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3108 lines
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3108 lines
114 KiB
Plaintext
============================
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LINUX KERNEL MEMORY BARRIERS
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============================
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By: David Howells <dhowells@redhat.com>
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Paul E. McKenney <paulmck@linux.ibm.com>
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Will Deacon <will.deacon@arm.com>
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Peter Zijlstra <peterz@infradead.org>
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==========
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DISCLAIMER
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==========
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This document is not a specification; it is intentionally (for the sake of
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brevity) and unintentionally (due to being human) incomplete. This document is
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meant as a guide to using the various memory barriers provided by Linux, but
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in case of any doubt (and there are many) please ask. Some doubts may be
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resolved by referring to the formal memory consistency model and related
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documentation at tools/memory-model/. Nevertheless, even this memory
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model should be viewed as the collective opinion of its maintainers rather
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than as an infallible oracle.
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To repeat, this document is not a specification of what Linux expects from
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hardware.
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The purpose of this document is twofold:
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(1) to specify the minimum functionality that one can rely on for any
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particular barrier, and
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(2) to provide a guide as to how to use the barriers that are available.
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Note that an architecture can provide more than the minimum requirement
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for any particular barrier, but if the architecture provides less than
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that, that architecture is incorrect.
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Note also that it is possible that a barrier may be a no-op for an
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architecture because the way that arch works renders an explicit barrier
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unnecessary in that case.
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========
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CONTENTS
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========
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(*) Abstract memory access model.
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- Device operations.
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- Guarantees.
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(*) What are memory barriers?
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- Varieties of memory barrier.
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- What may not be assumed about memory barriers?
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- Data dependency barriers (historical).
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- Control dependencies.
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- SMP barrier pairing.
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- Examples of memory barrier sequences.
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- Read memory barriers vs load speculation.
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- Multicopy atomicity.
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(*) Explicit kernel barriers.
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- Compiler barrier.
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- CPU memory barriers.
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(*) Implicit kernel memory barriers.
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- Lock acquisition functions.
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- Interrupt disabling functions.
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- Sleep and wake-up functions.
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- Miscellaneous functions.
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(*) Inter-CPU acquiring barrier effects.
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- Acquires vs memory accesses.
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(*) Where are memory barriers needed?
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- Interprocessor interaction.
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- Atomic operations.
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- Accessing devices.
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- Interrupts.
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(*) Kernel I/O barrier effects.
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(*) Assumed minimum execution ordering model.
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(*) The effects of the cpu cache.
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- Cache coherency.
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- Cache coherency vs DMA.
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- Cache coherency vs MMIO.
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(*) The things CPUs get up to.
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- And then there's the Alpha.
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- Virtual Machine Guests.
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(*) Example uses.
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- Circular buffers.
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(*) References.
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============================
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ABSTRACT MEMORY ACCESS MODEL
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============================
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Consider the following abstract model of the system:
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: :
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: :
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: :
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+-------+ : +--------+ : +-------+
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| | : | | : | |
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| | : | | : | |
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| CPU 1 |<----->| Memory |<----->| CPU 2 |
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| | : | | : | |
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| | : | | : | |
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+-------+ : +--------+ : +-------+
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^ : ^ : ^
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| : | : |
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| : | : |
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| : v : |
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| : +--------+ : |
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| : | | : |
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| : | | : |
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+---------->| Device |<----------+
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: | | :
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: | | :
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: +--------+ :
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: :
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Each CPU executes a program that generates memory access operations. In the
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abstract CPU, memory operation ordering is very relaxed, and a CPU may actually
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perform the memory operations in any order it likes, provided program causality
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appears to be maintained. Similarly, the compiler may also arrange the
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instructions it emits in any order it likes, provided it doesn't affect the
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apparent operation of the program.
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So in the above diagram, the effects of the memory operations performed by a
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CPU are perceived by the rest of the system as the operations cross the
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interface between the CPU and rest of the system (the dotted lines).
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For example, consider the following sequence of events:
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CPU 1 CPU 2
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=============== ===============
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{ A == 1; B == 2 }
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A = 3; x = B;
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B = 4; y = A;
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The set of accesses as seen by the memory system in the middle can be arranged
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in 24 different combinations:
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STORE A=3, STORE B=4, y=LOAD A->3, x=LOAD B->4
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STORE A=3, STORE B=4, x=LOAD B->4, y=LOAD A->3
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STORE A=3, y=LOAD A->3, STORE B=4, x=LOAD B->4
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STORE A=3, y=LOAD A->3, x=LOAD B->2, STORE B=4
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STORE A=3, x=LOAD B->2, STORE B=4, y=LOAD A->3
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STORE A=3, x=LOAD B->2, y=LOAD A->3, STORE B=4
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STORE B=4, STORE A=3, y=LOAD A->3, x=LOAD B->4
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STORE B=4, ...
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...
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and can thus result in four different combinations of values:
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x == 2, y == 1
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x == 2, y == 3
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x == 4, y == 1
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x == 4, y == 3
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Furthermore, the stores committed by a CPU to the memory system may not be
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perceived by the loads made by another CPU in the same order as the stores were
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committed.
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As a further example, consider this sequence of events:
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CPU 1 CPU 2
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=============== ===============
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{ A == 1, B == 2, C == 3, P == &A, Q == &C }
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B = 4; Q = P;
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P = &B D = *Q;
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There is an obvious data dependency here, as the value loaded into D depends on
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the address retrieved from P by CPU 2. At the end of the sequence, any of the
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following results are possible:
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(Q == &A) and (D == 1)
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(Q == &B) and (D == 2)
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(Q == &B) and (D == 4)
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Note that CPU 2 will never try and load C into D because the CPU will load P
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into Q before issuing the load of *Q.
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DEVICE OPERATIONS
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-----------------
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Some devices present their control interfaces as collections of memory
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locations, but the order in which the control registers are accessed is very
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important. For instance, imagine an ethernet card with a set of internal
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registers that are accessed through an address port register (A) and a data
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port register (D). To read internal register 5, the following code might then
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be used:
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*A = 5;
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x = *D;
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but this might show up as either of the following two sequences:
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STORE *A = 5, x = LOAD *D
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x = LOAD *D, STORE *A = 5
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the second of which will almost certainly result in a malfunction, since it set
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the address _after_ attempting to read the register.
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GUARANTEES
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----------
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There are some minimal guarantees that may be expected of a CPU:
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(*) On any given CPU, dependent memory accesses will be issued in order, with
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respect to itself. This means that for:
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Q = READ_ONCE(P); D = READ_ONCE(*Q);
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the CPU will issue the following memory operations:
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Q = LOAD P, D = LOAD *Q
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and always in that order. However, on DEC Alpha, READ_ONCE() also
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emits a memory-barrier instruction, so that a DEC Alpha CPU will
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instead issue the following memory operations:
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Q = LOAD P, MEMORY_BARRIER, D = LOAD *Q, MEMORY_BARRIER
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Whether on DEC Alpha or not, the READ_ONCE() also prevents compiler
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mischief.
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(*) Overlapping loads and stores within a particular CPU will appear to be
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ordered within that CPU. This means that for:
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a = READ_ONCE(*X); WRITE_ONCE(*X, b);
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the CPU will only issue the following sequence of memory operations:
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a = LOAD *X, STORE *X = b
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And for:
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WRITE_ONCE(*X, c); d = READ_ONCE(*X);
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the CPU will only issue:
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STORE *X = c, d = LOAD *X
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(Loads and stores overlap if they are targeted at overlapping pieces of
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memory).
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And there are a number of things that _must_ or _must_not_ be assumed:
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(*) It _must_not_ be assumed that the compiler will do what you want
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with memory references that are not protected by READ_ONCE() and
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WRITE_ONCE(). Without them, the compiler is within its rights to
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do all sorts of "creative" transformations, which are covered in
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the COMPILER BARRIER section.
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(*) It _must_not_ be assumed that independent loads and stores will be issued
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in the order given. This means that for:
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X = *A; Y = *B; *D = Z;
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we may get any of the following sequences:
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X = LOAD *A, Y = LOAD *B, STORE *D = Z
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X = LOAD *A, STORE *D = Z, Y = LOAD *B
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Y = LOAD *B, X = LOAD *A, STORE *D = Z
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Y = LOAD *B, STORE *D = Z, X = LOAD *A
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STORE *D = Z, X = LOAD *A, Y = LOAD *B
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STORE *D = Z, Y = LOAD *B, X = LOAD *A
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(*) It _must_ be assumed that overlapping memory accesses may be merged or
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discarded. This means that for:
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X = *A; Y = *(A + 4);
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we may get any one of the following sequences:
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X = LOAD *A; Y = LOAD *(A + 4);
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Y = LOAD *(A + 4); X = LOAD *A;
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{X, Y} = LOAD {*A, *(A + 4) };
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And for:
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*A = X; *(A + 4) = Y;
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we may get any of:
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STORE *A = X; STORE *(A + 4) = Y;
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STORE *(A + 4) = Y; STORE *A = X;
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STORE {*A, *(A + 4) } = {X, Y};
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And there are anti-guarantees:
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(*) These guarantees do not apply to bitfields, because compilers often
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generate code to modify these using non-atomic read-modify-write
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sequences. Do not attempt to use bitfields to synchronize parallel
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algorithms.
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(*) Even in cases where bitfields are protected by locks, all fields
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in a given bitfield must be protected by one lock. If two fields
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in a given bitfield are protected by different locks, the compiler's
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non-atomic read-modify-write sequences can cause an update to one
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field to corrupt the value of an adjacent field.
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(*) These guarantees apply only to properly aligned and sized scalar
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variables. "Properly sized" currently means variables that are
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the same size as "char", "short", "int" and "long". "Properly
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aligned" means the natural alignment, thus no constraints for
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"char", two-byte alignment for "short", four-byte alignment for
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"int", and either four-byte or eight-byte alignment for "long",
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on 32-bit and 64-bit systems, respectively. Note that these
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guarantees were introduced into the C11 standard, so beware when
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using older pre-C11 compilers (for example, gcc 4.6). The portion
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of the standard containing this guarantee is Section 3.14, which
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defines "memory location" as follows:
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memory location
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either an object of scalar type, or a maximal sequence
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of adjacent bit-fields all having nonzero width
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NOTE 1: Two threads of execution can update and access
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separate memory locations without interfering with
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each other.
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NOTE 2: A bit-field and an adjacent non-bit-field member
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are in separate memory locations. The same applies
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to two bit-fields, if one is declared inside a nested
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structure declaration and the other is not, or if the two
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are separated by a zero-length bit-field declaration,
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or if they are separated by a non-bit-field member
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declaration. It is not safe to concurrently update two
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bit-fields in the same structure if all members declared
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between them are also bit-fields, no matter what the
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sizes of those intervening bit-fields happen to be.
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=========================
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WHAT ARE MEMORY BARRIERS?
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=========================
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As can be seen above, independent memory operations are effectively performed
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in random order, but this can be a problem for CPU-CPU interaction and for I/O.
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What is required is some way of intervening to instruct the compiler and the
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CPU to restrict the order.
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Memory barriers are such interventions. They impose a perceived partial
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ordering over the memory operations on either side of the barrier.
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Such enforcement is important because the CPUs and other devices in a system
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can use a variety of tricks to improve performance, including reordering,
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deferral and combination of memory operations; speculative loads; speculative
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branch prediction and various types of caching. Memory barriers are used to
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override or suppress these tricks, allowing the code to sanely control the
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interaction of multiple CPUs and/or devices.
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VARIETIES OF MEMORY BARRIER
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---------------------------
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Memory barriers come in four basic varieties:
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(1) Write (or store) memory barriers.
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A write memory barrier gives a guarantee that all the STORE operations
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specified before the barrier will appear to happen before all the STORE
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operations specified after the barrier with respect to the other
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components of the system.
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A write barrier is a partial ordering on stores only; it is not required
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to have any effect on loads.
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A CPU can be viewed as committing a sequence of store operations to the
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memory system as time progresses. All stores _before_ a write barrier
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will occur _before_ all the stores after the write barrier.
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[!] Note that write barriers should normally be paired with read or data
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dependency barriers; see the "SMP barrier pairing" subsection.
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(2) Data dependency barriers.
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A data dependency barrier is a weaker form of read barrier. In the case
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where two loads are performed such that the second depends on the result
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of the first (eg: the first load retrieves the address to which the second
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load will be directed), a data dependency barrier would be required to
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make sure that the target of the second load is updated after the address
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obtained by the first load is accessed.
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A data dependency barrier is a partial ordering on interdependent loads
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only; it is not required to have any effect on stores, independent loads
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or overlapping loads.
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As mentioned in (1), the other CPUs in the system can be viewed as
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committing sequences of stores to the memory system that the CPU being
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considered can then perceive. A data dependency barrier issued by the CPU
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under consideration guarantees that for any load preceding it, if that
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load touches one of a sequence of stores from another CPU, then by the
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time the barrier completes, the effects of all the stores prior to that
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touched by the load will be perceptible to any loads issued after the data
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dependency barrier.
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See the "Examples of memory barrier sequences" subsection for diagrams
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showing the ordering constraints.
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[!] Note that the first load really has to have a _data_ dependency and
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not a control dependency. If the address for the second load is dependent
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on the first load, but the dependency is through a conditional rather than
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actually loading the address itself, then it's a _control_ dependency and
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a full read barrier or better is required. See the "Control dependencies"
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subsection for more information.
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[!] Note that data dependency barriers should normally be paired with
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write barriers; see the "SMP barrier pairing" subsection.
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(3) Read (or load) memory barriers.
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A read barrier is a data dependency barrier plus a guarantee that all the
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LOAD operations specified before the barrier will appear to happen before
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all the LOAD operations specified after the barrier with respect to the
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other components of the system.
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A read barrier is a partial ordering on loads only; it is not required to
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have any effect on stores.
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Read memory barriers imply data dependency barriers, and so can substitute
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for them.
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[!] Note that read barriers should normally be paired with write barriers;
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see the "SMP barrier pairing" subsection.
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(4) General memory barriers.
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A general memory barrier gives a guarantee that all the LOAD and STORE
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operations specified before the barrier will appear to happen before all
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the LOAD and STORE operations specified after the barrier with respect to
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the other components of the system.
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A general memory barrier is a partial ordering over both loads and stores.
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General memory barriers imply both read and write memory barriers, and so
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can substitute for either.
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And a couple of implicit varieties:
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(5) ACQUIRE operations.
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This acts as a one-way permeable barrier. It guarantees that all memory
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operations after the ACQUIRE operation will appear to happen after the
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ACQUIRE operation with respect to the other components of the system.
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ACQUIRE operations include LOCK operations and both smp_load_acquire()
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and smp_cond_load_acquire() operations.
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Memory operations that occur before an ACQUIRE operation may appear to
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happen after it completes.
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An ACQUIRE operation should almost always be paired with a RELEASE
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operation.
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(6) RELEASE operations.
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This also acts as a one-way permeable barrier. It guarantees that all
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memory operations before the RELEASE operation will appear to happen
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before the RELEASE operation with respect to the other components of the
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system. RELEASE operations include UNLOCK operations and
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smp_store_release() operations.
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Memory operations that occur after a RELEASE operation may appear to
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happen before it completes.
