On atomic types (atomic_t atomic64_t and atomic_long_t).
  
  The atomic type provides an interface to the architecture's means of atomic
  RMW operations between CPUs (atomic operations on MMIO are not supported and
  can lead to fatal traps on some platforms).
  
  API
  ---
  
  The 'full' API consists of (atomic64_ and atomic_long_ prefixes omitted for
  brevity):
  
  Non-RMW ops:
  
    atomic_read(), atomic_set()
    atomic_read_acquire(), atomic_set_release()
  
  
  RMW atomic operations:
  
  Arithmetic:
  
    atomic_{add,sub,inc,dec}()
    atomic_{add,sub,inc,dec}_return{,_relaxed,_acquire,_release}()
    atomic_fetch_{add,sub,inc,dec}{,_relaxed,_acquire,_release}()
  
  
  Bitwise:
  
    atomic_{and,or,xor,andnot}()
    atomic_fetch_{and,or,xor,andnot}{,_relaxed,_acquire,_release}()
  
  
  Swap:
  
    atomic_xchg{,_relaxed,_acquire,_release}()
    atomic_cmpxchg{,_relaxed,_acquire,_release}()
    atomic_try_cmpxchg{,_relaxed,_acquire,_release}()
  
  
  Reference count (but please see refcount_t):
  
    atomic_add_unless(), atomic_inc_not_zero()
    atomic_sub_and_test(), atomic_dec_and_test()
  
  
  Misc:
  
    atomic_inc_and_test(), atomic_add_negative()
    atomic_dec_unless_positive(), atomic_inc_unless_negative()
  
  
  Barriers:
  
    smp_mb__{before,after}_atomic()

  TYPES (signed vs unsigned)
  -----
  
  While atomic_t, atomic_long_t and atomic64_t use int, long and s64
  respectively (for hysterical raisins), the kernel uses -fno-strict-overflow
  (which implies -fwrapv) and defines signed overflow to behave like
  2s-complement.
  
  Therefore, an explicitly unsigned variant of the atomic ops is strictly
  unnecessary and we can simply cast, there is no UB.
  
  There was a bug in UBSAN prior to GCC-8 that would generate UB warnings for
  signed types.
  
  With this we also conform to the C/C++ _Atomic behaviour and things like
  P1236R1.

  
  SEMANTICS
  ---------
  
  Non-RMW ops:
  
  The non-RMW ops are (typically) regular LOADs and STOREs and are canonically
  implemented using READ_ONCE(), WRITE_ONCE(), smp_load_acquire() and

  smp_store_release() respectively. Therefore, if you find yourself only using
  the Non-RMW operations of atomic_t, you do not in fact need atomic_t at all
  and are doing it wrong.

  A note for the implementation of atomic_set{}() is that it must not break the
  atomicity of the RMW ops. That is:

    C Atomic-RMW-ops-are-atomic-WRT-atomic_set

  
    {

      atomic_t v = ATOMIC_INIT(1);

    }

    P0(atomic_t *v)

    {

      (void)atomic_add_unless(v, 1, 0);

    }

    P1(atomic_t *v)

    {
      atomic_set(v, 0);
    }
  
    exists
    (v=2)
  
  In this case we would expect the atomic_set() from CPU1 to either happen
  before the atomic_add_unless(), in which case that latter one would no-op, or
  _after_ in which case we'd overwrite its result. In no case is "2" a valid
  outcome.
  
  This is typically true on 'normal' platforms, where a regular competing STORE
  will invalidate a LL/SC or fail a CMPXCHG.
  
  The obvious case where this is not so is when we need to implement atomic ops
  with a lock:
  
    CPU0						CPU1
  
    atomic_add_unless(v, 1, 0);
      lock();
      ret = READ_ONCE(v->counter); // == 1
  						atomic_set(v, 0);
      if (ret != u)				  WRITE_ONCE(v->counter, 0);
        WRITE_ONCE(v->counter, ret + 1);
      unlock();
  
  the typical solution is to then implement atomic_set{}() with atomic_xchg().
  
  
  RMW ops:
  
  These come in various forms:
  
   - plain operations without return value: atomic_{}()
  
   - operations which return the modified value: atomic_{}_return()
  
     these are limited to the arithmetic operations because those are
     reversible. Bitops are irreversible and therefore the modified value
     is of dubious utility.
  
   - operations which return the original value: atomic_fetch_{}()
  
   - swap operations: xchg(), cmpxchg() and try_cmpxchg()
  
   - misc; the special purpose operations that are commonly used and would,
     given the interface, normally be implemented using (try_)cmpxchg loops but
     are time critical and can, (typically) on LL/SC architectures, be more
     efficiently implemented.
  
  All these operations are SMP atomic; that is, the operations (for a single
  atomic variable) can be fully ordered and no intermediate state is lost or
  visible.
  
  
  ORDERING  (go read memory-barriers.txt first)
  --------
  
  The rule of thumb:
  
   - non-RMW operations are unordered;
  
   - RMW operations that have no return value are unordered;
  
   - RMW operations that have a return value are fully ordered;
  
   - RMW operations that are conditional are unordered on FAILURE,
     otherwise the above rules apply.
  
  Except of course when an operation has an explicit ordering like:
  
   {}_relaxed: unordered
   {}_acquire: the R of the RMW (or atomic_read) is an ACQUIRE
   {}_release: the W of the RMW (or atomic_set)  is a  RELEASE
  
  Where 'unordered' is against other memory locations. Address dependencies are
  not defeated.
  
  Fully ordered primitives are ordered against everything prior and everything
  subsequent. Therefore a fully ordered primitive is like having an smp_mb()
  before and an smp_mb() after the primitive.
  
