the +independent EM framework in Documentation/power/energy-model.rst. + +Please also note that the scheduling domains need to be re-built after the +EM has been registered in order to start EAS. + + +6.3 - Energy Model complexity +^^^^^^^^^^^^^^^^^^^^^^^^^^^^^ + +The task wake-up path is very latency-sensitive. When the EM of a platform is +too complex (too many CPUs, too many performance domains, too many performance +states, ...), the cost of using it in the wake-up path can become prohibitive. +The energy-aware wake-up algorithm has a complexity of: + + C = Nd * (Nc + Ns) + +with: Nd the number of performance domains; Nc the number of CPUs; and Ns the +total number of OPPs (ex: for two perf. domains with 4 OPPs each, Ns = 8). + +A complexity check is performed at the root domain level, when scheduling +domains are built. EAS will not start on a root domain if its C happens to be +higher than the completely arbitrary EM_MAX_COMPLEXITY threshold (2048 at the +time of writing). + +If you really want to use EAS but the complexity of your platform's Energy +Model is too high to be used with a single root domain, you're left with only +two possible options: + + 1. split your system into separate, smaller, root domains using exclusive + cpusets and enable EAS locally on each of them. This option has the + benefit to work out of the box but the drawback of preventing load + balance between root domains, which can result in an unbalanced system + overall; + 2. submit patches to reduce the complexity of the EAS wake-up algorithm, + hence enabling it to cope with larger EMs in reasonable time. + + +6.4 - Schedutil governor +^^^^^^^^^^^^^^^^^^^^^^^^ + +EAS tries to predict at which OPP will the CPUs be running in the close future +in order to estimate their energy consumption. To do so, it is assumed that OPPs +of CPUs follow their utilization. + +Although it is very difficult to provide hard guarantees regarding the accuracy +of this assumption in practice (because the hardware might not do what it is +told to do, for example), schedutil as opposed to other CPUFreq governors at +least _requests_ frequencies calculated using the utilization signals. +Consequently, the only sane governor to use together with EAS is schedutil, +because it is the only one providing some degree of consistency between +frequency requests and energy predictions. + +Using EAS with any other governor than schedutil is not supported. + + +6.5 Scale-invariant utilization signals +^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^ + +In order to make accurate prediction across CPUs and for all performance +states, EAS needs frequency-invariant and CPU-invariant PELT signals. These can +be obtained using the architecture-defined arch_scale{cpu,freq}_capacity() +callbacks. + +Using EAS on a platform that doesn't implement these two callbacks is not +supported. + + +6.6 Multithreading (SMT) +^^^^^^^^^^^^^^^^^^^^^^^^ + +EAS in its current form is SMT unaware and is not able to leverage +multithreaded hardware to save energy. EAS considers threads as independent +CPUs, which can actually be counter-productive for both performance and energy. + +EAS on SMT is not supported. diff --git a/Documentation/scheduler/sched-nice-design.rst b/Documentation/scheduler/sched-nice-design.rst new file mode 100644 index 000000000..0571f1b47 --- /dev/null +++ b/Documentation/scheduler/sched-nice-design.rst @@ -0,0 +1,112 @@ +===================== +Scheduler Nice Design +===================== + +This document explains the thinking about the revamped and streamlined +nice-levels implementation in the new Linux scheduler. + +Nice levels were always pretty weak under Linux and people continuously +pestered us to make nice +19 tasks use up much less CPU time. + +Unfortunately that was not that easy to implement under the old +scheduler, (otherwise we'd have done it long ago) because nice level +support was historically coupled to timeslice length, and timeslice +units were driven by the HZ tick, so the smallest timeslice was 1/HZ. + +In the O(1) scheduler (in 2003) we changed negative nice levels to be +much stronger than they were before in 2.4 (and people were happy about +that change), and we also intentionally calibrated the linear timeslice +rule so that nice +19 level would be _exactly_ 1 jiffy. To better +understand