diff options
Diffstat (limited to 'Documentation')
21 files changed, 2183 insertions, 168 deletions
diff --git a/Documentation/DocBook/Makefile b/Documentation/DocBook/Makefile index 7d87dd7..5a2882d 100644 --- a/Documentation/DocBook/Makefile +++ b/Documentation/DocBook/Makefile @@ -2,7 +2,7 @@ # This makefile is used to generate the kernel documentation, # primarily based on in-line comments in various source files. # See Documentation/kernel-doc-nano-HOWTO.txt for instruction in how -# to ducument the SRC - and how to read it. +# to document the SRC - and how to read it. # To add a new book the only step required is to add the book to the # list of DOCBOOKS. diff --git a/Documentation/DocBook/kernel-api.tmpl b/Documentation/DocBook/kernel-api.tmpl index 8c9c670..ca02e04 100644 --- a/Documentation/DocBook/kernel-api.tmpl +++ b/Documentation/DocBook/kernel-api.tmpl @@ -322,7 +322,6 @@ X!Earch/i386/kernel/mca.c <chapter id="sysfs"> <title>The Filesystem for Exporting Kernel Objects</title> !Efs/sysfs/file.c -!Efs/sysfs/dir.c !Efs/sysfs/symlink.c !Efs/sysfs/bin.c </chapter> diff --git a/Documentation/acpi-hotkey.txt b/Documentation/acpi-hotkey.txt index 744f1ae..38040fa 100644 --- a/Documentation/acpi-hotkey.txt +++ b/Documentation/acpi-hotkey.txt @@ -30,7 +30,7 @@ specific hotkey(event)) echo "event_num:event_type:event_argument" > /proc/acpi/hotkey/action. The result of the execution of this aml method is -attached to /proc/acpi/hotkey/poll_method, which is dnyamically +attached to /proc/acpi/hotkey/poll_method, which is dynamically created. Please use command "cat /proc/acpi/hotkey/polling_method" to retrieve it. diff --git a/Documentation/feature-removal-schedule.txt b/Documentation/feature-removal-schedule.txt index 495858b..59d0c74 100644 --- a/Documentation/feature-removal-schedule.txt +++ b/Documentation/feature-removal-schedule.txt @@ -127,13 +127,6 @@ Who: Christoph Hellwig <hch@lst.de> --------------------------- -What: EXPORT_SYMBOL(lookup_hash) -When: January 2006 -Why: Too low-level interface. Use lookup_one_len or lookup_create instead. -Who: Christoph Hellwig <hch@lst.de> - ---------------------------- - What: CONFIG_FORCED_INLINING When: June 2006 Why: Config option is there to see if gcc is good enough. (in january @@ -241,3 +234,15 @@ Why: The USB subsystem has changed a lot over time, and it has been Who: Greg Kroah-Hartman <gregkh@suse.de> --------------------------- + +What: find_trylock_page +When: January 2007 +Why: The interface no longer has any callers left in the kernel. It + is an odd interface (compared with other find_*_page functions), in + that it does not take a refcount to the page, only the page lock. + It should be replaced with find_get_page or find_lock_page if possible. + This feature removal can be reevaluated if users of the interface + cannot cleanly use something else. +Who: Nick Piggin <npiggin@suse.de> + +--------------------------- diff --git a/Documentation/fujitsu/frv/kernel-ABI.txt b/Documentation/fujitsu/frv/kernel-ABI.txt index 0ed9b0a..8b0a5fc 100644 --- a/Documentation/fujitsu/frv/kernel-ABI.txt +++ b/Documentation/fujitsu/frv/kernel-ABI.txt @@ -1,17 +1,19 @@ - ================================= - INTERNAL KERNEL ABI FOR FR-V ARCH - ================================= - -The internal FRV kernel ABI is not quite the same as the userspace ABI. A number of the registers -are used for special purposed, and the ABI is not consistent between modules vs core, and MMU vs -no-MMU. - -This partly stems from the fact that FRV CPUs do not have a separate supervisor stack pointer, and -most of them do not have any scratch registers, thus requiring at least one general purpose -register to be clobbered in such an event. Also, within the kernel core, it is possible to simply -jump or call directly between functions using a relative offset. This cannot be extended to modules -for the displacement is likely to be too far. Thus in modules the address of a function to call -must be calculated in a register and then used, requiring two extra instructions. + ================================= + INTERNAL KERNEL ABI FOR FR-V ARCH + ================================= + +The internal FRV kernel ABI is not quite the same as the