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The use of ACQUIRE and RELEASE operations generally precludes the need
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for other sorts of memory barrier. In addition, a RELEASE+ACQUIRE pair is
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-not- guaranteed to act as a full memory barrier. However, after an
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ACQUIRE on a given variable, all memory accesses preceding any prior
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RELEASE on that same variable are guaranteed to be visible. In other
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words, within a given variable's critical section, all accesses of all
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previous critical sections for that variable are guaranteed to have
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completed.
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This means that ACQUIRE acts as a minimal "acquire" operation and
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RELEASE acts as a minimal "release" operation.
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A subset of the atomic operations described in atomic_t.txt have ACQUIRE and
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RELEASE variants in addition to fully-ordered and relaxed (no barrier
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semantics) definitions. For compound atomics performing both a load and a
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store, ACQUIRE semantics apply only to the load and RELEASE semantics apply
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only to the store portion of the operation.
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Memory barriers are only required where there's a possibility of interaction
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between two CPUs or between a CPU and a device. If it can be guaranteed that
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there won't be any such interaction in any particular piece of code, then
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memory barriers are unnecessary in that piece of code.
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|
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Note that these are the _minimum_ guarantees. Different architectures may give
|
|
more substantial guarantees, but they may _not_ be relied upon outside of arch
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specific code.
|
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|
WHAT MAY NOT BE ASSUMED ABOUT MEMORY BARRIERS?
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|
----------------------------------------------
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There are certain things that the Linux kernel memory barriers do not guarantee:
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|
|
(*) There is no guarantee that any of the memory accesses specified before a
|
|
memory barrier will be _complete_ by the completion of a memory barrier
|
|
instruction; the barrier can be considered to draw a line in that CPU's
|
|
access queue that accesses of the appropriate type may not cross.
|
|
|
|
(*) There is no guarantee that issuing a memory barrier on one CPU will have
|
|
any direct effect on another CPU or any other hardware in the system. The
|
|
indirect effect will be the order in which the second CPU sees the effects
|
|
of the first CPU's accesses occur, but see the next point:
|
|
|
|
(*) There is no guarantee that a CPU will see the correct order of effects
|
|
from a second CPU's accesses, even _if_ the second CPU uses a memory
|
|
barrier, unless the first CPU _also_ uses a matching memory barrier (see
|
|
the subsection on "SMP Barrier Pairing").
|
|
|
|
(*) There is no guarantee that some intervening piece of off-the-CPU
|
|
hardware[*] will not reorder the memory accesses. CPU cache coherency
|
|
mechanisms should propagate the indirect effects of a memory barrier
|
|
between CPUs, but might not do so in order.
|
|
|
|
[*] For information on bus mastering DMA and coherency please read:
|
|
|
|
Documentation/driver-api/pci/pci.rst
|
|
Documentation/DMA-API-HOWTO.txt
|
|
Documentation/DMA-API.txt
|
|
|
|
|
|
DATA DEPENDENCY BARRIERS (HISTORICAL)
|
|
-------------------------------------
|
|
|
|
As of v4.15 of the Linux kernel, an smp_read_barrier_depends() was
|
|
added to READ_ONCE(), which means that about the only people who
|
|
need to pay attention to this section are those working on DEC Alpha
|
|
architecture-specific code and those working on READ_ONCE() itself.
|
|
For those who need it, and for those who are interested in the history,
|
|
here is the story of data-dependency barriers.
|
|
|
|
The usage requirements of data dependency barriers are a little subtle, and
|
|
it's not always obvious that they're needed. To illustrate, consider the
|
|
following sequence of events:
|
|
|
|
CPU 1 CPU 2
|
|
=============== ===============
|
|
{ A == 1, B == 2, C == 3, P == &A, Q == &C }
|
|
B = 4;
|
|
<write barrier>
|
|
WRITE_ONCE(P, &B)
|
|
Q = READ_ONCE(P);
|
|
D = *Q;
|
|
|
|
There's a clear data dependency here, and it would seem that by the end of the
|
|
sequence, Q must be either &A or &B, and that:
|
|
|
|
(Q == &A) implies (D == 1)
|
|
(Q == &B) implies (D == 4)
|
|
|
|
But! CPU 2's perception of P may be updated _before_ its perception of B, thus
|
|
leading to the following situation:
|
|
|
|
(Q == &B) and (D == 2) ????
|
|
|
|
While this may seem like a failure of coherency or causality maintenance, it
|
|
isn't, and this behaviour can be observed on certain real CPUs (such as the DEC
|
|
Alpha).
|
|
|
|
To deal with this, a data dependency barrier or better must be inserted
|
|
between the address load and the data load:
|
|
|
|
CPU 1 CPU 2
|
|
=============== ===============
|
|
{ A == 1, B == 2, C == 3, P == &A, Q == &C }
|
|
B = 4;
|
|
<write barrier>
|
|
WRITE_ONCE(P, &B);
|
|
Q = READ_ONCE(P);
|
|
<data dependency barrier>
|
|
D = *Q;
|
|
|
|
This enforces the occurrence of one of the two implications, and prevents the
|
|
third possibility from arising.
|
|
|
|
|
|
[!] Note that this extremely counterintuitive situation arises most easily on
|
|
machines with split caches, so that, for example, one cache bank processes
|
|
even-numbered cache lines and the other bank processes odd-numbered cache
|
|
lines. The pointer P might be stored in an odd-numbered cache line, and the
|
|
variable B might be stored in an even-numbered cache line. Then, if the
|
|
even-numbered bank of the reading CPU's cache is extremely busy while the
|
|
odd-numbered bank is idle, one can see the new value of the pointer P (&B),
|
|
but the old value of the variable B (2).
|
|
|
|
|
|
A data-dependency barrier is not required to order dependent writes
|
|
because the CPUs that the Linux kernel supports don't do writes
|
|
until they are certain (1) that the write will actually happen, (2)
|
|
of the location of the write, and (3) of the value to be written.
|
|
But please carefully read the "CONTROL DEPENDENCIES" section and the
|
|
Documentation/RCU/rcu_dereference.txt file: The compiler can and does
|
|
break dependencies in a great many highly creative ways.
|
|
|
|
CPU 1 CPU 2
|
|
=============== ===============
|
|
{ A == 1, B == 2, C = 3, P == &A, Q == &C }
|
|
B = 4;
|
|
<write barrier>
|
|
WRITE_ONCE(P, &B);
|
|
Q = READ_ONCE(P);
|
|
WRITE_ONCE(*Q, 5);
|
|
|
|
Therefore, no data-dependency barrier is required to order the read into
|
|
Q with the store into *Q. In other words, this outcome is prohibited,
|
|
even without a data-dependency barrier:
|
|
|
|
(Q == &B) && (B == 4)
|
|
|
|
Please note that this pattern should be rare. After all, the whole point
|
|
of dependency ordering is to -prevent- writes to the data structure, along
|
|
with the expensive cache misses associated with those writes. This pattern
|
|
can be used to record rare error conditions and the like, and the CPUs'
|
|
naturally occurring ordering prevents such records from being lost.
|
|
|
|
|
|
Note well that the ordering provided by a data dependency is local to
|
|
the CPU containing it. See the section on "Multicopy atomicity" for
|
|
more information.
|
|
|
|
|
|
The data dependency barrier is very important to the RCU system,
|
|
for example. See rcu_assign_pointer() and rcu_dereference() in
|
|
include/linux/rcupdate.h. This permits the current target of an RCU'd
|
|
pointer to be replaced with a new modified target, without the replacement
|
|
target appearing to be incompletely initialised.
|
|
|
|
See also the subsection on "Cache Coherency" for a more thorough example.
|
|
|
|
|
|
CONTROL DEPENDENCIES
|
|
--------------------
|
|
|
|
Control dependencies can be a bit tricky because current compilers do
|
|
not understand them. The purpose of this section is to help you prevent
|
|
the compiler's ignorance from breaking your code.
|
|
|
|
A load-load control dependency requires a full read memory barrier, not
|
|
simply a data dependency barrier to make it work correctly. Consider the
|
|
following bit of code:
|
|
|
|
q = READ_ONCE(a);
|
|
if (q) {
|
|
<data dependency barrier> /* BUG: No data dependency!!! */
|
|
p = READ_ONCE(b);
|
|
}
|
|
|
|
This will not have the desired effect because there is no actual data
|
|
dependency, but rather a control dependency that the CPU may short-circuit
|
|
by attempting to predict the outcome in advance, so that other CPUs see
|
|
the load from b as having happened before the load from a. In such a
|
|
case what's actually required is:
|
|
|
|
q = READ_ONCE(a);
|
|
if (q) {
|
|
<read barrier>
|
|
p = READ_ONCE(b);
|
|
}
|
|
|
|
However, stores are not speculated. This means that ordering -is- provided
|
|
for load-store control dependencies, as in the following example:
|
|
|
|
q = READ_ONCE(a);
|
|
if (q) {
|
|
WRITE_ONCE(b, 1);
|
|
}
|
|
|
|
Control dependencies pair normally with other types of barriers.
|
|
That said, please note that neither READ_ONCE() nor WRITE_ONCE()
|
|
are optional! Without the READ_ONCE(), the compiler might combine the
|
|
load from 'a' with other loads from 'a'. Without the WRITE_ONCE(),
|
|
the compiler might combine the store to 'b' with other stores to 'b'.
|
|
Either can result in highly counterintuitive effects on ordering.
|
|
|
|
Worse yet, if the compiler is able to prove (say) that the value of
|
|
variable 'a' is always non-zero, it would be well within its rights
|
|
to optimize the original example by eliminating the "if" statement
|
|
as follows:
|
|
|
|
q = a;
|
|
b = 1; /* BUG: Compiler and CPU can both reorder!!! */
|
|
|
|
So don't leave out the READ_ONCE().
|
|
|
|
It is tempting to try to enforce ordering on identical stores on both
|
|
branches of the "if" statement as follows:
|
|
|
|
q = READ_ONCE(a);
|
|
if (q) {
|
|
barrier();
|
|
WRITE_ONCE(b, 1);
|
|
do_something();
|
|
} else {
|
|
barrier();
|
|
WRITE_ONCE(b, 1);
|
|
do_something_else();
|
|
}
|
|
|
|
Unfortunately, current compilers will transform this as follows at high
|
|
optimization levels:
|
|
|
|
q = READ_ONCE(a);
|
|
barrier();
|
|
WRITE_ONCE(b, 1); /* BUG: No ordering vs. load from a!!! */
|
|
if (q) {
|
|
/* WRITE_ONCE(b, 1); -- moved up, BUG!!! */
|
|
do_something();
|
|
} else {
|
|
/* WRITE_ONCE(b, 1); -- moved up, BUG!!! */
|
|
do_something_else();
|
|
}
|
|
|
|
Now there is no conditional between the load from 'a' and the store to
|
|
'b', which means that the CPU is within its rights to reorder them:
|
|
The conditional is absolutely required, and must be present in the
|
|
assembly code even after all compiler optimizations have been applied.
|
|
Therefore, if you need ordering in this example, you need explicit
|
|
memory barriers, for example, smp_store_release():
|
|
|
|
q = READ_ONCE(a);
|
|
if (q) {
|
|
smp_store_release(&b, 1);
|
|
do_something();
|
|
} else {
|
|
smp_store_release(&b, 1);
|
|
do_something_else();
|
|
}
|
|
|
|
In contrast, without explicit memory barriers, two-legged-if control
|
|
ordering is guaranteed only when the stores differ, for example:
|
|
|
|
q = READ_ONCE(a);
|
|
if (q) {
|
|
WRITE_ONCE(b, 1);
|
|
do_something();
|
|
} else {
|
|
WRITE_ONCE(b, 2);
|
|
do_something_else();
|
|
}
|
|
|
|
The initial READ_ONCE() is still required to prevent the compiler from
|
|
proving the value of 'a'.
|
|
|
|
In addition, you need to be careful what you do with the local variable 'q',
|
|
otherwise the compiler might be able to guess the value and again remove
|
|
the needed conditional. For example:
|
|
|
|
q = READ_ONCE(a);
|
|
if (q % MAX) {
|
|
WRITE_ONCE(b, 1);
|
|
do_something();
|
|
} else {
|
|
WRITE_ONCE(b, 2);
|
|
do_something_else();
|
|
}
|
|
|
|
If MAX is defined to be 1, then the compiler knows that (q % MAX) is
|
|
equal to zero, in which case the compiler is within its rights to
|
|
transform the above code into the following:
|
|
|
|
q = READ_ONCE(a);
|
|
WRITE_ONCE(b, 2);
|
|
do_something_else();
|
|
|
|
Given this transformation, the CPU is not required to respect the ordering
|
|
between the load from variable 'a' and the store to variable 'b'. It is
|
|
tempting to add a barrier(), but this does not help. The conditional
|
|
is gone, and the barrier won't bring it back. Therefore, if you are
|
|
relying on this ordering, you should make sure that MAX is greater than
|
|
one, perhaps as follows:
|
|
|
|
q = READ_ONCE(a);
|
|
BUILD_BUG_ON(MAX <= 1); /* Order load from a with store to b. */
|
|
if (q % MAX) {
|
|
WRITE_ONCE(b, 1);
|
|
do_something();
|
|
} else {
|
|
WRITE_ONCE(b, 2);
|
|
do_something_else();
|
|
}
|
|
|
|
Please note once again that the stores to 'b' differ. If they were
|
|
identical, as noted earlier, the compiler could pull this store outside
|
|
of the 'if' statement.
|
|
|
|
You must also be careful not to rely too much on boolean short-circuit
|
|
evaluation. Consider this example:
|
|
|
|
q = READ_ONCE(a);
|
|
if (q || 1 > 0)
|
|
WRITE_ONCE(b, 1);
|
|
|
|
Because the first condition cannot fault and the second condition is
|
|
always true, the compiler can transform this example as following,
|
|
defeating control dependency:
|
|
|
|
q = READ_ONCE(a);
|
|
WRITE_ONCE(b, 1);
|
|
|
|
This example underscores the need to ensure that the compiler cannot
|
|
out-guess your code. More generally, although READ_ONCE() does force
|
|
the compiler to actually emit code for a given load, it does not force
|
|
the compiler to use the results.
|
|
|
|
In addition, control dependencies apply only to the then-clause and
|
|
else-clause of the if-statement in question. In particular, it does
|
|
not necessarily apply to code following the if-statement:
|
|
|
|
q = READ_ONCE(a);
|
|
if (q) {
|
|
WRITE_ONCE(b, 1);
|
|
} else {
|
|
WRITE_ONCE(b, 2);
|
|
}
|
|
WRITE_ONCE(c, 1); /* BUG: No ordering against the read from 'a'. */
|
|
|
|
It is tempting to argue that there in fact is ordering because the
|
|
compiler cannot reorder volatile accesses and also cannot reorder
|
|
the writes to 'b' with the condition. Unfortunately for this line
|
|
of reasoning, the compiler might compile the two writes to 'b' as
|
|
conditional-move instructions, as in this fanciful pseudo-assembly
|
|
language:
|
|
|
|
ld r1,a
|
|
cmp r1,$0
|
|
cmov,ne r4,$1
|
|
cmov,eq r4,$2
|
|
st r4,b
|
|
st $1,c
|
|
|
|
A weakly ordered CPU would have no dependency of any sort between the load
|
|
from 'a' and the store to 'c'. The control dependencies would extend
|
|
only to the pair of cmov instructions and the store depending on them.
|
|
In short, control dependencies apply only to the stores in the then-clause
|
|
and else-clause of the if-statement in question (including functions
|
|
invoked by those two clauses), not to code following that if-statement.