  
  The barriers:
  
    smp_mb__{before,after}_atomic()

  only apply to the RMW atomic ops and can be used to augment/upgrade the
  ordering inherent to the op. These barriers act almost like a full smp_mb():
  smp_mb__before_atomic() orders all earlier accesses against the RMW op
  itself and all accesses following it, and smp_mb__after_atomic() orders all
  later accesses against the RMW op and all accesses preceding it. However,
  accesses between the smp_mb__{before,after}_atomic() and the RMW op are not
  ordered, so it is advisable to place the barrier right next to the RMW atomic
  op whenever possible.

  
  These helper barriers exist because architectures have varying implicit
  ordering on their SMP atomic primitives. For example our TSO architectures
  provide full ordered atomics and these barriers are no-ops.

  NOTE: when the atomic RmW ops are fully ordered, they should also imply a
  compiler barrier.

  Thus:
  
    atomic_fetch_add();
  
  is equivalent to:
  
    smp_mb__before_atomic();
    atomic_fetch_add_relaxed();
    smp_mb__after_atomic();
  
  However the atomic_fetch_add() might be implemented more efficiently.
  
  Further, while something like:
  
    smp_mb__before_atomic();
    atomic_dec(&X);
  
  is a 'typical' RELEASE pattern, the barrier is strictly stronger than

  a RELEASE because it orders preceding instructions against both the read
  and write parts of the atomic_dec(), and against all following instructions
  as well. Similarly, something like:

    atomic_inc(&X);
    smp_mb__after_atomic();
  
  is an ACQUIRE pattern (though very much not typical), but again the barrier is
  strictly stronger than ACQUIRE. As illustrated:

    C Atomic-RMW+mb__after_atomic-is-stronger-than-acquire

  
    {
    }

    P0(int *x, atomic_t *y)

    {
      r0 = READ_ONCE(*x);
      smp_rmb();
      r1 = atomic_read(y);
    }

    P1(int *x, atomic_t *y)

    {
      atomic_inc(y);
      smp_mb__after_atomic();
      WRITE_ONCE(*x, 1);
    }
  
    exists

    (0:r0=1 /\ 0:r1=0)

  
  This should not happen; but a hypothetical atomic_inc_acquire() --
  (void)atomic_fetch_inc_acquire() for instance -- would allow the outcome,

  because it would not order the W part of the RMW against the following
  WRITE_ONCE.  Thus:

    P0			P1

  
  			t = LL.acq *y (0)
  			t++;
  			*x = 1;
    r0 = *x (1)
    RMB
    r1 = *y (0)
  			SC *y, t;

  is allowed.

  
  
  CMPXCHG vs TRY_CMPXCHG
  ----------------------
  
    int atomic_cmpxchg(atomic_t *ptr, int old, int new);
    bool atomic_try_cmpxchg(atomic_t *ptr, int *oldp, int new);
  
  Both provide the same functionality, but try_cmpxchg() can lead to more
  compact code. The functions relate like:
  
    bool atomic_try_cmpxchg(atomic_t *ptr, int *oldp, int new)
    {
      int ret, old = *oldp;
      ret = atomic_cmpxchg(ptr, old, new);
      if (ret != old)
        *oldp = ret;
      return ret == old;
    }
  
  and:
  
    int atomic_cmpxchg(atomic_t *ptr, int old, int new)
    {
      (void)atomic_try_cmpxchg(ptr, &old, new);
      return old;
    }
  
  Usage:
  
    old = atomic_read(&v);			old = atomic_read(&v);
    for (;;) {					do {
      new = func(old);				  new = func(old);
      tmp = atomic_cmpxchg(&v, old, new);		} while (!atomic_try_cmpxchg(&v, &old, new));
      if (tmp == old)
        break;
      old = tmp;
    }
  
  NB. try_cmpxchg() also generates better code on some platforms (notably x86)
  where the function more closely matches the hardware instruction.

  
  
  FORWARD PROGRESS
  ----------------
  
  In general strong forward progress is expected of all unconditional atomic
  operations -- those in the Arithmetic and Bitwise classes and xchg(). However
  a fair amount of code also requires forward progress from the conditional
  atomic operations.
  
  Specifically 'simple' cmpxchg() loops are expected to not starve one another
  indefinitely. However, this is not evident on LL/SC architectures, because
  while an LL/SC architecure 'can/should/must' provide forward progress
  guarantees between competing LL/SC sections, such a guarantee does not
  transfer to cmpxchg() implemented using LL/SC. Consider:
  
    old = atomic_read(&v);
    do {
      new = func(old);
    } while (!atomic_try_cmpxchg(&v, &old, new));
  
  which on LL/SC becomes something like:
  
    old = atomic_read(&v);
    do {
      new = func(old);
    } while (!({
      volatile asm ("1: LL  %[oldval], %[v]
  "
                    "   CMP %[oldval], %[old]
  "
                    "   BNE 2f
  "
                    "   SC  %[new], %[v]
  "
                    "   BNE 1b
  "
                    "2:
  "
                    : [oldval] "=&r" (oldval), [v] "m" (v)
  		  : [old] "r" (old), [new] "r" (new)
                    : "memory");
      success = (oldval == old);
      if (!success)
        old = oldval;
      success; }));
  
  However, even the forward branch from the failed compare can cause the LL/SC
  to fail on some architectures, let alone whatever the compiler makes of the C
  loop body. As a result there is no guarantee what so ever the cacheline
  containing @v will stay on the local CPU and progress is made.
  
  Even native CAS architectures can fail to provide forward progress for their
  primitive (See Sparc64 for an example).
  
  Such implementations are strongly encouraged to add exponential backoff loops
  to a failed CAS in order to ensure some progress. Affected architectures are
  also strongly encouraged to inspect/audit the atomic fallbacks, refcount_t and
  their locking primitives.