userspace ABI. A +number of the registers are used for special purposed, and the ABI is not +consistent between modules vs core, and MMU vs no-MMU. + +This partly stems from the fact that FRV CPUs do not have a separate +supervisor stack pointer, and most of them do not have any scratch +registers, thus requiring at least one general purpose register to be +clobbered in such an event. Also, within the kernel core, it is possible to +simply jump or call directly between functions using a relative offset. +This cannot be extended to modules for the displacement is likely to be too +far. Thus in modules the address of a function to call must be calculated +in a register and then used, requiring two extra instructions. This document has the following sections: @@ -39,7 +41,8 @@ When a system call is made, the following registers are effective: CPU OPERATING MODES =================== -The FR-V CPU has three basic operating modes. In order of increasing capability: +The FR-V CPU has three basic operating modes. In order of increasing +capability: (1) User mode. @@ -47,42 +50,46 @@ The FR-V CPU has three basic operating modes. In order of increasing capability: (2) Kernel mode. - Normal kernel mode. There are many additional control registers available that may be - accessed in this mode, in addition to all the stuff available to user mode. This has two - submodes: + Normal kernel mode. There are many additional control registers + available that may be accessed in this mode, in addition to all the + stuff available to user mode. This has two submodes: (a) Exceptions enabled (PSR.T == 1). - Exceptions will invoke the appropriate normal kernel mode handler. On entry to the - handler, the PSR.T bit will be cleared. + Exceptions will invoke the appropriate normal kernel mode + handler. On entry to the handler, the PSR.T bit will be cleared. (b) Exceptions disabled (PSR.T == 0). - No exceptions or interrupts may happen. Any mandatory exceptions will cause the CPU to - halt unless the CPU is told to jump into debug mode instead. + No exceptions or interrupts may happen. Any mandatory exceptions + will cause the CPU to halt unless the CPU is told to jump into + debug mode instead. (3) Debug mode. - No exceptions may happen in this mode. Memory protection and management exceptions will be - flagged for later consideration, but the exception handler won't be invoked. Debugging traps - such as hardware breakpoints and watchpoints will be ignored. This mode is entered only by - debugging events obtained from the other two modes. + No exceptions may happen in this mode. Memory protection and + management exceptions will be flagged for later consideration, but + the exception handler won't be invoked. Debugging traps such as + hardware breakpoints and watchpoints will be ignored. This mode is + entered only by debugging events obtained from the other two modes. - All kernel mode registers may be accessed, plus a few extra debugging specific registers. + All kernel mode registers may be accessed, plus a few extra debugging + specific registers. ================================= INTERNAL KERNEL-MODE REGISTER ABI ================================= -There are a number of permanent register assignments that are set up by entry.S in the exception -prologue. Note that there is a complete set of exception prologues for each of user->kernel -transition and kernel->kernel transition. There are also user->debug and kernel->debug mode -transition prologues. +There are a number of permanent register assignments that are set up by +entry.S in the exception prologue. Note that there is a complete set of +exception prologues for each of user->kernel transition and kernel->kernel +transition. There are also user->debug and kernel->debug mode transition +prologues. REGISTER FLAVOUR USE - =============== ======= ==================================================== + =============== ======= ============================================== GR1 Supervisor stack pointer GR15 Current thread info pointer GR16 GP-Rel base register for small data @@ -92,10 +99,12 @@ transition prologues. GR31 NOMMU Destroyed by debug mode entry GR31 MMU Destroyed by TLB miss kernel mode entry CCR.ICC2 Virtual interrupt