|
|
|
|
|
|
Note well that the ordering provided by a control dependency is local
|
|
to the CPU containing it. See the section on "Multicopy atomicity"
|
|
for more information.
|
|
|
|
|
|
In summary:
|
|
|
|
(*) Control dependencies can order prior loads against later stores.
|
|
However, they do -not- guarantee any other sort of ordering:
|
|
Not prior loads against later loads, nor prior stores against
|
|
later anything. If you need these other forms of ordering,
|
|
use smp_rmb(), smp_wmb(), or, in the case of prior stores and
|
|
later loads, smp_mb().
|
|
|
|
(*) If both legs of the "if" statement begin with identical stores to
|
|
the same variable, then those stores must be ordered, either by
|
|
preceding both of them with smp_mb() or by using smp_store_release()
|
|
to carry out the stores. Please note that it is -not- sufficient
|
|
to use barrier() at beginning of each leg of the "if" statement
|
|
because, as shown by the example above, optimizing compilers can
|
|
destroy the control dependency while respecting the letter of the
|
|
barrier() law.
|
|
|
|
(*) Control dependencies require at least one run-time conditional
|
|
between the prior load and the subsequent store, and this
|
|
conditional must involve the prior load. If the compiler is able
|
|
to optimize the conditional away, it will have also optimized
|
|
away the ordering. Careful use of READ_ONCE() and WRITE_ONCE()
|
|
can help to preserve the needed conditional.
|
|
|
|
(*) Control dependencies require that the compiler avoid reordering the
|
|
dependency into nonexistence. Careful use of READ_ONCE() or
|
|
atomic{,64}_read() can help to preserve your control dependency.
|
|
Please see the COMPILER BARRIER section for more information.
|
|
|
|
(*) Control dependencies apply only to the then-clause and else-clause
|
|
of the if-statement containing the control dependency, including
|
|
any functions that these two clauses call. Control dependencies
|
|
do -not- apply to code following the if-statement containing the
|
|
control dependency.
|
|
|
|
(*) Control dependencies pair normally with other types of barriers.
|
|
|
|
(*) Control dependencies do -not- provide multicopy atomicity. If you
|
|
need all the CPUs to see a given store at the same time, use smp_mb().
|
|
|
|
(*) Compilers do not understand control dependencies. It is therefore
|
|
your job to ensure that they do not break your code.
|
|
|
|
|
|
SMP BARRIER PAIRING
|
|
-------------------
|
|
|
|
When dealing with CPU-CPU interactions, certain types of memory barrier should
|
|
always be paired. A lack of appropriate pairing is almost certainly an error.
|
|
|
|
General barriers pair with each other, though they also pair with most
|
|
other types of barriers, albeit without multicopy atomicity. An acquire
|
|
barrier pairs with a release barrier, but both may also pair with other
|
|
barriers, including of course general barriers. A write barrier pairs
|
|
with a data dependency barrier, a control dependency, an acquire barrier,
|
|
a release barrier, a read barrier, or a general barrier. Similarly a
|
|
read barrier, control dependency, or a data dependency barrier pairs
|
|
with a write barrier, an acquire barrier, a release barrier, or a
|
|
general barrier:
|
|
|
|
CPU 1 CPU 2
|
|
=============== ===============
|
|
WRITE_ONCE(a, 1);
|
|
<write barrier>
|
|
WRITE_ONCE(b, 2); x = READ_ONCE(b);
|
|
<read barrier>
|
|
y = READ_ONCE(a);
|
|
|
|
Or:
|
|
|
|
CPU 1 CPU 2
|
|
=============== ===============================
|
|
a = 1;
|
|
<write barrier>
|
|
WRITE_ONCE(b, &a); x = READ_ONCE(b);
|
|
<data dependency barrier>
|
|
y = *x;
|
|
|
|
Or even:
|
|
|
|
CPU 1 CPU 2
|
|
=============== ===============================
|
|
r1 = READ_ONCE(y);
|
|
<general barrier>
|
|
WRITE_ONCE(x, 1); if (r2 = READ_ONCE(x)) {
|
|
<implicit control dependency>
|
|
WRITE_ONCE(y, 1);
|
|
}
|
|
|
|
assert(r1 == 0 || r2 == 0);
|
|
|
|
Basically, the read barrier always has to be there, even though it can be of
|
|
the "weaker" type.
|
|
|
|
[!] Note that the stores before the write barrier would normally be expected to
|
|
match the loads after the read barrier or the data dependency barrier, and vice
|
|
versa:
|
|
|
|
CPU 1 CPU 2
|
|
=================== ===================
|
|
WRITE_ONCE(a, 1); }---- --->{ v = READ_ONCE(c);
|
|
WRITE_ONCE(b, 2); } \ / { w = READ_ONCE(d);
|
|
<write barrier> \ <read barrier>
|
|
WRITE_ONCE(c, 3); } / \ { x = READ_ONCE(a);
|
|
WRITE_ONCE(d, 4); }---- --->{ y = READ_ONCE(b);
|
|
|
|
|
|
EXAMPLES OF MEMORY BARRIER SEQUENCES
|
|
------------------------------------
|
|
|
|
Firstly, write barriers act as partial orderings on store operations.
|
|
Consider the following sequence of events:
|
|
|
|
CPU 1
|
|
=======================
|
|
STORE A = 1
|
|
STORE B = 2
|
|
STORE C = 3
|
|
<write barrier>
|
|
STORE D = 4
|
|
STORE E = 5
|
|
|
|
This sequence of events is committed to the memory coherence system in an order
|
|
that the rest of the system might perceive as the unordered set of { STORE A,
|
|
STORE B, STORE C } all occurring before the unordered set of { STORE D, STORE E
|
|
}:
|
|
|
|
+-------+ : :
|
|
| | +------+
|
|
| |------>| C=3 | } /\
|
|
| | : +------+ }----- \ -----> Events perceptible to
|
|
| | : | A=1 | } \/ the rest of the system
|
|
| | : +------+ }
|
|
| CPU 1 | : | B=2 | }
|
|
| | +------+ }
|
|
| | wwwwwwwwwwwwwwww } <--- At this point the write barrier
|
|
| | +------+ } requires all stores prior to the
|
|
| | : | E=5 | } barrier to be committed before
|
|
| | : +------+ } further stores may take place
|
|
| |------>| D=4 | }
|
|
| | +------+
|
|
+-------+ : :
|
|
|
|
|
| Sequence in which stores are committed to the
|
|
| memory system by CPU 1
|
|
V
|
|
|
|
|
|
Secondly, data dependency barriers act as partial orderings on data-dependent
|
|
loads. Consider the following sequence of events:
|
|
|
|
CPU 1 CPU 2
|
|
======================= =======================
|
|
{ B = 7; X = 9; Y = 8; C = &Y }
|
|
STORE A = 1
|
|
STORE B = 2
|
|
<write barrier>
|
|
STORE C = &B LOAD X
|
|
STORE D = 4 LOAD C (gets &B)
|
|
LOAD *C (reads B)
|
|
|
|
Without intervention, CPU 2 may perceive the events on CPU 1 in some
|
|
effectively random order, despite the write barrier issued by CPU 1:
|
|
|
|
+-------+ : : : :
|
|
| | +------+ +-------+ | Sequence of update
|
|
| |------>| B=2 |----- --->| Y->8 | | of perception on
|
|
| | : +------+ \ +-------+ | CPU 2
|
|
| CPU 1 | : | A=1 | \ --->| C->&Y | V
|
|
| | +------+ | +-------+
|
|
| | wwwwwwwwwwwwwwww | : :
|
|
| | +------+ | : :
|
|
| | : | C=&B |--- | : : +-------+
|
|
| | : +------+ \ | +-------+ | |
|
|
| |------>| D=4 | ----------->| C->&B |------>| |
|
|
| | +------+ | +-------+ | |
|
|
+-------+ : : | : : | |
|
|
| : : | |
|
|
| : : | CPU 2 |
|
|
| +-------+ | |
|
|
Apparently incorrect ---> | | B->7 |------>| |
|
|
perception of B (!) | +-------+ | |
|
|
| : : | |
|
|
| +-------+ | |
|
|
The load of X holds ---> \ | X->9 |------>| |
|
|
up the maintenance \ +-------+ | |
|
|
of coherence of B ----->| B->2 | +-------+
|
|
+-------+
|
|
: :
|
|
|
|
|
|
In the above example, CPU 2 perceives that B is 7, despite the load of *C
|
|
(which would be B) coming after the LOAD of C.
|
|
|
|
If, however, a data dependency barrier were to be placed between the load of C
|
|
and the load of *C (ie: B) on CPU 2:
|
|
|
|
CPU 1 CPU 2
|
|
======================= =======================
|
|
{ B = 7; X = 9; Y = 8; C = &Y }
|
|
STORE A = 1
|
|
STORE B = 2
|
|
<write barrier>
|
|
STORE C = &B LOAD X
|
|
STORE D = 4 LOAD C (gets &B)
|
|
<data dependency barrier>
|
|
LOAD *C (reads B)
|
|
|
|
then the following will occur:
|
|
|
|
+-------+ : : : :
|
|
| | +------+ +-------+
|
|
| |------>| B=2 |----- --->| Y->8 |
|
|
| | : +------+ \ +-------+
|
|
| CPU 1 | : | A=1 | \ --->| C->&Y |
|
|
| | +------+ | +-------+
|
|
| | wwwwwwwwwwwwwwww | : :
|
|
| | +------+ | : :
|
|
| | : | C=&B |--- | : : +-------+
|
|
| | : +------+ \ | +-------+ | |
|
|
| |------>| D=4 | ----------->| C->&B |------>| |
|
|
| | +------+ | +-------+ | |
|
|
+-------+ : : | : : | |
|
|
| : : | |
|
|
| : : | CPU 2 |
|
|
| +-------+ | |
|
|
| | X->9 |------>| |
|
|
| +-------+ | |
|
|
Makes sure all effects ---> \ ddddddddddddddddd | |
|
|
prior to the store of C \ +-------+ | |
|
|
are perceptible to ----->| B->2 |------>| |
|
|
subsequent loads +-------+ | |
|
|
: : +-------+
|
|
|
|
|
|
And thirdly, a read barrier acts as a partial order on loads. Consider the
|
|
following sequence of events:
|
|
|
|
CPU 1 CPU 2
|
|
======================= =======================
|
|
{ A = 0, B = 9 }
|
|
STORE A=1
|
|
<write barrier>
|
|
STORE B=2
|
|
LOAD B
|
|
LOAD A
|
|
|
|
Without intervention, CPU 2 may then choose to perceive the events on CPU 1 in
|
|
some effectively random order, despite the write barrier issued by CPU 1:
|
|
|
|
+-------+ : : : :
|
|
| | +------+ +-------+
|
|
| |------>| A=1 |------ --->| A->0 |
|
|
| | +------+ \ +-------+
|
|
| CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
|
|
| | +------+ | +-------+
|
|
| |------>| B=2 |--- | : :
|
|
| | +------+ \ | : : +-------+
|
|
+-------+ : : \ | +-------+ | |
|
|
---------->| B->2 |------>| |
|
|
| +-------+ | CPU 2 |
|
|
| | A->0 |------>| |
|
|
| +-------+ | |
|
|
| : : +-------+
|
|
\ : :
|
|
\ +-------+
|
|
---->| A->1 |
|
|
+-------+
|
|
: :
|
|
|
|
|
|
If, however, a read barrier were to be placed between the load of B and the
|
|
load of A on CPU 2:
|
|
|
|
CPU 1 CPU 2
|
|
======================= =======================
|
|
{ A = 0, B = 9 }
|
|
STORE A=1
|
|
<write barrier>
|
|
STORE B=2
|
|
LOAD B
|
|
<read barrier>
|
|
LOAD A
|
|
|
|
then the partial ordering imposed by CPU 1 will be perceived correctly by CPU
|
|
2:
|
|
|
|
+-------+ : : : :
|
|
| | +------+ +-------+
|
|
| |------>| A=1 |------ --->| A->0 |
|
|
| | +------+ \ +-------+
|
|
| CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
|
|
| | +------+ | +-------+
|
|
| |------>| B=2 |--- | : :
|
|
| | +------+ \ | : : +-------+
|
|
+-------+ : : \ | +-------+ | |
|
|
---------->| B->2 |------>| |
|
|
| +-------+ | CPU 2 |
|
|
| : : | |
|
|
| : : | |
|
|
At this point the read ----> \ rrrrrrrrrrrrrrrrr | |
|
|
barrier causes all effects \ +-------+ | |
|
|
prior to the storage of B ---->| A->1 |------>| |
|
|
to be perceptible to CPU 2 +-------+ | |
|
|
: : +-------+
|
|
|
|
|
|
To illustrate this more completely, consider what could happen if the code
|
|
contained a load of A either side of the read barrier:
|
|
|
|
CPU 1 CPU 2
|
|
======================= =======================
|
|
{ A = 0, B = 9 }
|
|
STORE A=1
|
|
<write barrier>
|
|
STORE B=2
|
|
LOAD B
|
|
LOAD A [first load of A]
|
|
<read barrier>
|
|
LOAD A [second load of A]
|
|
|
|
Even though the two loads of A both occur after the load of B, they may both
|
|
come up with different values:
|
|
|
|
+-------+ : : : :
|
|
| | +------+ +-------+
|
|
| |------>| A=1 |------ --->| A->0 |
|
|
| | +------+ \ +-------+
|
|
| CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
|
|
| | +------+ | +-------+
|
|
| |------>| B=2 |--- | : :
|
|
| | +------+ \ | : : +-------+
|
|
+-------+ : : \ | +-------+ | |
|
|
---------->| B->2 |------>| |
|
|
| +-------+ | CPU 2 |
|
|
| : : | |
|
|
| : : | |
|
|
| +-------+ | |
|
|
| | A->0 |------>| 1st |
|
|
| +-------+ | |
|
|
At this point the read ----> \ rrrrrrrrrrrrrrrrr | |
|
|
barrier causes all effects \ +-------+ | |
|
|
prior to the storage of B ---->| A->1 |------>| 2nd |
|
|
to be perceptible to CPU 2 +-------+ | |
|
|
: : +-------+
|
|
|
|
|
|
But it may be that the update to A from CPU 1 becomes perceptible to CPU 2
|
|
before the read barrier completes anyway:
|
|
|
|
+-------+ : : : :
|
|
| | +------+ +-------+
|
|
| |------>| A=1 |------ --->| A->0 |
|
|
| | +------+ \ +-------+
|
|
| CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
|
|
| | +------+ | +-------+
|
|
| |------>| B=2 |--- | : :
|
|
| | +------+ \ | : : +-------+
|
|
+-------+ : : \ | +-------+ | |
|
|
---------->| B->2 |------>| |
|
|
| +-------+ | CPU 2 |
|
|
| : : | |
|
|
\ : : | |
|
|
\ +-------+ | |
|
|
---->| A->1 |------>| 1st |
|
|
+-------+ | |
|
|
rrrrrrrrrrrrrrrrr | |
|
|
+-------+ | |
|
|
| A->1 |------>| 2nd |
|
|
+-------+ | |
|
|
: : +-------+
|
|
|
|
|
|
The guarantee is that the second load will always come up with A == 1 if the
|
|
load of B came up with B == 2. No such guarantee exists for the first load of
|
|
A; that may come up with either A == 0 or A == 1.