it, the timeslice graph went like this (cheesy ASCII art +alert!):: + + + A + \ | [timeslice length] + \ | + \ | + \ | + \ | + \|___100msecs + |^ . _ + | ^ . _ + | ^ . _ + -*----------------------------------*-----> [nice level] + -20 | +19 + | + | + +So that if someone wanted to really renice tasks, +19 would give a much +bigger hit than the normal linear rule would do. (The solution of +changing the ABI to extend priorities was discarded early on.) + +This approach worked to some degree for some time, but later on with +HZ=1000 it caused 1 jiffy to be 1 msec, which meant 0.1% CPU usage which +we felt to be a bit excessive. Excessive _not_ because it's too small of +a CPU utilization, but because it causes too frequent (once per +millisec) rescheduling. (and would thus trash the cache, etc. Remember, +this was long ago when hardware was weaker and caches were smaller, and +people were running number crunching apps at nice +19.) + +So for HZ=1000 we changed nice +19 to 5msecs, because that felt like the +right minimal granularity - and this translates to 5% CPU utilization. +But the fundamental HZ-sensitive property for nice+19 still remained, +and we never got a single complaint about nice +19 being too _weak_ in +terms of CPU utilization, we only got complaints about it (still) being +too _strong_ :-) + +To sum it up: we always wanted to make nice levels more consistent, but +within the constraints of HZ and jiffies and their nasty design level +coupling to timeslices and granularity it was not really viable. + +The second (less frequent but still periodically occurring) complaint +about Linux's nice level support was its assymetry around the origo +(which you can see demonstrated in the picture above), or more +accurately: the fact that nice level behavior depended on the _absolute_ +nice level as well, while the nice API itself is fundamentally +"relative": + + int nice(int inc); + + asmlinkage long sys_nice(int increment) + +(the first one is the glibc API, the second one is the syscall API.) +Note that the 'inc' is relative to the current nice level. Tools like +bash's "nice" command mirror this relative API. + +With the old scheduler, if you for example started a niced task with +1 +and another task with +2, the CPU split between the two tasks would +depend on the nice level of the parent shell - if it was at nice -10 the +CPU split was different than if it was at +5 or +10. + +A third complaint against Linux's nice level support was that negative +nice levels were not 'punchy enough', so lots of people had to resort to +run audio (and other multimedia) apps under RT priorities such as +SCHED_FIFO. But this caused other problems: SCHED_FIFO is not starvation +proof, and a buggy SCHED_FIFO app can also lock up the system for good. + +The new scheduler in v2.6.23 addresses all three types of complaints: + +To address the first complaint (of nice levels being not "punchy" +enough), the scheduler was decoupled from 'time slice' and HZ concepts +(and granularity was made a separate concept from nice levels) and thus +it was possible to implement better and more consistent nice +19 +support: with the new scheduler nice +19 tasks get a HZ-independent +1.5%, instead of the variable 3%-5%-9% range they got in the old +scheduler. + +To address the second complaint (of nice levels not being consistent), +the new scheduler makes nice(1) have the same CPU utilization effect on +tasks, regardless of their absolute nice levels. So on the new +scheduler, running a nice +10 and a nice 11 task has the same CPU +utilization "split" between them as running a nice -5 and a nice -4 +task. (one will get 55% of the CPU, the other 45%.) That is why nice +levels were changed to be "multiplicative" (or exponential) - that way +it does not matter which nice level you start out from, the 'relative +result' will always be the same. + +The third complaint (of negative nice levels not being "punchy" enough +and forcing audio apps to run under the more dangerous SCHED_FIFO +scheduling policy) is addressed by the new scheduler almost +automatically: stronger negative nice levels are an automatic +side-effect of the recalibrated dynamic range of nice levels. diff --git a/Documentation/scheduler/sched-pelt.c b/Documentation/scheduler/sched-pelt.c new file mode 100644 index 000000000..7238b3559 --- /dev/null +++ b/Documentation/scheduler/sched-pelt.c @@ -0,0 +1,109 @@ +/* + * The following program is used to generate the constants for + * computing sched averages. + * + * ============================================================== + * C program (compile with -lm) + * ============================================================== + */ + +#include <math.h> +#include <stdio.h> + +#define HALFLIFE 32 +#define SHIFT 32 + +double y; + +void calc_runnable_avg_yN_inv(void) +{ + int i; + unsigned int x; + + /* To silence -Wunused-but-set-variable warnings. */ + printf("static const u32 runnable_avg_yN_inv[] __maybe_unused = {"); + for (i = 0; i < HALFLIFE; i++) { + x = ((1UL<<32)-1)*pow(y, i); + + if (i % 6 == 0) printf("\n\t"); + printf("0x%8x, ", x); + } + printf("\n};\n\n"); +} + +int sum = 1024; + +void calc_runnable_avg_yN_sum(void) +{ + int i; + + printf("static const u32 runnable_avg_yN_sum[] = {\n\t 0,"); + for (i = 1; i <= HALFLIFE; i++) { + if (i == 1) + sum *= y; + else + sum = sum*y + 1024*y; + + if (i % 11 == 0) + printf("\n\t"); + + printf("%5d,", sum); + } + printf("\n};\n\n"); +} + +int n = -1; +/* first period */ +long max = 1024; + +void calc_converged_max(void) +{ + long last = 0, y_inv = ((1UL<<32)-1)*y; + + for (; ; n++) { + if (n > -1) + max = ((max*y_inv)>>SHIFT) + 1024; + /* + * This is the same as: + * max = max*y + 1024; + */ + + if (last == max) + break; + + last = max; + } + n--; + printf("#define LOAD_AVG_PERIOD %d\n", HALFLIFE); + printf("#define LOAD_AVG_MAX %ld\n", max); +// printf("#define LOAD_AVG_MAX_N %d\n\n", n); +} + +void calc_accumulated_sum_32(void) +{ + int i, x = sum; + + printf("static const u32 __accumulated_sum_N32[] = {\n\t 0,"); + for (i = 1; i <= n/HALFLIFE+1; i++) { + if (i > 1) + x = x/2 + sum; + + if (i % 6 == 0) + printf("\n\t"); + + printf("%6d,", x); + } + printf("\n};\n\n"); +} + +void main(void) +{ + printf("/* Generated by Documentation/scheduler/sched-pelt; do not modify. */\n\n"); + + y = pow(0.5, 1/(double)HALFLIFE); + + calc_runnable_avg_yN_inv(); +// calc_runnable_avg_yN_sum(); + calc_converged_max(); +// calc_accumulated_sum_32(); +} diff --git a/Documentation/scheduler/sched-rt-group.rst b/Documentation/scheduler/sched-rt-group.rst new file mode 100644 index 000000000..655a096ec --- /dev/null +++ b/Documentation/scheduler/sched-rt-group.rst @@ -0,0 +1,185 @@ +========================== +Real-Time group scheduling +========================== + +.. CONTENTS + + 0. WARNING + 1. Overview + 1.1 The problem + 1.2 The solution + 2. The interface + 2.1 System-wide settings + 2.2 Default behaviour + 2.3 Basis for grouping tasks + 3. Future plans + + +0. WARNING +========== + + Fiddling with these settings can result in an unstable system, the knobs are + root only and assumes root knows what he is doing. + +Most notable: + + * very small values in sched_rt_period_us can result in an unstable + system when the period is smaller than either the available hrtimer + resolution, or the time it takes to handle the budget refresh itself. + + * very small values in sched_rt_runtime_us can result in an unstable + system when the runtime is so small the system has difficulty making + forward progress (NOTE: the migration thread and kstopmachine both + are real-time processes). + +1. Overview +=========== + + +1.1 The problem +--------------- + +Realtime scheduling is all about determinism, a group has to be able to rely on +the amount of bandwidth (eg. CPU time) being constant. In order to schedule +multiple groups of realtime tasks, each group must be assigned a fixed portion +of the CPU time available. Without a minimum guarantee a realtime group can +obviously fall short. A fuzzy upper limit is of no use since it cannot be +relied upon. Which leaves us with just the single fixed portion. + +1.2 The solution +---------------- + +CPU time is divided by means of specifying how much time can be spent running +in a given period. We allocate this "run time" for each realtime group which +the other realtime groups will not be permitted to use. + +Any time not allocated to a realtime group will be used to run normal priority +tasks (SCHED_OTHER). Any allocated run time not used will also be picked up by +SCHED_OTHER. + +Let's consider an example: a frame fixed realtime renderer must deliver 25 +frames a second, which yields a period