disablement tracking - CCCR.CC3 Cleared by exception prologue (atomic op emulation) + CCCR.CC3 Cleared by exception prologue + (atomic op emulation) SCR0 MMU See mmu-layout.txt. SCR1 MMU See mmu-layout.txt. - SCR2 MMU Save for EAR0 (destroyed by icache insns in debug mode) + SCR2 MMU Save for EAR0 (destroyed by icache insns + in debug mode) SCR3 MMU Save for GR31 during debug exceptions DAMR/IAMR NOMMU Fixed memory protection layout. DAMR/IAMR MMU See mmu-layout.txt. @@ -104,18 +113,21 @@ transition prologues. Certain registers are also used or modified across function calls: REGISTER CALL RETURN - =============== =============================== =============================== + =============== =============================== ====================== GR0 Fixed Zero - GR2 Function call frame pointer GR3 Special Preserved GR3-GR7 - Clobbered - GR8 Function call arg #1 Return value (or clobbered) - GR9 Function call arg #2 Return value MSW (or clobbered) + GR8 Function call arg #1 Return value + (or clobbered) + GR9 Function call arg #2 Return value MSW + (or clobbered) GR10-GR13 Function call arg #3-#6 Clobbered GR14 - Clobbered GR15-GR16 Special Preserved GR17-GR27 - Preserved - GR28-GR31 Special Only accessed explicitly + GR28-GR31 Special Only accessed + explicitly LR Return address after CALL Clobbered CCR/CCCR - Mostly Clobbered @@ -124,46 +136,53 @@ Certain registers are also used or modified across function calls: INTERNAL DEBUG-MODE REGISTER ABI ================================ -This is the same as the kernel-mode register ABI for functions calls. The difference is that in -debug-mode there's a different stack and a different exception frame. Almost all the global -registers from kernel-mode (including the stack pointer) may be changed. +This is the same as the kernel-mode register ABI for functions calls. The +difference is that in debug-mode there's a different stack and a different +exception frame. Almost all the global registers from kernel-mode +(including the stack pointer) may be changed. REGISTER FLAVOUR USE - =============== ======= ==================================================== + =============== ======= ============================================== GR1 Debug stack pointer GR16 GP-Rel base register for small data - GR31 Current debug exception frame pointer (__debug_frame) + GR31 Current debug exception frame pointer + (__debug_frame) SCR3 MMU Saved value of GR31 -Note that debug mode is able to interfere with the kernel's emulated atomic ops, so it must be -exceedingly careful not to do any that would interact with the main kernel in this regard. Hence -the debug mode code (gdbstub) is almost completely self-contained. The only external code used is -the sprintf family of functions. +Note that debug mode is able to interfere with the kernel's emulated atomic +ops, so it must be exceedingly careful not to do any that would interact +with the main kernel in this regard. Hence the debug mode code (gdbstub) is +almost completely self-contained. The only external code used is the +sprintf family of functions. -Futhermore, break.S is so complicated because single-step mode does not switch off on entry to an -exception. That means unless manually disabled, single-stepping will blithely go on stepping into -things like interrupts. See gdbstub.txt for more information. +Futhermore, break.S is so complicated because single-step mode does not +switch off on entry to an exception. That means unless manually disabled, +single-stepping will blithely go on stepping into things like interrupts. +See gdbstub.txt for more information. ========================== VIRTUAL INTERRUPT HANDLING ========================== -Because accesses to the PSR is so slow, and to disable interrupts we have to access it twice (once -to read and once to write), we don't actually disable interrupts at all if we don't have to. What -we do instead is use the ICC2 condition code flags to note virtual disablement, such that if we -then do take an interrupt, we note the flag, really disable interrupts, set another flag and resume -execution at the point the interrupt happened. Setting condition flags as a side effect of an -arithmetic or logical instruction is really