|
|
|
|
|
|
READ MEMORY BARRIERS VS LOAD SPECULATION
|
|
----------------------------------------
|
|
|
|
Many CPUs speculate with loads: that is they see that they will need to load an
|
|
item from memory, and they find a time where they're not using the bus for any
|
|
other loads, and so do the load in advance - even though they haven't actually
|
|
got to that point in the instruction execution flow yet. This permits the
|
|
actual load instruction to potentially complete immediately because the CPU
|
|
already has the value to hand.
|
|
|
|
It may turn out that the CPU didn't actually need the value - perhaps because a
|
|
branch circumvented the load - in which case it can discard the value or just
|
|
cache it for later use.
|
|
|
|
Consider:
|
|
|
|
CPU 1 CPU 2
|
|
======================= =======================
|
|
LOAD B
|
|
DIVIDE } Divide instructions generally
|
|
DIVIDE } take a long time to perform
|
|
LOAD A
|
|
|
|
Which might appear as this:
|
|
|
|
: : +-------+
|
|
+-------+ | |
|
|
--->| B->2 |------>| |
|
|
+-------+ | CPU 2 |
|
|
: :DIVIDE | |
|
|
+-------+ | |
|
|
The CPU being busy doing a ---> --->| A->0 |~~~~ | |
|
|
division speculates on the +-------+ ~ | |
|
|
LOAD of A : : ~ | |
|
|
: :DIVIDE | |
|
|
: : ~ | |
|
|
Once the divisions are complete --> : : ~-->| |
|
|
the CPU can then perform the : : | |
|
|
LOAD with immediate effect : : +-------+
|
|
|
|
|
|
Placing a read barrier or a data dependency barrier just before the second
|
|
load:
|
|
|
|
CPU 1 CPU 2
|
|
======================= =======================
|
|
LOAD B
|
|
DIVIDE
|
|
DIVIDE
|
|
<read barrier>
|
|
LOAD A
|
|
|
|
will force any value speculatively obtained to be reconsidered to an extent
|
|
dependent on the type of barrier used. If there was no change made to the
|
|
speculated memory location, then the speculated value will just be used:
|
|
|
|
: : +-------+
|
|
+-------+ | |
|
|
--->| B->2 |------>| |
|
|
+-------+ | CPU 2 |
|
|
: :DIVIDE | |
|
|
+-------+ | |
|
|
The CPU being busy doing a ---> --->| A->0 |~~~~ | |
|
|
division speculates on the +-------+ ~ | |
|
|
LOAD of A : : ~ | |
|
|
: :DIVIDE | |
|
|
: : ~ | |
|
|
: : ~ | |
|
|
rrrrrrrrrrrrrrrr~ | |
|
|
: : ~ | |
|
|
: : ~-->| |
|
|
: : | |
|
|
: : +-------+
|
|
|
|
|
|
but if there was an update or an invalidation from another CPU pending, then
|
|
the speculation will be cancelled and the value reloaded:
|
|
|
|
: : +-------+
|
|
+-------+ | |
|
|
--->| B->2 |------>| |
|
|
+-------+ | CPU 2 |
|
|
: :DIVIDE | |
|
|
+-------+ | |
|
|
The CPU being busy doing a ---> --->| A->0 |~~~~ | |
|
|
division speculates on the +-------+ ~ | |
|
|
LOAD of A : : ~ | |
|
|
: :DIVIDE | |
|
|
: : ~ | |
|
|
: : ~ | |
|
|
rrrrrrrrrrrrrrrrr | |
|
|
+-------+ | |
|
|
The speculation is discarded ---> --->| A->1 |------>| |
|
|
and an updated value is +-------+ | |
|
|
retrieved : : +-------+
|
|
|
|
|
|
MULTICOPY ATOMICITY
|
|
--------------------
|
|
|
|
Multicopy atomicity is a deeply intuitive notion about ordering that is
|
|
not always provided by real computer systems, namely that a given store
|
|
becomes visible at the same time to all CPUs, or, alternatively, that all
|
|
CPUs agree on the order in which all stores become visible. However,
|
|
support of full multicopy atomicity would rule out valuable hardware
|
|
optimizations, so a weaker form called ``other multicopy atomicity''
|
|
instead guarantees only that a given store becomes visible at the same
|
|
time to all -other- CPUs. The remainder of this document discusses this
|
|
weaker form, but for brevity will call it simply ``multicopy atomicity''.
|
|
|
|
The following example demonstrates multicopy atomicity:
|
|
|
|
CPU 1 CPU 2 CPU 3
|
|
======================= ======================= =======================
|
|
{ X = 0, Y = 0 }
|
|
STORE X=1 r1=LOAD X (reads 1) LOAD Y (reads 1)
|
|
<general barrier> <read barrier>
|
|
STORE Y=r1 LOAD X
|
|
|
|
Suppose that CPU 2's load from X returns 1, which it then stores to Y,
|
|
and CPU 3's load from Y returns 1. This indicates that CPU 1's store
|
|
to X precedes CPU 2's load from X and that CPU 2's store to Y precedes
|
|
CPU 3's load from Y. In addition, the memory barriers guarantee that
|
|
CPU 2 executes its load before its store, and CPU 3 loads from Y before
|
|
it loads from X. The question is then "Can CPU 3's load from X return 0?"
|
|
|
|
Because CPU 3's load from X in some sense comes after CPU 2's load, it
|
|
is natural to expect that CPU 3's load from X must therefore return 1.
|
|
This expectation follows from multicopy atomicity: if a load executing
|
|
on CPU B follows a load from the same variable executing on CPU A (and
|
|
CPU A did not originally store the value which it read), then on
|
|
multicopy-atomic systems, CPU B's load must return either the same value
|
|
that CPU A's load did or some later value. However, the Linux kernel
|
|
does not require systems to be multicopy atomic.
|
|
|
|
The use of a general memory barrier in the example above compensates
|
|
for any lack of multicopy atomicity. In the example, if CPU 2's load
|
|
from X returns 1 and CPU 3's load from Y returns 1, then CPU 3's load
|
|
from X must indeed also return 1.
|
|
|
|
However, dependencies, read barriers, and write barriers are not always
|
|
able to compensate for non-multicopy atomicity. For example, suppose
|
|
that CPU 2's general barrier is removed from the above example, leaving
|
|
only the data dependency shown below:
|
|
|
|
CPU 1 CPU 2 CPU 3
|
|
======================= ======================= =======================
|
|
{ X = 0, Y = 0 }
|
|
STORE X=1 r1=LOAD X (reads 1) LOAD Y (reads 1)
|
|
<data dependency> <read barrier>
|
|
STORE Y=r1 LOAD X (reads 0)
|
|
|
|
This substitution allows non-multicopy atomicity to run rampant: in
|
|
this example, it is perfectly legal for CPU 2's load from X to return 1,
|
|
CPU 3's load from Y to return 1, and its load from X to return 0.
|
|
|
|
The key point is that although CPU 2's data dependency orders its load
|
|
and store, it does not guarantee to order CPU 1's store. Thus, if this
|
|
example runs on a non-multicopy-atomic system where CPUs 1 and 2 share a
|
|
store buffer or a level of cache, CPU 2 might have early access to CPU 1's
|
|
writes. General barriers are therefore required to ensure that all CPUs
|
|
agree on the combined order of multiple accesses.
|
|
|
|
General barriers can compensate not only for non-multicopy atomicity,
|
|
but can also generate additional ordering that can ensure that -all-
|
|
CPUs will perceive the same order of -all- operations. In contrast, a
|
|
chain of release-acquire pairs do not provide this additional ordering,
|
|
which means that only those CPUs on the chain are guaranteed to agree
|
|
on the combined order of the accesses. For example, switching to C code
|
|
in deference to the ghost of Herman Hollerith:
|
|
|
|
int u, v, x, y, z;
|
|
|
|
void cpu0(void)
|
|
{
|
|
r0 = smp_load_acquire(&x);
|
|
WRITE_ONCE(u, 1);
|
|
smp_store_release(&y, 1);
|
|
}
|
|
|
|
void cpu1(void)
|
|
{
|
|
r1 = smp_load_acquire(&y);
|
|
r4 = READ_ONCE(v);
|
|
r5 = READ_ONCE(u);
|
|
smp_store_release(&z, 1);
|
|
}
|
|
|
|
void cpu2(void)
|
|
{
|
|
r2 = smp_load_acquire(&z);
|
|
smp_store_release(&x, 1);
|
|
}
|
|
|
|
void cpu3(void)
|
|
{
|
|
WRITE_ONCE(v, 1);
|
|
smp_mb();
|
|
r3 = READ_ONCE(u);
|
|
}
|
|
|
|
Because cpu0(), cpu1(), and cpu2() participate in a chain of
|
|
smp_store_release()/smp_load_acquire() pairs, the following outcome
|
|
is prohibited:
|
|
|
|
r0 == 1 && r1 == 1 && r2 == 1
|
|
|
|
Furthermore, because of the release-acquire relationship between cpu0()
|
|
and cpu1(), cpu1() must see cpu0()'s writes, so that the following
|
|
outcome is prohibited:
|
|
|
|
r1 == 1 && r5 == 0
|
|
|
|
However, the ordering provided by a release-acquire chain is local
|
|
to the CPUs participating in that chain and does not apply to cpu3(),
|
|
at least aside from stores. Therefore, the following outcome is possible:
|
|
|
|
r0 == 0 && r1 == 1 && r2 == 1 && r3 == 0 && r4 == 0
|
|
|
|
As an aside, the following outcome is also possible:
|
|
|
|
r0 == 0 && r1 == 1 && r2 == 1 && r3 == 0 && r4 == 0 && r5 == 1
|
|
|
|
Although cpu0(), cpu1(), and cpu2() will see their respective reads and
|
|
writes in order, CPUs not involved in the release-acquire chain might
|
|
well disagree on the order. This disagreement stems from the fact that
|
|
the weak memory-barrier instructions used to implement smp_load_acquire()
|
|
and smp_store_release() are not required to order prior stores against
|
|
subsequent loads in all cases. This means that cpu3() can see cpu0()'s
|
|
store to u as happening -after- cpu1()'s load from v, even though
|
|
both cpu0() and cpu1() agree that these two operations occurred in the
|
|
intended order.
|
|
|
|
However, please keep in mind that smp_load_acquire() is not magic.
|
|
In particular, it simply reads from its argument with ordering. It does
|
|
-not- ensure that any particular value will be read. Therefore, the
|
|
following outcome is possible:
|
|
|
|
r0 == 0 && r1 == 0 && r2 == 0 && r5 == 0
|
|
|
|
Note that this outcome can happen even on a mythical sequentially
|
|
consistent system where nothing is ever reordered.
|
|
|
|
To reiterate, if your code requires full ordering of all operations,
|
|
use general barriers throughout.
|
|
|
|
|
|
========================
|
|
EXPLICIT KERNEL BARRIERS
|
|
========================
|
|
|
|
The Linux kernel has a variety of different barriers that act at different
|
|
levels:
|
|
|
|
(*) Compiler barrier.
|
|
|
|
(*) CPU memory barriers.
|
|
|
|
|
|
COMPILER BARRIER
|
|
----------------
|
|
|
|
The Linux kernel has an explicit compiler barrier function that prevents the
|
|
compiler from moving the memory accesses either side of it to the other side:
|
|
|
|
barrier();
|
|
|
|
This is a general barrier -- there are no read-read or write-write
|
|
variants of barrier(). However, READ_ONCE() and WRITE_ONCE() can be
|
|
thought of as weak forms of barrier() that affect only the specific
|
|
accesses flagged by the READ_ONCE() or WRITE_ONCE().
|
|
|
|
The barrier() function has the following effects:
|
|
|
|
(*) Prevents the compiler from reordering accesses following the
|
|
barrier() to precede any accesses preceding the barrier().
|
|
One example use for this property is to ease communication between
|
|
interrupt-handler code and the code that was interrupted.
|
|
|
|
(*) Within a loop, forces the compiler to load the variables used
|
|
in that loop's conditional on each pass through that loop.
|
|
|
|
The READ_ONCE() and WRITE_ONCE() functions can prevent any number of
|
|
optimizations that, while perfectly safe in single-threaded code, can
|
|
be fatal in concurrent code. Here are some examples of these sorts
|
|
of optimizations:
|
|
|
|
(*) The compiler is within its rights to reorder loads and stores
|
|
to the same variable, and in some cases, the CPU is within its
|
|
rights to reorder loads to the same variable. This means that
|
|
the following code:
|
|
|
|
a[0] = x;
|
|
a[1] = x;
|
|
|
|
Might result in an older value of x stored in a[1] than in a[0].
|
|
Prevent both the compiler and the CPU from doing this as follows:
|
|
|
|
a[0] = READ_ONCE(x);
|
|
a[1] = READ_ONCE(x);
|
|
|
|
In short, READ_ONCE() and WRITE_ONCE() provide cache coherence for
|
|
accesses from multiple CPUs to a single variable.
|
|
|
|
(*) The compiler is within its rights to merge successive loads from
|
|
the same variable. Such merging can cause the compiler to "optimize"
|
|
the following code:
|
|
|
|
while (tmp = a)
|
|
do_something_with(tmp);
|
|
|
|
into the following code, which, although in some sense legitimate
|
|
for single-threaded code, is almost certainly not what the developer
|
|
intended:
|
|
|
|
if (tmp = a)
|
|
for (;;)
|
|
do_something_with(tmp);
|
|
|
|
Use READ_ONCE() to prevent the compiler from doing this to you:
|
|
|
|
while (tmp = READ_ONCE(a))
|
|
do_something_with(tmp);
|
|
|
|
(*) The compiler is within its rights to reload a variable, for example,
|
|
in cases where high register pressure prevents the compiler from
|
|
keeping all data of interest in registers. The compiler might
|
|
therefore optimize the variable 'tmp' out of our previous example:
|
|
|
|
while (tmp = a)
|
|
do_something_with(tmp);
|
|
|
|
This could result in the following code, which is perfectly safe in
|
|
single-threaded code, but can be fatal in concurrent code:
|
|
|
|
while (a)
|
|
do_something_with(a);
|
|
|
|
For example, the optimized version of this code could result in
|
|
passing a zero to do_something_with() in the case where the variable
|
|
a was modified by some other CPU between the "while" statement and
|
|
the call to do_something_with().
|
|
|
|
Again, use READ_ONCE() to prevent the compiler from doing this:
|
|
|
|
while (tmp = READ_ONCE(a))
|
|
do_something_with(tmp);
|
|
|
|
Note that if the compiler runs short of registers, it might save
|
|
tmp onto the stack. The overhead of this saving and later restoring
|
|
is why compilers reload variables. Doing so is perfectly safe for
|
|
single-threaded code, so you need to tell the compiler about cases
|
|
where it is not safe.
|
|
|
|
(*) The compiler is within its rights to omit a load entirely if it knows
|
|
what the value will be. For example, if the compiler can prove that
|
|
the value of variable 'a' is always zero, it can optimize this code:
|
|
|
|
while (tmp = a)
|
|
do_something_with(tmp);
|
|
|
|
Into this:
|
|
|
|
do { } while (0);
|
|
|
|
This transformation is a win for single-threaded code because it
|
|
gets rid of a load and a branch. The problem is that the compiler
|
|
will carry out its proof assuming that the current CPU is the only
|
|
one updating variable 'a'. If variable 'a' is shared, then the
|
|
compiler's proof will be erroneous. Use READ_ONCE() to tell the
|
|
compiler that it doesn't know as much as it thinks it does:
|
|
|
|
while (tmp = READ_ONCE(a))
|
|
do_something_with(tmp);
|
|
|
|
But please note that the compiler is also closely watching what you
|
|
do with the value after the READ_ONCE(). For example, suppose you
|
|
do the following and MAX is a preprocessor macro with the value 1:
|
|
|
|
while ((tmp = READ_ONCE(a)) % MAX)
|
|
do_something_with(tmp);
|
|
|
|
Then the compiler knows that the result of the "%" operator applied
|
|
to MAX will always be zero, again allowing the compiler to optimize
|
|
the code into near-nonexistence. (It will still load from the
|
|
variable 'a'.)