of 0.04s per frame. Now say it will also +have to play some music and respond to input, leaving it with around 80% CPU +time dedicated for the graphics. We can then give this group a run time of 0.8 +* 0.04s = 0.032s. + +This way the graphics group will have a 0.04s period with a 0.032s run time +limit. Now if the audio thread needs to refill the DMA buffer every 0.005s, but +needs only about 3% CPU time to do so, it can do with a 0.03 * 0.005s = +0.00015s. So this group can be scheduled with a period of 0.005s and a run time +of 0.00015s. + +The remaining CPU time will be used for user input and other tasks. Because +realtime tasks have explicitly allocated the CPU time they need to perform +their tasks, buffer underruns in the graphics or audio can be eliminated. + +NOTE: the above example is not fully implemented yet. We still +lack an EDF scheduler to make non-uniform periods usable. + + +2. The Interface +================ + + +2.1 System wide settings +------------------------ + +The system wide settings are configured under the /proc virtual file system: + +/proc/sys/kernel/sched_rt_period_us: + The scheduling period that is equivalent to 100% CPU bandwidth + +/proc/sys/kernel/sched_rt_runtime_us: + A global limit on how much time realtime scheduling may use. Even without + CONFIG_RT_GROUP_SCHED enabled, this will limit time reserved to realtime + processes. With CONFIG_RT_GROUP_SCHED it signifies the total bandwidth + available to all realtime groups. + + * Time is specified in us because the interface is s32. This gives an + operating range from 1us to about 35 minutes. + * sched_rt_period_us takes values from 1 to INT_MAX. + * sched_rt_runtime_us takes values from -1 to (INT_MAX - 1). + * A run time of -1 specifies runtime == period, ie. no limit. + + +2.2 Default behaviour +--------------------- + +The default values for sched_rt_period_us (1000000 or 1s) and +sched_rt_runtime_us (950000 or 0.95s). This gives 0.05s to be used by +SCHED_OTHER (non-RT tasks). These defaults were chosen so that a run-away +realtime tasks will not lock up the machine but leave a little time to recover +it. By setting runtime to -1 you'd get the old behaviour back. + +By default all bandwidth is assigned to the root group and new groups get the +period from /proc/sys/kernel/sched_rt_period_us and a run time of 0. If you +want to assign bandwidth to another group, reduce the root group's bandwidth +and assign some or all of the difference to another group. + +Realtime group scheduling means you have to assign a portion of total CPU +bandwidth to the group before it will accept realtime tasks. Therefore you will +not be able to run realtime tasks as any user other than root until you have +done that, even if the user has the rights to run processes with realtime +priority! + + +2.3 Basis for grouping tasks +---------------------------- + +Enabling CONFIG_RT_GROUP_SCHED lets you explicitly allocate real +CPU bandwidth to task groups. + +This uses the cgroup virtual file system and "<cgroup>/cpu.rt_runtime_us" +to control the CPU time reserved for each control group. + +For more information on working with control groups, you should read +Documentation/admin-guide/cgroup-v1/cgroups.rst as well. + +Group settings are checked against the following limits in order to keep the +configuration schedulable: + + \Sum_{i} runtime_{i} / global_period <= global_runtime / global_period + +For now, this can be simplified to just the following (but see Future plans): + + \Sum_{i} runtime_{i} <= global_runtime + + +3. Future plans +=============== + +There is work in progress to make the scheduling period for each group +("<cgroup>/cpu.rt_period_us") configurable as well. + +The constraint on the period is that a subgroup must have a smaller or +equal period to its parent. But realistically its not very useful _yet_ +as its prone to starvation without deadline scheduling. + +Consider two sibling groups A and B; both have 50% bandwidth, but A's +period is twice the length of B's. + +* group A: period=100000us, runtime=50000us + + - this runs for 0.05s once every 0.1s + +* group B: period= 50000us, runtime=25000us + + - this runs for 0.025s twice every 0.1s (or once every 0.05 sec). + +This means that currently a while (1) loop in A will run for the full period of +B and can starve B's tasks (assuming