fast. This use of the ICC2 only occurs within the +Because accesses to the PSR is so slow, and to disable interrupts we have +to access it twice (once to read and once to write), we don't actually +disable interrupts at all if we don't have to. What we do instead is use +the ICC2 condition code flags to note virtual disablement, such that if we +then do take an interrupt, we note the flag, really disable interrupts, set +another flag and resume execution at the point the interrupt happened. +Setting condition flags as a side effect of an arithmetic or logical +instruction is really fast. This use of the ICC2 only occurs within the kernel - it does not affect userspace. The flags we use are: (*) CCR.ICC2.Z [Zero flag] - Set to virtually disable interrupts, clear when interrupts are virtually enabled. Can be - modified by logical instructions without affecting the Carry flag. + Set to virtually disable interrupts, clear when interrupts are + virtually enabled. Can be modified by logical instructions without + affecting the Carry flag. (*) CCR.ICC2.C [Carry flag] @@ -176,8 +195,9 @@ What happens is this: ICC2.Z is 0, ICC2.C is 1. - (2) An interrupt occurs. The exception prologue examines ICC2.Z and determines that nothing needs - doing. This is done simply with an unlikely BEQ instruction. + (2) An interrupt occurs. The exception prologue examines ICC2.Z and + determines that nothing needs doing. This is done simply with an + unlikely BEQ instruction. (3) The interrupts are disabled (local_irq_disable) @@ -187,48 +207,56 @@ What happens is this: ICC2.Z would be set to 0. - A TIHI #2 instruction (trap #2 if condition HI - Z==0 && C==0) would be used to trap if - interrupts were now virtually enabled, but physically disabled - which they're not, so the - trap isn't taken. The kernel would then be back to state (1). + A TIHI #2 instruction (trap #2 if condition HI - Z==0 && C==0) would + be used to trap if interrupts were now virtually enabled, but + physically disabled - which they're not, so the trap isn't taken. The + kernel would then be back to state (1). - (5) An interrupt occurs. The exception prologue examines ICC2.Z and determines that the interrupt - shouldn't actually have happened. It jumps aside, and there disabled interrupts by setting - PSR.PIL to 14 and then it clears ICC2.C. + (5) An interrupt occurs. The exception prologue examines ICC2.Z and + determines that the interrupt shouldn't actually have happened. It + jumps aside, and there disabled interrupts by setting PSR.PIL to 14 + and then it clears ICC2.C. (6) If interrupts were then saved and disabled again (local_irq_save): - ICC2.Z would be shifted into the save variable and masked off (giving a 1). + ICC2.Z would be shifted into the save variable and masked off + (giving a 1). - ICC2.Z would then be set to 1 (thus unchanged), and ICC2.C would be unaffected (ie: 0). + ICC2.Z would then be set to 1 (thus unchanged), and ICC2.C would be + unaffected (ie: 0). (7) If interrupts were then restored from state (6) (local_irq_restore): - ICC2.Z would be set to indicate the result of XOR'ing the saved value (ie: 1) with 1, which - gives a result of 0 - thus leaving ICC2.Z set. + ICC2.Z would be set to indicate the result of XOR'ing the saved + value (ie: 1) with 1, which gives a result of 0 - thus leaving + ICC2.Z set. ICC2.C would remain unaffected (ie: 0). - A TIHI #2 instruction would be used to again assay the current state, but this would do - nothing as Z==1. + A TIHI #2 instruction would be used to again assay the current state, + but this would do nothing as Z==1. (8) If interrupts were then enabled (local_irq_enable): - ICC2.Z would be cleared. ICC2.C would be left unaffected. Both flags would now be 0. + ICC2.Z would be cleared. ICC2.C would be left unaffected. Both + flags would now be 0. - A TIHI #2 instruction again issued to assay the current state would then trap as both Z==0 - [interrupts virtually enabled] and C==0 [interrupts really disabled] would then be true. + A TIHI #2 instruction again issued to assay the current state would + then trap as both Z==0 [interrupts virtually enabled] and C==0 + [interrupts really disabled] would then be true. - (9) The trap #2 handler would