|
|
|
|
(*) Similarly, the compiler is within its rights to omit a store entirely
|
|
if it knows that the variable already has the value being stored.
|
|
Again, the compiler assumes that the current CPU is the only one
|
|
storing into the variable, which can cause the compiler to do the
|
|
wrong thing for shared variables. For example, suppose you have
|
|
the following:
|
|
|
|
a = 0;
|
|
... Code that does not store to variable a ...
|
|
a = 0;
|
|
|
|
The compiler sees that the value of variable 'a' is already zero, so
|
|
it might well omit the second store. This would come as a fatal
|
|
surprise if some other CPU might have stored to variable 'a' in the
|
|
meantime.
|
|
|
|
Use WRITE_ONCE() to prevent the compiler from making this sort of
|
|
wrong guess:
|
|
|
|
WRITE_ONCE(a, 0);
|
|
... Code that does not store to variable a ...
|
|
WRITE_ONCE(a, 0);
|
|
|
|
(*) The compiler is within its rights to reorder memory accesses unless
|
|
you tell it not to. For example, consider the following interaction
|
|
between process-level code and an interrupt handler:
|
|
|
|
void process_level(void)
|
|
{
|
|
msg = get_message();
|
|
flag = true;
|
|
}
|
|
|
|
void interrupt_handler(void)
|
|
{
|
|
if (flag)
|
|
process_message(msg);
|
|
}
|
|
|
|
There is nothing to prevent the compiler from transforming
|
|
process_level() to the following, in fact, this might well be a
|
|
win for single-threaded code:
|
|
|
|
void process_level(void)
|
|
{
|
|
flag = true;
|
|
msg = get_message();
|
|
}
|
|
|
|
If the interrupt occurs between these two statement, then
|
|
interrupt_handler() might be passed a garbled msg. Use WRITE_ONCE()
|
|
to prevent this as follows:
|
|
|
|
void process_level(void)
|
|
{
|
|
WRITE_ONCE(msg, get_message());
|
|
WRITE_ONCE(flag, true);
|
|
}
|
|
|
|
void interrupt_handler(void)
|
|
{
|
|
if (READ_ONCE(flag))
|
|
process_message(READ_ONCE(msg));
|
|
}
|
|
|
|
Note that the READ_ONCE() and WRITE_ONCE() wrappers in
|
|
interrupt_handler() are needed if this interrupt handler can itself
|
|
be interrupted by something that also accesses 'flag' and 'msg',
|
|
for example, a nested interrupt or an NMI. Otherwise, READ_ONCE()
|
|
and WRITE_ONCE() are not needed in interrupt_handler() other than
|
|
for documentation purposes. (Note also that nested interrupts
|
|
do not typically occur in modern Linux kernels, in fact, if an
|
|
interrupt handler returns with interrupts enabled, you will get a
|
|
WARN_ONCE() splat.)
|
|
|
|
You should assume that the compiler can move READ_ONCE() and
|
|
WRITE_ONCE() past code not containing READ_ONCE(), WRITE_ONCE(),
|
|
barrier(), or similar primitives.
|
|
|
|
This effect could also be achieved using barrier(), but READ_ONCE()
|
|
and WRITE_ONCE() are more selective: With READ_ONCE() and
|
|
WRITE_ONCE(), the compiler need only forget the contents of the
|
|
indicated memory locations, while with barrier() the compiler must
|
|
discard the value of all memory locations that it has currented
|
|
cached in any machine registers. Of course, the compiler must also
|
|
respect the order in which the READ_ONCE()s and WRITE_ONCE()s occur,
|
|
though the CPU of course need not do so.
|
|
|
|
(*) The compiler is within its rights to invent stores to a variable,
|
|
as in the following example:
|
|
|
|
if (a)
|
|
b = a;
|
|
else
|
|
b = 42;
|
|
|
|
The compiler might save a branch by optimizing this as follows:
|
|
|
|
b = 42;
|
|
if (a)
|
|
b = a;
|
|
|
|
In single-threaded code, this is not only safe, but also saves
|
|
a branch. Unfortunately, in concurrent code, this optimization
|
|
could cause some other CPU to see a spurious value of 42 -- even
|
|
if variable 'a' was never zero -- when loading variable 'b'.
|
|
Use WRITE_ONCE() to prevent this as follows:
|
|
|
|
if (a)
|
|
WRITE_ONCE(b, a);
|
|
else
|
|
WRITE_ONCE(b, 42);
|
|
|
|
The compiler can also invent loads. These are usually less
|
|
damaging, but they can result in cache-line bouncing and thus in
|
|
poor performance and scalability. Use READ_ONCE() to prevent
|
|
invented loads.
|
|
|
|
(*) For aligned memory locations whose size allows them to be accessed
|
|
with a single memory-reference instruction, prevents "load tearing"
|
|
and "store tearing," in which a single large access is replaced by
|
|
multiple smaller accesses. For example, given an architecture having
|
|
16-bit store instructions with 7-bit immediate fields, the compiler
|
|
might be tempted to use two 16-bit store-immediate instructions to
|
|
implement the following 32-bit store:
|
|
|
|
p = 0x00010002;
|
|
|
|
Please note that GCC really does use this sort of optimization,
|
|
which is not surprising given that it would likely take more
|
|
than two instructions to build the constant and then store it.
|
|
This optimization can therefore be a win in single-threaded code.
|
|
In fact, a recent bug (since fixed) caused GCC to incorrectly use
|
|
this optimization in a volatile store. In the absence of such bugs,
|
|
use of WRITE_ONCE() prevents store tearing in the following example:
|
|
|
|
WRITE_ONCE(p, 0x00010002);
|
|
|
|
Use of packed structures can also result in load and store tearing,
|
|
as in this example:
|
|
|
|
struct __attribute__((__packed__)) foo {
|
|
short a;
|
|
int b;
|
|
short c;
|
|
};
|
|
struct foo foo1, foo2;
|
|
...
|
|
|
|
foo2.a = foo1.a;
|
|
foo2.b = foo1.b;
|
|
foo2.c = foo1.c;
|
|
|
|
Because there are no READ_ONCE() or WRITE_ONCE() wrappers and no
|
|
volatile markings, the compiler would be well within its rights to
|
|
implement these three assignment statements as a pair of 32-bit
|
|
loads followed by a pair of 32-bit stores. This would result in
|
|
load tearing on 'foo1.b' and store tearing on 'foo2.b'. READ_ONCE()
|
|
and WRITE_ONCE() again prevent tearing in this example:
|
|
|
|
foo2.a = foo1.a;
|
|
WRITE_ONCE(foo2.b, READ_ONCE(foo1.b));
|
|
foo2.c = foo1.c;
|
|
|
|
All that aside, it is never necessary to use READ_ONCE() and
|
|
WRITE_ONCE() on a variable that has been marked volatile. For example,
|
|
because 'jiffies' is marked volatile, it is never necessary to
|
|
say READ_ONCE(jiffies). The reason for this is that READ_ONCE() and
|
|
WRITE_ONCE() are implemented as volatile casts, which has no effect when
|
|
its argument is already marked volatile.
|
|
|
|
Please note that these compiler barriers have no direct effect on the CPU,
|
|
which may then reorder things however it wishes.
|
|
|
|
|
|
CPU MEMORY BARRIERS
|
|
-------------------
|
|
|
|
The Linux kernel has eight basic CPU memory barriers:
|
|
|
|
TYPE MANDATORY SMP CONDITIONAL
|
|
=============== ======================= ===========================
|
|
GENERAL mb() smp_mb()
|
|
WRITE wmb() smp_wmb()
|
|
READ rmb() smp_rmb()
|
|
DATA DEPENDENCY READ_ONCE()
|
|
|
|
|
|
All memory barriers except the data dependency barriers imply a compiler
|
|
barrier. Data dependencies do not impose any additional compiler ordering.
|
|
|
|
Aside: In the case of data dependencies, the compiler would be expected
|
|
to issue the loads in the correct order (eg. `a[b]` would have to load
|
|
the value of b before loading a[b]), however there is no guarantee in
|
|
the C specification that the compiler may not speculate the value of b
|
|
(eg. is equal to 1) and load a before b (eg. tmp = a[1]; if (b != 1)
|
|
tmp = a[b]; ). There is also the problem of a compiler reloading b after
|
|
having loaded a[b], thus having a newer copy of b than a[b]. A consensus
|
|
has not yet been reached about these problems, however the READ_ONCE()
|
|
macro is a good place to start looking.
|
|
|
|
SMP memory barriers are reduced to compiler barriers on uniprocessor compiled
|
|
systems because it is assumed that a CPU will appear to be self-consistent,
|
|
and will order overlapping accesses correctly with respect to itself.
|
|
However, see the subsection on "Virtual Machine Guests" below.
|
|
|
|
[!] Note that SMP memory barriers _must_ be used to control the ordering of
|
|
references to shared memory on SMP systems, though the use of locking instead
|
|
is sufficient.
|
|
|
|
Mandatory barriers should not be used to control SMP effects, since mandatory
|
|
barriers impose unnecessary overhead on both SMP and UP systems. They may,
|
|
however, be used to control MMIO effects on accesses through relaxed memory I/O
|
|
windows. These barriers are required even on non-SMP systems as they affect
|
|
the order in which memory operations appear to a device by prohibiting both the
|
|
compiler and the CPU from reordering them.
|
|
|
|
|
|
There are some more advanced barrier functions:
|
|
|
|
(*) smp_store_mb(var, value)
|
|
|
|
This assigns the value to the variable and then inserts a full memory
|
|
barrier after it. It isn't guaranteed to insert anything more than a
|
|
compiler barrier in a UP compilation.
|
|
|
|
|
|
(*) smp_mb__before_atomic();
|
|
(*) smp_mb__after_atomic();
|
|
|
|
These are for use with atomic (such as add, subtract, increment and
|
|
decrement) functions that don't return a value, especially when used for
|
|
reference counting. These functions do not imply memory barriers.
|
|
|
|
These are also used for atomic bitop functions that do not return a
|
|
value (such as set_bit and clear_bit).
|
|
|
|
As an example, consider a piece of code that marks an object as being dead
|
|
and then decrements the object's reference count:
|
|
|
|
obj->dead = 1;
|
|
smp_mb__before_atomic();
|
|
atomic_dec(&obj->ref_count);
|
|
|
|
This makes sure that the death mark on the object is perceived to be set
|
|
*before* the reference counter is decremented.
|
|
|
|
See Documentation/atomic_{t,bitops}.txt for more information.
|
|
|
|
|
|
(*) dma_wmb();
|
|
(*) dma_rmb();
|
|
|
|
These are for use with consistent memory to guarantee the ordering
|
|
of writes or reads of shared memory accessible to both the CPU and a
|
|
DMA capable device.
|
|
|
|
For example, consider a device driver that shares memory with a device
|
|
and uses a descriptor status value to indicate if the descriptor belongs
|
|
to the device or the CPU, and a doorbell to notify it when new
|
|
descriptors are available:
|
|
|
|
if (desc->status != DEVICE_OWN) {
|
|
/* do not read data until we own descriptor */
|
|
dma_rmb();
|
|
|
|
/* read/modify data */
|
|
read_data = desc->data;
|
|
desc->data = write_data;
|
|
|
|
/* flush modifications before status update */
|
|
dma_wmb();
|
|
|
|
/* assign ownership */
|
|
desc->status = DEVICE_OWN;
|
|
|
|
/* notify device of new descriptors */
|
|
writel(DESC_NOTIFY, doorbell);
|
|
}
|
|
|
|
The dma_rmb() allows us guarantee the device has released ownership
|
|
before we read the data from the descriptor, and the dma_wmb() allows
|
|
us to guarantee the data is written to the descriptor before the device
|
|
can see it now has ownership. Note that, when using writel(), a prior
|
|
wmb() is not needed to guarantee that the cache coherent memory writes
|
|
have completed before writing to the MMIO region. The cheaper
|
|
writel_relaxed() does not provide this guarantee and must not be used
|
|
here.
|
|
|
|
See the subsection "Kernel I/O barrier effects" for more information on
|
|
relaxed I/O accessors and the Documentation/DMA-API.txt file for more
|
|
information on consistent memory.
|
|
|
|
|
|
===============================
|
|
IMPLICIT KERNEL MEMORY BARRIERS
|
|
===============================
|
|
|
|
Some of the other functions in the linux kernel imply memory barriers, amongst
|
|
which are locking and scheduling functions.
|
|
|
|
This specification is a _minimum_ guarantee; any particular architecture may
|
|
provide more substantial guarantees, but these may not be relied upon outside
|
|
of arch specific code.
|
|
|
|
|
|
LOCK ACQUISITION FUNCTIONS
|
|
--------------------------
|
|
|
|
The Linux kernel has a number of locking constructs:
|
|
|
|
(*) spin locks
|
|
(*) R/W spin locks
|
|
(*) mutexes
|
|
(*) semaphores
|
|
(*) R/W semaphores
|
|
|
|
In all cases there are variants on "ACQUIRE" operations and "RELEASE" operations
|
|
for each construct. These operations all imply certain barriers:
|
|
|
|
(1) ACQUIRE operation implication:
|
|
|
|
Memory operations issued after the ACQUIRE will be completed after the
|
|
ACQUIRE operation has completed.
|
|
|
|
Memory operations issued before the ACQUIRE may be completed after
|
|
the ACQUIRE operation has completed.
|
|
|
|
(2) RELEASE operation implication:
|
|
|
|
Memory operations issued before the RELEASE will be completed before the
|
|
RELEASE operation has completed.
|
|
|
|
Memory operations issued after the RELEASE may be completed before the
|
|
RELEASE operation has completed.
|
|
|
|
(3) ACQUIRE vs ACQUIRE implication:
|
|
|
|
All ACQUIRE operations issued before another ACQUIRE operation will be
|
|
completed before that ACQUIRE operation.
|
|
|
|
(4) ACQUIRE vs RELEASE implication:
|
|
|
|
All ACQUIRE operations issued before a RELEASE operation will be
|
|
completed before the RELEASE operation.
|
|
|
|
(5) Failed conditional ACQUIRE implication:
|
|
|
|
Certain locking variants of the ACQUIRE operation may fail, either due to
|
|
being unable to get the lock immediately, or due to receiving an unblocked
|
|
signal while asleep waiting for the lock to become available. Failed
|
|
locks do not imply any sort of barrier.
|
|
|
|
[!] Note: one of the consequences of lock ACQUIREs and RELEASEs being only
|
|
one-way barriers is that the effects of instructions outside of a critical
|
|
section may seep into the inside of the critical section.