they are of lower priority) for a whole +period. + +The next project will be SCHED_EDF (Earliest Deadline First scheduling) to bring +full deadline scheduling to the linux kernel. Deadline scheduling the above +groups and treating end of the period as a deadline will ensure that they both +get their allocated time. + +Implementing SCHED_EDF might take a while to complete. Priority Inheritance is +the biggest challenge as the current linux PI infrastructure is geared towards +the limited static priority levels 0-99. With deadline scheduling you need to +do deadline inheritance (since priority is inversely proportional to the +deadline delta (deadline - now)). + +This means the whole PI machinery will have to be reworked - and that is one of +the most complex pieces of code we have. diff --git a/Documentation/scheduler/sched-stats.rst b/Documentation/scheduler/sched-stats.rst new file mode 100644 index 000000000..dd9b99a02 --- /dev/null +++ b/Documentation/scheduler/sched-stats.rst @@ -0,0 +1,167 @@ +==================== +Scheduler Statistics +==================== + +Version 15 of schedstats dropped counters for some sched_yield: +yld_exp_empty, yld_act_empty and yld_both_empty. Otherwise, it is +identical to version 14. + +Version 14 of schedstats includes support for sched_domains, which hit the +mainline kernel in 2.6.20 although it is identical to the stats from version +12 which was in the kernel from 2.6.13-2.6.19 (version 13 never saw a kernel +release). Some counters make more sense to be per-runqueue; other to be +per-domain. Note that domains (and their associated information) will only +be pertinent and available on machines utilizing CONFIG_SMP. + +In version 14 of schedstat, there is at least one level of domain +statistics for each cpu listed, and there may well be more than one +domain. Domains have no particular names in this implementation, but +the highest numbered one typically arbitrates balancing across all the +cpus on the machine, while domain0 is the most tightly focused domain, +sometimes balancing only between pairs of cpus. At this time, there +are no architectures which need more than three domain levels. The first +field in the domain stats is a bit map indicating which cpus are affected +by that domain. + +These fields are counters, and only increment. Programs which make use +of these will need to start with a baseline observation and then calculate +the change in the counters at each subsequent observation. A perl script +which does this for many of the fields is available at + + http://eaglet.pdxhosts.com/rick/linux/schedstat/ + +Note that any such script will necessarily be version-specific, as the main +reason to change versions is changes in the output format. For those wishing +to write their own scripts, the fields are described here. + +CPU statistics +-------------- +cpu<N> 1 2 3 4 5 6 7 8 9 + +First field is a sched_yield() statistic: + + 1) # of times sched_yield() was called + +Next three are schedule() statistics: + + 2) This field is a legacy array expiration count field used in the O(1) + scheduler. We kept it for ABI compatibility, but it is always set to zero. + 3) # of times schedule() was called + 4) # of times schedule() left the processor idle + +Next two are try_to_wake_up() statistics: + + 5) # of times try_to_wake_up() was called + 6) # of times try_to_wake_up() was called to wake up the local cpu + +Next three are statistics describing scheduling latency: + + 7) sum of all time spent running by tasks on this processor (in jiffies) + 8) sum of all time spent waiting to run by tasks on this processor (in + jiffies) + 9) # of timeslices run on this cpu + + +Domain statistics +----------------- +One of these is produced per domain for each cpu described. (Note that if +CONFIG_SMP is not defined, *no* domains are utilized and these lines +will not appear in the output.) + +domain<N> <cpumask> 1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16 17 18 19 20 21 22 23 24 25 26 27 28 29 30 31 32 33 34 35 36 + +The first field is a bit mask indicating what cpus this domain operates over. + +The next 24 are a variety of load_balance() statistics in grouped into types +of idleness (idle, busy, and newly idle): + + 1) # of times in this domain load_balance() was called when the + cpu was idle + 2) # of times in this domain