simply enable hardware interrupts (set PSR.PIL to 0), set ICC2.C to - 1 and return. + (9) The trap #2 handler would simply enable hardware interrupts + (set PSR.PIL to 0), set ICC2.C to 1 and return. (10) Immediately upon returning, the pending interrupt would be taken. -(11) The interrupt handler would take the path of actually processing the interrupt (ICC2.Z is - clear, BEQ fails as per step (2)). +(11) The interrupt handler would take the path of actually processing the + interrupt (ICC2.Z is clear, BEQ fails as per step (2)). -(12) The interrupt handler would then set ICC2.C to 1 since hardware interrupts are definitely - enabled - or else the kernel wouldn't be here. +(12) The interrupt handler would then set ICC2.C to 1 since hardware + interrupts are definitely enabled - or else the kernel wouldn't be here. (13) On return from the interrupt handler, things would be back to state (1). -This trap (#2) is only available in kernel mode. In user mode it will result in SIGILL. +This trap (#2) is only available in kernel mode. In user mode it will +result in SIGILL. diff --git a/Documentation/input/joystick-parport.txt b/Documentation/input/joystick-parport.txt index 88a011c..d537c48 100644 --- a/Documentation/input/joystick-parport.txt +++ b/Documentation/input/joystick-parport.txt @@ -36,12 +36,12 @@ with them. All NES and SNES use the same synchronous serial protocol, clocked from the computer's side (and thus timing insensitive). To allow up to 5 NES -and/or SNES gamepads connected to the parallel port at once, the output -lines of the parallel port are shared, while one of 5 available input lines -is assigned to each gamepad. +and/or SNES gamepads and/or SNES mice connected to the parallel port at once, +the output lines of the parallel port are shared, while one of 5 available +input lines is assigned to each gamepad. This protocol is handled by the gamecon.c driver, so that's the one -you'll use for NES and SNES gamepads. +you'll use for NES, SNES gamepads and SNES mice. The main problem with PC parallel ports is that they don't have +5V power source on any of their pins. So, if you want a reliable source of power @@ -106,7 +106,7 @@ A, Turbo B, Select and Start, and is connected through 5 wires, then it is either a NES or NES clone and will work with this connection. SNES gamepads also use 5 wires, but have more buttons. They will work as well, of course. -Pinout for NES gamepads Pinout for SNES gamepads +Pinout for NES gamepads Pinout for SNES gamepads and mice +----> Power +-----------------------\ | 7 | o o o o | x x o | 1 @@ -454,6 +454,7 @@ uses the following kernel/module command line: 6 | N64 pad 7 | Sony PSX controller 8 | Sony PSX DDR controller + 9 | SNES mouse The exact type of the PSX controller type is autoprobed when used so hot swapping should work (but is not recomended). diff --git a/Documentation/kernel-parameters.txt b/Documentation/kernel-parameters.txt index f8cb55c..b3a6187 100644 --- a/Documentation/kernel-parameters.txt +++ b/Documentation/kernel-parameters.txt @@ -1,4 +1,4 @@ -February 2003 Kernel Parameters v2.5.59 + Kernel Parameters ~~~~~~~~~~~~~~~~~ The following is a consolidated list of the kernel parameters as implemented @@ -17,9 +17,17 @@ are specified on the kernel command line with the module name plus usbcore.blinkenlights=1 -The text in square brackets at the beginning of the description states the -restrictions on the kernel for the said kernel parameter to be valid. The -restrictions referred to are that the relevant option is valid if: +This document may not be entirely up to date and comprehensive. The command +"modinfo -p ${modulename}" shows a current list of all parameters of a loadable +module. Loadable modules, after being loaded into the running kernel, also +reveal their parameters in /sys/module/${modulename}/parameters/. Some of these +parameters may be changed at runtime by the command +"echo -n ${value} > /sys/module/${modulename}/parameters/${parm}". + +The parameters listed below are only valid if certain kernel build options were +enabled and if respective hardware is present. The text in square brackets at +the beginning of each description states the restrictions within which a +parameter