|
|
|
|
An ACQUIRE followed by a RELEASE may not be assumed to be full memory barrier
|
|
because it is possible for an access preceding the ACQUIRE to happen after the
|
|
ACQUIRE, and an access following the RELEASE to happen before the RELEASE, and
|
|
the two accesses can themselves then cross:
|
|
|
|
*A = a;
|
|
ACQUIRE M
|
|
RELEASE M
|
|
*B = b;
|
|
|
|
may occur as:
|
|
|
|
ACQUIRE M, STORE *B, STORE *A, RELEASE M
|
|
|
|
When the ACQUIRE and RELEASE are a lock acquisition and release,
|
|
respectively, this same reordering can occur if the lock's ACQUIRE and
|
|
RELEASE are to the same lock variable, but only from the perspective of
|
|
another CPU not holding that lock. In short, a ACQUIRE followed by an
|
|
RELEASE may -not- be assumed to be a full memory barrier.
|
|
|
|
Similarly, the reverse case of a RELEASE followed by an ACQUIRE does
|
|
not imply a full memory barrier. Therefore, the CPU's execution of the
|
|
critical sections corresponding to the RELEASE and the ACQUIRE can cross,
|
|
so that:
|
|
|
|
*A = a;
|
|
RELEASE M
|
|
ACQUIRE N
|
|
*B = b;
|
|
|
|
could occur as:
|
|
|
|
ACQUIRE N, STORE *B, STORE *A, RELEASE M
|
|
|
|
It might appear that this reordering could introduce a deadlock.
|
|
However, this cannot happen because if such a deadlock threatened,
|
|
the RELEASE would simply complete, thereby avoiding the deadlock.
|
|
|
|
Why does this work?
|
|
|
|
One key point is that we are only talking about the CPU doing
|
|
the reordering, not the compiler. If the compiler (or, for
|
|
that matter, the developer) switched the operations, deadlock
|
|
-could- occur.
|
|
|
|
But suppose the CPU reordered the operations. In this case,
|
|
the unlock precedes the lock in the assembly code. The CPU
|
|
simply elected to try executing the later lock operation first.
|
|
If there is a deadlock, this lock operation will simply spin (or
|
|
try to sleep, but more on that later). The CPU will eventually
|
|
execute the unlock operation (which preceded the lock operation
|
|
in the assembly code), which will unravel the potential deadlock,
|
|
allowing the lock operation to succeed.
|
|
|
|
But what if the lock is a sleeplock? In that case, the code will
|
|
try to enter the scheduler, where it will eventually encounter
|
|
a memory barrier, which will force the earlier unlock operation
|
|
to complete, again unraveling the deadlock. There might be
|
|
a sleep-unlock race, but the locking primitive needs to resolve
|
|
such races properly in any case.
|
|
|
|
Locks and semaphores may not provide any guarantee of ordering on UP compiled
|
|
systems, and so cannot be counted on in such a situation to actually achieve
|
|
anything at all - especially with respect to I/O accesses - unless combined
|
|
with interrupt disabling operations.
|
|
|
|
See also the section on "Inter-CPU acquiring barrier effects".
|
|
|
|
|
|
As an example, consider the following:
|
|
|
|
*A = a;
|
|
*B = b;
|
|
ACQUIRE
|
|
*C = c;
|
|
*D = d;
|
|
RELEASE
|
|
*E = e;
|
|
*F = f;
|
|
|
|
The following sequence of events is acceptable:
|
|
|
|
ACQUIRE, {*F,*A}, *E, {*C,*D}, *B, RELEASE
|
|
|
|
[+] Note that {*F,*A} indicates a combined access.
|
|
|
|
But none of the following are:
|
|
|
|
{*F,*A}, *B, ACQUIRE, *C, *D, RELEASE, *E
|
|
*A, *B, *C, ACQUIRE, *D, RELEASE, *E, *F
|
|
*A, *B, ACQUIRE, *C, RELEASE, *D, *E, *F
|
|
*B, ACQUIRE, *C, *D, RELEASE, {*F,*A}, *E
|
|
|
|
|
|
|
|
INTERRUPT DISABLING FUNCTIONS
|
|
-----------------------------
|
|
|
|
Functions that disable interrupts (ACQUIRE equivalent) and enable interrupts
|
|
(RELEASE equivalent) will act as compiler barriers only. So if memory or I/O
|
|
barriers are required in such a situation, they must be provided from some
|
|
other means.
|
|
|
|
|
|
SLEEP AND WAKE-UP FUNCTIONS
|
|
---------------------------
|
|
|
|
Sleeping and waking on an event flagged in global data can be viewed as an
|
|
interaction between two pieces of data: the task state of the task waiting for
|
|
the event and the global data used to indicate the event. To make sure that
|
|
these appear to happen in the right order, the primitives to begin the process
|
|
of going to sleep, and the primitives to initiate a wake up imply certain
|
|
barriers.
|
|
|
|
Firstly, the sleeper normally follows something like this sequence of events:
|
|
|
|
for (;;) {
|
|
set_current_state(TASK_UNINTERRUPTIBLE);
|
|
if (event_indicated)
|
|
break;
|
|
schedule();
|
|
}
|
|
|
|
A general memory barrier is interpolated automatically by set_current_state()
|
|
after it has altered the task state:
|
|
|
|
CPU 1
|
|
===============================
|
|
set_current_state();
|
|
smp_store_mb();
|
|
STORE current->state
|
|
<general barrier>
|
|
LOAD event_indicated
|
|
|
|
set_current_state() may be wrapped by:
|
|
|
|
prepare_to_wait();
|
|
prepare_to_wait_exclusive();
|
|
|
|
which therefore also imply a general memory barrier after setting the state.
|
|
The whole sequence above is available in various canned forms, all of which
|
|
interpolate the memory barrier in the right place:
|
|
|
|
wait_event();
|
|
wait_event_interruptible();
|
|
wait_event_interruptible_exclusive();
|
|
wait_event_interruptible_timeout();
|
|
wait_event_killable();
|
|
wait_event_timeout();
|
|
wait_on_bit();
|
|
wait_on_bit_lock();
|
|
|
|
|
|
Secondly, code that performs a wake up normally follows something like this:
|
|
|
|
event_indicated = 1;
|
|
wake_up(&event_wait_queue);
|
|
|
|
or:
|
|
|
|
event_indicated = 1;
|
|
wake_up_process(event_daemon);
|
|
|
|
A general memory barrier is executed by wake_up() if it wakes something up.
|
|
If it doesn't wake anything up then a memory barrier may or may not be
|
|
executed; you must not rely on it. The barrier occurs before the task state
|
|
is accessed, in particular, it sits between the STORE to indicate the event
|
|
and the STORE to set TASK_RUNNING:
|
|
|
|
CPU 1 (Sleeper) CPU 2 (Waker)
|
|
=============================== ===============================
|
|
set_current_state(); STORE event_indicated
|
|
smp_store_mb(); wake_up();
|
|
STORE current->state ...
|
|
<general barrier> <general barrier>
|
|
LOAD event_indicated if ((LOAD task->state) & TASK_NORMAL)
|
|
STORE task->state
|
|
|
|
where "task" is the thread being woken up and it equals CPU 1's "current".
|
|
|
|
To repeat, a general memory barrier is guaranteed to be executed by wake_up()
|
|
if something is actually awakened, but otherwise there is no such guarantee.
|
|
To see this, consider the following sequence of events, where X and Y are both
|
|
initially zero:
|
|
|
|
CPU 1 CPU 2
|
|
=============================== ===============================
|
|
X = 1; Y = 1;
|
|
smp_mb(); wake_up();
|
|
LOAD Y LOAD X
|
|
|
|
If a wakeup does occur, one (at least) of the two loads must see 1. If, on
|
|
the other hand, a wakeup does not occur, both loads might see 0.
|
|
|
|
wake_up_process() always executes a general memory barrier. The barrier again
|
|
occurs before the task state is accessed. In particular, if the wake_up() in
|
|
the previous snippet were replaced by a call to wake_up_process() then one of
|
|
the two loads would be guaranteed to see 1.
|
|
|
|
The available waker functions include:
|
|
|
|
complete();
|
|
wake_up();
|
|
wake_up_all();
|
|
wake_up_bit();
|
|
wake_up_interruptible();
|
|
wake_up_interruptible_all();
|
|
wake_up_interruptible_nr();
|
|
wake_up_interruptible_poll();
|
|
wake_up_interruptible_sync();
|
|
wake_up_interruptible_sync_poll();
|
|
wake_up_locked();
|
|
wake_up_locked_poll();
|
|
wake_up_nr();
|
|
wake_up_poll();
|
|
wake_up_process();
|
|
|
|
In terms of memory ordering, these functions all provide the same guarantees of
|
|
a wake_up() (or stronger).
|
|
|
|
[!] Note that the memory barriers implied by the sleeper and the waker do _not_
|
|
order multiple stores before the wake-up with respect to loads of those stored
|
|
values after the sleeper has called set_current_state(). For instance, if the
|
|
sleeper does:
|
|
|
|
set_current_state(TASK_INTERRUPTIBLE);
|
|
if (event_indicated)
|
|
break;
|
|
__set_current_state(TASK_RUNNING);
|
|
do_something(my_data);
|
|
|
|
and the waker does:
|
|
|
|
my_data = value;
|
|
event_indicated = 1;
|
|
wake_up(&event_wait_queue);
|
|
|
|
there's no guarantee that the change to event_indicated will be perceived by
|
|
the sleeper as coming after the change to my_data. In such a circumstance, the
|
|
code on both sides must interpolate its own memory barriers between the
|
|
separate data accesses. Thus the above sleeper ought to do:
|
|
|
|
set_current_state(TASK_INTERRUPTIBLE);
|
|
if (event_indicated) {
|
|
smp_rmb();
|
|
do_something(my_data);
|
|
}
|
|
|
|
and the waker should do:
|
|
|
|
my_data = value;
|
|
smp_wmb();
|
|
event_indicated = 1;
|
|
wake_up(&event_wait_queue);
|
|
|
|
|
|
MISCELLANEOUS FUNCTIONS
|
|
-----------------------
|
|
|
|
Other functions that imply barriers:
|
|
|
|
(*) schedule() and similar imply full memory barriers.
|
|
|
|
|
|
===================================
|
|
INTER-CPU ACQUIRING BARRIER EFFECTS
|
|
===================================
|
|
|
|
On SMP systems locking primitives give a more substantial form of barrier: one
|
|
that does affect memory access ordering on other CPUs, within the context of
|
|
conflict on any particular lock.
|
|
|
|
|
|
ACQUIRES VS MEMORY ACCESSES
|
|
---------------------------
|
|
|
|
Consider the following: the system has a pair of spinlocks (M) and (Q), and
|
|
three CPUs; then should the following sequence of events occur:
|
|
|
|
CPU 1 CPU 2
|
|
=============================== ===============================
|
|
WRITE_ONCE(*A, a); WRITE_ONCE(*E, e);
|
|
ACQUIRE M ACQUIRE Q
|
|
WRITE_ONCE(*B, b); WRITE_ONCE(*F, f);
|
|
WRITE_ONCE(*C, c); WRITE_ONCE(*G, g);
|
|
RELEASE M RELEASE Q
|
|
WRITE_ONCE(*D, d); WRITE_ONCE(*H, h);
|
|
|
|
Then there is no guarantee as to what order CPU 3 will see the accesses to *A
|
|
through *H occur in, other than the constraints imposed by the separate locks
|
|
on the separate CPUs. It might, for example, see:
|
|
|
|
*E, ACQUIRE M, ACQUIRE Q, *G, *C, *F, *A, *B, RELEASE Q, *D, *H, RELEASE M
|
|
|
|
But it won't see any of:
|
|
|
|
*B, *C or *D preceding ACQUIRE M
|
|
*A, *B or *C following RELEASE M
|
|
*F, *G or *H preceding ACQUIRE Q
|
|
*E, *F or *G following RELEASE Q
|
|
|
|
|
|
=================================
|
|
WHERE ARE MEMORY BARRIERS NEEDED?
|
|
=================================
|
|
|
|
Under normal operation, memory operation reordering is generally not going to
|
|
be a problem as a single-threaded linear piece of code will still appear to
|
|
work correctly, even if it's in an SMP kernel. There are, however, four
|
|
circumstances in which reordering definitely _could_ be a problem:
|
|
|
|
(*) Interprocessor interaction.
|
|
|
|
(*) Atomic operations.
|
|
|
|
(*) Accessing devices.
|
|
|
|
(*) Interrupts.
|
|
|
|
|
|
INTERPROCESSOR INTERACTION
|
|
--------------------------
|
|
|
|
When there's a system with more than one processor, more than one CPU in the
|
|
system may be working on the same data set at the same time. This can cause
|
|
synchronisation problems, and the usual way of dealing with them is to use
|
|
locks. Locks, however, are quite expensive, and so it may be preferable to
|
|
operate without the use of a lock if at all possible. In such a case
|
|
operations that affect both CPUs may have to be carefully ordered to prevent
|
|
a malfunction.
|
|
|
|
Consider, for example, the R/W semaphore slow path. Here a waiting process is
|
|
queued on the semaphore, by virtue of it having a piece of its stack linked to
|
|
the semaphore's list of waiting processes:
|
|
|
|
struct rw_semaphore {
|
|
...
|
|
spinlock_t lock;
|
|
struct list_head waiters;
|
|
};
|
|
|
|
struct rwsem_waiter {
|
|
struct list_head list;
|
|
struct task_struct *task;
|
|
};
|
|
|
|
To wake up a particular waiter, the up_read() or up_write() functions have to:
|
|
|
|
(1) read the next pointer from this waiter's record to know as to where the
|
|
next waiter record is;
|
|
|
|
(2) read the pointer to the waiter's task structure;
|
|
|
|
(3) clear the task pointer to tell the waiter it has been given the semaphore;
|
|
|
|
(4) call wake_up_process() on the task; and
|
|
|
|
(5) release the reference held on the waiter's task struct.
|
|
|
|
In other words, it has to perform this sequence of events:
|
|
|
|
LOAD waiter->list.next;
|
|
LOAD waiter->task;
|
|
STORE waiter->task;
|
|
CALL wakeup
|
|
RELEASE task
|
|
|
|
and if any of these steps occur out of order, then the whole thing may
|
|
malfunction.
|
|
|
|
Once it has queued itself and dropped the semaphore lock, the waiter does not
|
|
get the lock again; it instead just waits for its task pointer to be cleared
|
|
before proceeding. Since the record is on the waiter's stack, this means that
|
|
if the task pointer is cleared _before_ the next pointer in the list is read,
|
|
another CPU might start processing the waiter and might clobber the waiter's
|
|
stack before the up*() function has a chance to read the next pointer.
|
|
|
|
Consider then what might happen to the above sequence of events:
|
|
|
|
CPU 1 CPU 2
|
|
=============================== ===============================
|
|
down_xxx()
|
|
Queue waiter
|
|
Sleep
|
|
up_yyy()
|
|
LOAD waiter->task;
|
|
STORE waiter->task;
|
|
Woken up by other event
|
|
<preempt>
|
|
Resume processing
|
|
down_xxx() returns
|
|
call foo()
|
|
foo() clobbers *waiter
|
|
</preempt>
|
|
LOAD waiter->list.next;
|
|
--- OOPS ---
|
|
|
|
This could be dealt with using the semaphore lock, but then the down_xxx()
|
|
function has to needlessly get the spinlock again after being woken up.