load_balance() checked but found + the load did not require balancing when the cpu was idle + 3) # of times in this domain load_balance() tried to move one or + more tasks and failed, when the cpu was idle + 4) sum of imbalances discovered (if any) with each call to + load_balance() in this domain when the cpu was idle + 5) # of times in this domain pull_task() was called when the cpu + was idle + 6) # of times in this domain pull_task() was called even though + the target task was cache-hot when idle + 7) # of times in this domain load_balance() was called but did + not find a busier queue while the cpu was idle + 8) # of times in this domain a busier queue was found while the + cpu was idle but no busier group was found + 9) # of times in this domain load_balance() was called when the + cpu was busy + 10) # of times in this domain load_balance() checked but found the + load did not require balancing when busy + 11) # of times in this domain load_balance() tried to move one or + more tasks and failed, when the cpu was busy + 12) sum of imbalances discovered (if any) with each call to + load_balance() in this domain when the cpu was busy + 13) # of times in this domain pull_task() was called when busy + 14) # of times in this domain pull_task() was called even though the + target task was cache-hot when busy + 15) # of times in this domain load_balance() was called but did not + find a busier queue while the cpu was busy + 16) # of times in this domain a busier queue was found while the cpu + was busy but no busier group was found + + 17) # of times in this domain load_balance() was called when the + cpu was just becoming idle + 18) # of times in this domain load_balance() checked but found the + load did not require balancing when the cpu was just becoming idle + 19) # of times in this domain load_balance() tried to move one or more + tasks and failed, when the cpu was just becoming idle + 20) sum of imbalances discovered (if any) with each call to + load_balance() in this domain when the cpu was just becoming idle + 21) # of times in this domain pull_task() was called when newly idle + 22) # of times in this domain pull_task() was called even though the + target task was cache-hot when just becoming idle + 23) # of times in this domain load_balance() was called but did not + find a busier queue while the cpu was just becoming idle + 24) # of times in this domain a busier queue was found while the cpu + was just becoming idle but no busier group was found + + Next three are active_load_balance() statistics: + + 25) # of times active_load_balance() was called + 26) # of times active_load_balance() tried to move a task and failed + 27) # of times active_load_balance() successfully moved a task + + Next three are sched_balance_exec() statistics: + + 28) sbe_cnt is not used + 29) sbe_balanced is not used + 30) sbe_pushed is not used + + Next three are sched_balance_fork() statistics: + + 31) sbf_cnt is not used + 32) sbf_balanced is not used + 33) sbf_pushed is not used + + Next three are try_to_wake_up() statistics: + + 34) # of times in this domain try_to_wake_up() awoke a task that + last ran on a different cpu in this domain + 35) # of times in this domain try_to_wake_up() moved a task to the + waking cpu because it was cache-cold on its own cpu anyway + 36) # of times in this domain try_to_wake_up() started passive balancing + +/proc/<pid>/schedstat +--------------------- +schedstats also adds a new /proc/<pid>/schedstat file to include some of +the same information on a per-process level. There are three fields in +this file correlating for that process to: + + 1) time spent on the cpu + 2) time spent waiting on a runqueue + 3) # of timeslices run on this cpu + +A program could be easily written to make use of these extra fields to +report on how well a particular process or set of processes is faring +under the scheduler's policies. A simple version of such a program is +available at + + http://eaglet.pdxhosts.com/rick/linux/schedstat/v12/latency.c diff --git a/Documentation/scheduler/text_files.rst b/Documentation/scheduler/text_files.rst new file mode 100644 index 000000000..0bc50307b --- /dev/null +++ b/Documentation/scheduler/text_files.rst @@ -0,0 +1,5 @@ +Scheduler pelt c program +------------------------ + +.. literalinclude:: sched-pelt.c + :language: c |