|
|
|
|
The way to deal with this is to insert a general SMP memory barrier:
|
|
|
|
LOAD waiter->list.next;
|
|
LOAD waiter->task;
|
|
smp_mb();
|
|
STORE waiter->task;
|
|
CALL wakeup
|
|
RELEASE task
|
|
|
|
In this case, the barrier makes a guarantee that all memory accesses before the
|
|
barrier will appear to happen before all the memory accesses after the barrier
|
|
with respect to the other CPUs on the system. It does _not_ guarantee that all
|
|
the memory accesses before the barrier will be complete by the time the barrier
|
|
instruction itself is complete.
|
|
|
|
On a UP system - where this wouldn't be a problem - the smp_mb() is just a
|
|
compiler barrier, thus making sure the compiler emits the instructions in the
|
|
right order without actually intervening in the CPU. Since there's only one
|
|
CPU, that CPU's dependency ordering logic will take care of everything else.
|
|
|
|
|
|
ATOMIC OPERATIONS
|
|
-----------------
|
|
|
|
While they are technically interprocessor interaction considerations, atomic
|
|
operations are noted specially as some of them imply full memory barriers and
|
|
some don't, but they're very heavily relied on as a group throughout the
|
|
kernel.
|
|
|
|
See Documentation/atomic_t.txt for more information.
|
|
|
|
|
|
ACCESSING DEVICES
|
|
-----------------
|
|
|
|
Many devices can be memory mapped, and so appear to the CPU as if they're just
|
|
a set of memory locations. To control such a device, the driver usually has to
|
|
make the right memory accesses in exactly the right order.
|
|
|
|
However, having a clever CPU or a clever compiler creates a potential problem
|
|
in that the carefully sequenced accesses in the driver code won't reach the
|
|
device in the requisite order if the CPU or the compiler thinks it is more
|
|
efficient to reorder, combine or merge accesses - something that would cause
|
|
the device to malfunction.
|
|
|
|
Inside of the Linux kernel, I/O should be done through the appropriate accessor
|
|
routines - such as inb() or writel() - which know how to make such accesses
|
|
appropriately sequential. While this, for the most part, renders the explicit
|
|
use of memory barriers unnecessary, if the accessor functions are used to refer
|
|
to an I/O memory window with relaxed memory access properties, then _mandatory_
|
|
memory barriers are required to enforce ordering.
|
|
|
|
See Documentation/driver-api/device-io.rst for more information.
|
|
|
|
|
|
INTERRUPTS
|
|
----------
|
|
|
|
A driver may be interrupted by its own interrupt service routine, and thus the
|
|
two parts of the driver may interfere with each other's attempts to control or
|
|
access the device.
|
|
|
|
This may be alleviated - at least in part - by disabling local interrupts (a
|
|
form of locking), such that the critical operations are all contained within
|
|
the interrupt-disabled section in the driver. While the driver's interrupt
|
|
routine is executing, the driver's core may not run on the same CPU, and its
|
|
interrupt is not permitted to happen again until the current interrupt has been
|
|
handled, thus the interrupt handler does not need to lock against that.
|
|
|
|
However, consider a driver that was talking to an ethernet card that sports an
|
|
address register and a data register. If that driver's core talks to the card
|
|
under interrupt-disablement and then the driver's interrupt handler is invoked:
|
|
|
|
LOCAL IRQ DISABLE
|
|
writew(ADDR, 3);
|
|
writew(DATA, y);
|
|
LOCAL IRQ ENABLE
|
|
<interrupt>
|
|
writew(ADDR, 4);
|
|
q = readw(DATA);
|
|
</interrupt>
|
|
|
|
The store to the data register might happen after the second store to the
|
|
address register if ordering rules are sufficiently relaxed:
|
|
|
|
STORE *ADDR = 3, STORE *ADDR = 4, STORE *DATA = y, q = LOAD *DATA
|
|
|
|
|
|
If ordering rules are relaxed, it must be assumed that accesses done inside an
|
|
interrupt disabled section may leak outside of it and may interleave with
|
|
accesses performed in an interrupt - and vice versa - unless implicit or
|
|
explicit barriers are used.
|
|
|
|
Normally this won't be a problem because the I/O accesses done inside such
|
|
sections will include synchronous load operations on strictly ordered I/O
|
|
registers that form implicit I/O barriers.
|
|
|
|
|
|
A similar situation may occur between an interrupt routine and two routines
|
|
running on separate CPUs that communicate with each other. If such a case is
|
|
likely, then interrupt-disabling locks should be used to guarantee ordering.
|
|
|
|
|
|
==========================
|
|
KERNEL I/O BARRIER EFFECTS
|
|
==========================
|
|
|
|
Interfacing with peripherals via I/O accesses is deeply architecture and device
|
|
specific. Therefore, drivers which are inherently non-portable may rely on
|
|
specific behaviours of their target systems in order to achieve synchronization
|
|
in the most lightweight manner possible. For drivers intending to be portable
|
|
between multiple architectures and bus implementations, the kernel offers a
|
|
series of accessor functions that provide various degrees of ordering
|
|
guarantees:
|
|
|
|
(*) readX(), writeX():
|
|
|
|
The readX() and writeX() MMIO accessors take a pointer to the
|
|
peripheral being accessed as an __iomem * parameter. For pointers
|
|
mapped with the default I/O attributes (e.g. those returned by
|
|
ioremap()), the ordering guarantees are as follows:
|
|
|
|
1. All readX() and writeX() accesses to the same peripheral are ordered
|
|
with respect to each other. This ensures that MMIO register accesses
|
|
by the same CPU thread to a particular device will arrive in program
|
|
order.
|
|
|
|
2. A writeX() issued by a CPU thread holding a spinlock is ordered
|
|
before a writeX() to the same peripheral from another CPU thread
|
|
issued after a later acquisition of the same spinlock. This ensures
|
|
that MMIO register writes to a particular device issued while holding
|
|
a spinlock will arrive in an order consistent with acquisitions of
|
|
the lock.
|
|
|
|
3. A writeX() by a CPU thread to the peripheral will first wait for the
|
|
completion of all prior writes to memory either issued by, or
|
|
propagated to, the same thread. This ensures that writes by the CPU
|
|
to an outbound DMA buffer allocated by dma_alloc_coherent() will be
|
|
visible to a DMA engine when the CPU writes to its MMIO control
|
|
register to trigger the transfer.
|
|
|
|
4. A readX() by a CPU thread from the peripheral will complete before
|
|
any subsequent reads from memory by the same thread can begin. This
|
|
ensures that reads by the CPU from an incoming DMA buffer allocated
|
|
by dma_alloc_coherent() will not see stale data after reading from
|
|
the DMA engine's MMIO status register to establish that the DMA
|
|
transfer has completed.
|
|
|
|
5. A readX() by a CPU thread from the peripheral will complete before
|
|
any subsequent delay() loop can begin execution on the same thread.
|
|
This ensures that two MMIO register writes by the CPU to a peripheral
|
|
will arrive at least 1us apart if the first write is immediately read
|
|
back with readX() and udelay(1) is called prior to the second
|
|
writeX():
|
|
|
|
writel(42, DEVICE_REGISTER_0); // Arrives at the device...
|
|
readl(DEVICE_REGISTER_0);
|
|
udelay(1);
|
|
writel(42, DEVICE_REGISTER_1); // ...at least 1us before this.
|
|
|
|
The ordering properties of __iomem pointers obtained with non-default
|
|
attributes (e.g. those returned by ioremap_wc()) are specific to the
|
|
underlying architecture and therefore the guarantees listed above cannot
|
|
generally be relied upon for accesses to these types of mappings.
|
|
|
|
(*) readX_relaxed(), writeX_relaxed():
|
|
|
|
These are similar to readX() and writeX(), but provide weaker memory
|
|
ordering guarantees. Specifically, they do not guarantee ordering with
|
|
respect to locking, normal memory accesses or delay() loops (i.e.
|
|
bullets 2-5 above) but they are still guaranteed to be ordered with
|
|
respect to other accesses from the same CPU thread to the same
|
|
peripheral when operating on __iomem pointers mapped with the default
|
|
I/O attributes.
|
|
|
|
(*) readsX(), writesX():
|
|
|
|
The readsX() and writesX() MMIO accessors are designed for accessing
|
|
register-based, memory-mapped FIFOs residing on peripherals that are not
|
|
capable of performing DMA. Consequently, they provide only the ordering
|
|
guarantees of readX_relaxed() and writeX_relaxed(), as documented above.
|
|
|
|
(*) inX(), outX():
|
|
|
|
The inX() and outX() accessors are intended to access legacy port-mapped
|
|
I/O peripherals, which may require special instructions on some
|
|
architectures (notably x86). The port number of the peripheral being
|
|
accessed is passed as an argument.
|
|
|
|
Since many CPU architectures ultimately access these peripherals via an
|
|
internal virtual memory mapping, the portable ordering guarantees
|
|
provided by inX() and outX() are the same as those provided by readX()
|
|
and writeX() respectively when accessing a mapping with the default I/O
|
|
attributes.
|
|
|
|
Device drivers may expect outX() to emit a non-posted write transaction
|
|
that waits for a completion response from the I/O peripheral before
|
|
returning. This is not guaranteed by all architectures and is therefore
|
|
not part of the portable ordering semantics.
|
|
|
|
(*) insX(), outsX():
|
|
|
|
As above, the insX() and outsX() accessors provide the same ordering
|
|
guarantees as readsX() and writesX() respectively when accessing a
|
|
mapping with the default I/O attributes.
|
|
|
|
(*) ioreadX(), iowriteX():
|
|
|
|
These will perform appropriately for the type of access they're actually
|
|
doing, be it inX()/outX() or readX()/writeX().
|
|
|
|
With the exception of the string accessors (insX(), outsX(), readsX() and
|
|
writesX()), all of the above assume that the underlying peripheral is
|
|
little-endian and will therefore perform byte-swapping operations on big-endian
|
|
architectures.
|
|
|
|
|
|
========================================
|
|
ASSUMED MINIMUM EXECUTION ORDERING MODEL
|
|
========================================
|
|
|
|
It has to be assumed that the conceptual CPU is weakly-ordered but that it will
|
|
maintain the appearance of program causality with respect to itself. Some CPUs
|
|
(such as i386 or x86_64) are more constrained than others (such as powerpc or
|
|
frv), and so the most relaxed case (namely DEC Alpha) must be assumed outside
|
|
of arch-specific code.
|
|
|
|
This means that it must be considered that the CPU will execute its instruction
|
|
stream in any order it feels like - or even in parallel - provided that if an
|
|
instruction in the stream depends on an earlier instruction, then that
|
|
earlier instruction must be sufficiently complete[*] before the later
|
|
instruction may proceed; in other words: provided that the appearance of
|
|
causality is maintained.
|
|
|
|
[*] Some instructions have more than one effect - such as changing the
|
|
condition codes, changing registers or changing memory - and different
|
|
instructions may depend on different effects.
|
|
|
|
A CPU may also discard any instruction sequence that winds up having no
|
|
ultimate effect. For example, if two adjacent instructions both load an
|
|
immediate value into the same register, the first may be discarded.
|
|
|
|
|
|
Similarly, it has to be assumed that compiler might reorder the instruction
|
|
stream in any way it sees fit, again provided the appearance of causality is
|
|
maintained.
|
|
|
|
|
|
============================
|
|
THE EFFECTS OF THE CPU CACHE
|
|
============================
|
|
|
|
The way cached memory operations are perceived across the system is affected to
|
|
a certain extent by the caches that lie between CPUs and memory, and by the
|
|
memory coherence system that maintains the consistency of state in the system.
|
|
|
|
As far as the way a CPU interacts with another part of the system through the
|
|
caches goes, the memory system has to include the CPU's caches, and memory
|
|
barriers for the most part act at the interface between the CPU and its cache
|
|
(memory barriers logically act on the dotted line in the following diagram):
|
|
|
|
<--- CPU ---> : <----------- Memory ----------->
|
|
:
|
|
+--------+ +--------+ : +--------+ +-----------+
|
|
| | | | : | | | | +--------+
|
|
| CPU | | Memory | : | CPU | | | | |
|
|
| Core |--->| Access |----->| Cache |<-->| | | |
|
|
| | | Queue | : | | | |--->| Memory |
|
|
| | | | : | | | | | |
|
|
+--------+ +--------+ : +--------+ | | | |
|
|
: | Cache | +--------+
|
|
: | Coherency |
|
|
: | Mechanism | +--------+
|
|
+--------+ +--------+ : +--------+ | | | |
|
|
| | | | : | | | | | |
|
|
| CPU | | Memory | : | CPU | | |--->| Device |
|
|
| Core |--->| Access |----->| Cache |<-->| | | |
|
|
| | | Queue | : | | | | | |
|
|
| | | | : | | | | +--------+
|
|
+--------+ +--------+ : +--------+ +-----------+
|
|
:
|
|
:
|
|
|
|
Although any particular load or store may not actually appear outside of the
|
|
CPU that issued it since it may have been satisfied within the CPU's own cache,
|
|
it will still appear as if the full memory access had taken place as far as the
|
|
other CPUs are concerned since the cache coherency mechanisms will migrate the
|
|
cacheline over to the accessing CPU and propagate the effects upon conflict.
|
|
|
|
The CPU core may execute instructions in any order it deems fit, provided the
|
|
expected program causality appears to be maintained. Some of the instructions
|
|
generate load and store operations which then go into the queue of memory
|
|
accesses to be performed. The core may place these in the queue in any order
|
|
it wishes, and continue execution until it is forced to wait for an instruction
|
|
to complete.
|
|
|
|
What memory barriers are concerned with is controlling the order in which
|
|
accesses cross from the CPU side of things to the memory side of things, and
|
|
the order in which the effects are perceived to happen by the other observers
|
|
in the system.
|
|
|
|
[!] Memory barriers are _not_ needed within a given CPU, as CPUs always see
|
|
their own loads and stores as if they had happened in program order.
|
|
|
|
[!] MMIO or other device accesses may bypass the cache system. This depends on
|
|
the properties of the memory window through which devices are accessed and/or
|
|
the use of any special device communication instructions the CPU may have.
|
|
|
|
|
|
CACHE COHERENCY
|
|
---------------
|
|
|
|
Life isn't quite as simple as it may appear above, however: for while the
|
|
caches are expected to be coherent, there's no guarantee that that coherency
|
|
will be ordered. This means that while changes made on one CPU will
|
|
eventually become visible on all CPUs, there's no guarantee that they will
|
|
become apparent in the same order on those other CPUs.
|
|
|
|
|
|
Consider dealing with a system that has a pair of CPUs (1 & 2), each of which
|
|
has a pair of parallel data caches (CPU 1 has A/B, and CPU 2 has C/D):
|
|
|
|
:
|
|
: +--------+
|
|
: +---------+ | |
|
|
+--------+ : +--->| Cache A |<------->| |
|
|
| | : | +---------+ | |
|
|
| CPU 1 |<---+ | |
|
|
| | : | +---------+ | |
|
|
+--------+ : +--->| Cache B |<------->| |
|
|
: +---------+ | |
|
|
: | Memory |
|
|
: +---------+ | System |
|
|
+--------+ : +--->| Cache C |<------->| |
|
|
| | : | +---------+ | |
|
|
| CPU 2 |<---+ | |
|
|
| | : | +---------+ | |
|
|
+--------+ : +--->| Cache D |<------->| |
|
|
: +---------+ | |
|
|
: +--------+
|
|
:
|
|
|
|
Imagine the system has the following properties:
|
|
|
|
(*) an odd-numbered cache line may be in cache A, cache C or it may still be
|
|
resident in memory;
|
|
|
|
(*) an even-numbered cache line may be in cache B, cache D or it may still be
|
|
resident in memory;
|
|
|
|
(*) while the CPU core is interrogating one cache, the other cache may be
|
|
making use of the bus to access the rest of the system - perhaps to
|
|
displace a dirty cacheline or to do a speculative load;
|
|
|
|
(*) each cache has a queue of operations that need to be applied to that cache
|
|
to maintain coherency with the rest of the system;
|
|
|
|
(*) the coherency queue is not flushed by normal loads to lines already
|
|
present in the cache, even though the contents of the queue may
|
|
potentially affect those loads.
|
|
|
|
Imagine, then, that two writes are made on the first CPU, with a write barrier
|
|
between them to guarantee that they will appear to reach that CPU's caches in
|
|
the requisite order:
|
|
|
|
CPU 1 CPU 2 COMMENT
|
|
=============== =============== =======================================
|
|
u == 0, v == 1 and p == &u, q == &u
|
|
v = 2;
|
|
smp_wmb(); Make sure change to v is visible before
|
|
change to p
|
|
<A:modify v=2> v is now in cache A exclusively
|
|
p = &v;
|
|
<B:modify p=&v> p is now in cache B exclusively
|
|
|
|
The write memory barrier forces the other CPUs in the system to perceive that
|
|
the local CPU's caches have apparently been updated in the correct order. But
|
|
now imagine that the second CPU wants to read those values:
|
|
|
|
CPU 1 CPU 2 COMMENT
|
|
=============== =============== =======================================
|
|
...
|
|
q = p;
|
|
x = *q;
|
|
|
|
The above pair of reads may then fail to happen in the expected order, as the
|
|
cacheline holding p may get updated in one of the second CPU's caches while
|
|
the update to the cacheline holding v is delayed in the other of the second
|
|
CPU's caches by some other cache event:
|
|
|
|
CPU 1 CPU 2 COMMENT
|
|
=============== =============== =======================================
|
|
u == 0, v == 1 and p == &u, q == &u
|
|
v = 2;
|
|
smp_wmb();
|
|
<A:modify v=2> <C:busy>
|
|
<C:queue v=2>
|
|
p = &v; q = p;
|
|
<D:request p>
|
|
<B:modify p=&v> <D:commit p=&v>
|
|
<D:read p>
|
|
x = *q;
|
|
<C:read *q> Reads from v before v updated in cache
|
|
<C:unbusy>
|
|
<C:commit v=2>
|
|
|
|
Basically, while both cachelines will be updated on CPU 2 eventually, there's
|
|
no guarantee that, without intervention, the order of update will be the same
|
|
as that committed on CPU 1.
|
|
|
|
|
|
To intervene, we need to interpolate a data dependency barrier or a read
|
|
barrier between the loads (which as of v4.15 is supplied unconditionally
|
|
by the READ_ONCE() macro). This will force the cache to commit its
|
|
coherency queue before processing any further requests:
|
|
|
|
CPU 1 CPU 2 COMMENT
|
|
=============== =============== =======================================
|
|
u == 0, v == 1 and p == &u, q == &u
|
|
v = 2;
|
|
smp_wmb();
|
|
<A:modify v=2> <C:busy>
|
|
<C:queue v=2>
|
|
p = &v; q = p;
|
|
<D:request p>
|
|
<B:modify p=&v> <D:commit p=&v>
|
|
<D:read p>
|
|
smp_read_barrier_depends()
|
|
<C:unbusy>
|
|
<C:commit v=2>
|
|
x = *q;
|
|
<C:read *q> Reads from v after v updated in cache
|
|
|
|
|
|
This sort of problem can be encountered on DEC Alpha processors as they have a
|
|
split cache that improves performance by making better use of the data bus.
|
|
While most CPUs do imply a data dependency barrier on the read when a memory
|
|
access depends on a read, not all do, so it may not be relied on.
|
|
|
|
Other CPUs may also have split caches, but must coordinate between the various
|
|
cachelets for normal memory accesses. The semantics of the Alpha removes the
|
|
need for hardware coordination in the absence of memory barriers, which
|
|
permitted Alpha to sport higher CPU clock rates back in the day. However,
|
|
please note that (again, as of v4.15) smp_read_barrier_depends() should not
|
|
be used except in Alpha arch-specific code and within the READ_ONCE() macro.
|
|
|
|
|
|
CACHE COHERENCY VS DMA
|
|
----------------------
|
|
|
|
Not all systems maintain cache coherency with respect to devices doing DMA. In
|
|
such cases, a device attempting DMA may obtain stale data from RAM because
|
|
dirty cache lines may be resident in the caches of various CPUs, and may not
|
|
have been written back to RAM yet. To deal with this, the appropriate part of
|
|
the kernel must flush the overlapping bits of cache on each CPU (and maybe
|
|
invalidate them as well).
|
|
|
|
In addition, the data DMA'd to RAM by a device may be overwritten by dirty
|
|
cache lines being written back to RAM from a CPU's cache after the device has
|
|
installed its own data, or cache lines present in the CPU's cache may simply
|
|
obscure the fact that RAM has been updated, until at such time as the cacheline
|
|
is discarded from the CPU's cache and reloaded. To deal with this, the
|
|
appropriate part of the kernel must invalidate the overlapping bits of the
|
|
cache on each CPU.
|
|
|
|
See Documentation/core-api/cachetlb.rst for more information on cache management.
|
|
|
|
|
|
CACHE COHERENCY VS MMIO
|
|
-----------------------
|
|
|
|
Memory mapped I/O usually takes place through memory locations that are part of
|
|
a window in the CPU's memory space that has different properties assigned than
|
|
the usual RAM directed window.
|
|
|
|
Amongst these properties is usually the fact that such accesses bypass the
|
|
caching entirely and go directly to the device buses. This means MMIO accesses
|
|
may, in effect, overtake accesses to cached memory that were emitted earlier.
|
|
A memory barrier isn't sufficient in such a case, but rather the cache must be
|
|
flushed between the cached memory write and the MMIO access if the two are in
|
|
any way dependent.
|
|
|
|
|
|
=========================
|
|
THE THINGS CPUS GET UP TO
|
|
=========================
|
|
|
|
A programmer might take it for granted that the CPU will perform memory
|
|
operations in exactly the order specified, so that if the CPU is, for example,
|
|
given the following piece of code to execute:
|
|
|
|
a = READ_ONCE(*A);
|
|
WRITE_ONCE(*B, b);
|
|
c = READ_ONCE(*C);
|
|
d = READ_ONCE(*D);
|
|
WRITE_ONCE(*E, e);
|
|
|
|
they would then expect that the CPU will complete the memory operation for each
|
|
instruction before moving on to the next one, leading to a definite sequence of
|
|
operations as seen by external observers in the system:
|
|
|
|
LOAD *A, STORE *B, LOAD *C, LOAD *D, STORE *E.
|
|
|
|
|
|
Reality is, of course, much messier. With many CPUs and compilers, the above
|
|
assumption doesn't hold because:
|
|
|
|
(*) loads are more likely to need to be completed immediately to permit
|
|
execution progress, whereas stores can often be deferred without a
|
|
problem;
|
|
|
|
(*) loads may be done speculatively, and the result discarded should it prove
|
|
to have been unnecessary;
|
|
|
|
(*) loads may be done speculatively, leading to the result having been fetched
|
|
at the wrong time in the expected sequence of events;
|
|
|
|
(*) the order of the memory accesses may be rearranged to promote better use
|
|
of the CPU buses and caches;
|
|
|
|
(*) loads and stores may be combined to improve performance when talking to
|
|
memory or I/O hardware that can do batched accesses of adjacent locations,
|
|
thus cutting down on transaction setup costs (memory and PCI devices may
|
|
both be able to do this); and
|
|
|
|
(*) the CPU's data cache may affect the ordering, and while cache-coherency
|
|
mechanisms may alleviate this - once the store has actually hit the cache
|
|
- there's no guarantee that the coherency management will be propagated in
|
|
order to other CPUs.
|
|
|
|
So what another CPU, say, might actually observe from the above piece of code
|
|
is:
|
|
|
|
LOAD *A, ..., LOAD {*C,*D}, STORE *E, STORE *B
|
|
|
|
(Where "LOAD {*C,*D}" is a combined load)
|
|
|
|
|
|
However, it is guaranteed that a CPU will be self-consistent: it will see its
|
|
_own_ accesses appear to be correctly ordered, without the need for a memory
|
|
barrier. For instance with the following code:
|
|
|
|
U = READ_ONCE(*A);
|
|
WRITE_ONCE(*A, V);
|
|
WRITE_ONCE(*A, W);
|
|
X = READ_ONCE(*A);
|
|
WRITE_ONCE(*A, Y);
|
|
Z = READ_ONCE(*A);
|
|
|
|
and assuming no intervention by an external influence, it can be assumed that
|
|
the final result will appear to be:
|
|
|
|
U == the original value of *A
|
|
X == W
|
|
Z == Y
|
|
*A == Y
|
|
|
|
The code above may cause the CPU to generate the full sequence of memory
|
|
accesses:
|
|
|
|
U=LOAD *A, STORE *A=V, STORE *A=W, X=LOAD *A, STORE *A=Y, Z=LOAD *A
|
|
|
|
in that order, but, without intervention, the sequence may have almost any
|
|
combination of elements combined or discarded, provided the program's view
|
|
of the world remains consistent. Note that READ_ONCE() and WRITE_ONCE()
|
|
are -not- optional in the above example, as there are architectures
|
|
where a given CPU might reorder successive loads to the same location.
|
|
On such architectures, READ_ONCE() and WRITE_ONCE() do whatever is
|
|
necessary to prevent this, for example, on Itanium the volatile casts
|
|
used by READ_ONCE() and WRITE_ONCE() cause GCC to emit the special ld.acq
|
|
and st.rel instructions (respectively) that prevent such reordering.
|
|
|
|
The compiler may also combine, discard or defer elements of the sequence before
|
|
the CPU even sees them.
|
|
|
|
For instance:
|
|
|
|
*A = V;
|
|
*A = W;
|
|
|
|
may be reduced to:
|
|
|
|
*A = W;
|
|
|
|
since, without either a write barrier or an WRITE_ONCE(), it can be
|
|
assumed that the effect of the storage of V to *A is lost. Similarly:
|
|
|
|
*A = Y;
|
|
Z = *A;
|
|
|
|
may, without a memory barrier or an READ_ONCE() and WRITE_ONCE(), be
|
|
reduced to:
|
|
|
|
*A = Y;
|
|
Z = Y;
|
|
|
|
and the LOAD operation never appear outside of the CPU.
|
|
|
|
|
|
AND THEN THERE'S THE ALPHA
|
|
--------------------------
|
|
|
|
The DEC Alpha CPU is one of the most relaxed CPUs there is. Not only that,
|
|
some versions of the Alpha CPU have a split data cache, permitting them to have
|
|
two semantically-related cache lines updated at separate times. This is where
|
|
the data dependency barrier really becomes necessary as this synchronises both
|
|
caches with the memory coherence system, thus making it seem like pointer
|
|
changes vs new data occur in the right order.
|
|
|
|
The Alpha defines the Linux kernel's memory model, although as of v4.15
|
|
the Linux kernel's addition of smp_read_barrier_depends() to READ_ONCE()
|
|
greatly reduced Alpha's impact on the memory model.
|
|
|
|
See the subsection on "Cache Coherency" above.
|
|
|
|
|
|
VIRTUAL MACHINE GUESTS
|
|
----------------------
|
|
|
|
Guests running within virtual machines might be affected by SMP effects even if
|
|
the guest itself is compiled without SMP support. This is an artifact of
|
|
interfacing with an SMP host while running an UP kernel. Using mandatory
|
|
barriers for this use-case would be possible but is often suboptimal.
|
|
|
|
To handle this case optimally, low-level virt_mb() etc macros are available.
|
|
These have the same effect as smp_mb() etc when SMP is enabled, but generate
|
|
identical code for SMP and non-SMP systems. For example, virtual machine guests
|
|
should use virt_mb() rather than smp_mb() when synchronizing against a
|
|
(possibly SMP) host.
|
|
|
|
These are equivalent to smp_mb() etc counterparts in all other respects,
|
|
in particular, they do not control MMIO effects: to control
|
|
MMIO effects, use mandatory barriers.
|
|
|
|
|
|
============
|
|
EXAMPLE USES
|
|
============
|
|
|
|
CIRCULAR BUFFERS
|
|
----------------
|
|
|
|
Memory barriers can be used to implement circular buffering without the need
|
|
of a lock to serialise the producer with the consumer. See:
|
|
|
|
Documentation/core-api/circular-buffers.rst
|
|
|
|
for details.
|
|
|
|
|
|
==========
|
|
REFERENCES
|
|
==========
|
|
|
|
Alpha AXP Architecture Reference Manual, Second Edition (Sites & Witek,
|
|
Digital Press)
|
|
Chapter 5.2: Physical Address Space Characteristics
|
|
Chapter 5.4: Caches and Write Buffers
|
|
Chapter 5.5: Data Sharing
|
|
Chapter 5.6: Read/Write Ordering
|
|
|
|
AMD64 Architecture Programmer's Manual Volume 2: System Programming
|
|
Chapter 7.1: Memory-Access Ordering
|
|
Chapter 7.4: Buffering and Combining Memory Writes
|
|
|
|
ARM Architecture Reference Manual (ARMv8, for ARMv8-A architecture profile)
|
|
Chapter B2: The AArch64 Application Level Memory Model
|
|
|
|
IA-32 Intel Architecture Software Developer's Manual, Volume 3:
|
|
System Programming Guide
|
|
Chapter 7.1: Locked Atomic Operations
|
|
Chapter 7.2: Memory Ordering
|
|
Chapter 7.4: Serializing Instructions
|
|
|
|
The SPARC Architecture Manual, Version 9
|
|
Chapter 8: Memory Models
|
|
Appendix D: Formal Specification of the Memory Models
|
|
Appendix J: Programming with the Memory Models
|
|
|
|
Storage in the PowerPC (Stone and Fitzgerald)
|
|
|
|
UltraSPARC Programmer Reference Manual
|
|
Chapter 5: Memory Accesses and Cacheability
|
|
Chapter 15: Sparc-V9 Memory Models
|
|
|
|
UltraSPARC III Cu User's Manual
|
|
Chapter 9: Memory Models
|
|
|
|
UltraSPARC IIIi Processor User's Manual
|
|
Chapter 8: Memory Models
|
|
|
|
UltraSPARC Architecture 2005
|
|
Chapter 9: Memory
|
|
Appendix D: Formal Specifications of the Memory Models
|
|
|
|
UltraSPARC T1 Supplement to the UltraSPARC Architecture 2005
|
|
Chapter 8: Memory Models
|
|
Appendix F: Caches and Cache Coherency
|
|
|
|
Solaris Internals, Core Kernel Architecture, p63-68:
|
|
Chapter 3.3: Hardware Considerations for Locks and
|
|
Synchronization
|
|
|
|
Unix Systems for Modern Architectures, Symmetric Multiprocessing and Caching
|
|
for Kernel Programmers:
|
|
Chapter 13: Other Memory Models
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Intel Itanium Architecture Software Developer's Manual: Volume 1:
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Section 2.6: Speculation
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Section 4.4: Memory Access
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