is applicable: ACPI ACPI support is enabled. ALSA ALSA sound support is enabled. @@ -1046,10 +1054,10 @@ running once the system is up. noltlbs [PPC] Do not use large page/tlb entries for kernel lowmem mapping on PPC40x. - nomce [IA-32] Machine Check Exception - nomca [IA-64] Disable machine check abort handling + nomce [IA-32] Machine Check Exception + noresidual [PPC] Don't use residual data on PReP machines. noresume [SWSUSP] Disables resume and restores original swap @@ -1682,20 +1690,6 @@ running once the system is up. ______________________________________________________________________ -Changelog: - -2000-06-?? Mr. Unknown - The last known update (for 2.4.0) - the changelog was not kept before. - -2002-11-24 Petr Baudis <pasky@ucw.cz> - Randy Dunlap <randy.dunlap@verizon.net> - Update for 2.5.49, description for most of the options introduced, - references to other documentation (C files, READMEs, ..), added S390, - PPC, SPARC, MTD, ALSA and OSS category. Minor corrections and - reformatting. - -2005-10-19 Randy Dunlap <rdunlap@xenotime.net> - Lots of typos, whitespace, some reformatting. TODO: diff --git a/Documentation/leds-class.txt b/Documentation/leds-class.txt new file mode 100644 index 0000000..8c35c04 --- /dev/null +++ b/Documentation/leds-class.txt @@ -0,0 +1,71 @@ +LED handling under Linux +======================== + +If you're reading this and thinking about keyboard leds, these are +handled by the input subsystem and the led class is *not* needed. + +In its simplest form, the LED class just allows control of LEDs from +userspace. LEDs appear in /sys/class/leds/. The brightness file will +set the brightness of the LED (taking a value 0-255). Most LEDs don't +have hardware brightness support so will just be turned on for non-zero +brightness settings. + +The class also introduces the optional concept of an LED trigger. A trigger +is a kernel based source of led events. Triggers can either be simple or +complex. A simple trigger isn't configurable and is designed to slot into +existing subsystems with minimal additional code. Examples are the ide-disk, +nand-disk and sharpsl-charge triggers. With led triggers disabled, the code +optimises away. + +Complex triggers whilst available to all LEDs have LED specific +parameters and work on a per LED basis. The timer trigger is an example. + +You can change triggers in a similar manner to the way an IO scheduler +is chosen (via /sys/class/leds/<device>/trigger). Trigger specific +parameters can appear in /sys/class/leds/<device> once a given trigger is +selected. + + +Design Philosophy +================= + +The underlying design philosophy is simplicity. LEDs are simple devices +and the aim is to keep a small amount of code giving as much functionality +as possible. Please keep this in mind when suggesting enhancements. + + +LED Device Naming +================= + +Is currently of the form: + +"devicename:colour" + +There have been calls for LED properties such as colour to be exported as +individual led class attributes. As a solution which doesn't incur as much +overhead, I suggest these become part of the device name. The naming scheme +above leaves scope for further attributes should they be needed. + + +Known Issues +============ + +The LED Trigger core cannot be a module as the simple trigger functions +would cause nightmare dependency issues. I see this as a minor issue +compared to the benefits the simple trigger functionality brings. The +rest of the LED subsystem can be modular. + +Some leds can be programmed to flash in hardware. As this isn't a generic +LED device property, this should be exported as a device specific sysfs +attribute rather than part of the class if this functionality is required. + + +Future Development +================== + +At the moment, a trigger can't be created specifically for a single LED. +There are a number of cases where a trigger might only be mappable to a +particular LED (ACPI?). The addition of triggers provided by the LED driver +should cover this option and be possible to add without breaking the +current interface. + diff --git a/Documentation/memory-barriers.txt b/Documentation/memory-barriers.txt new file mode 100644 index 0000000..f855031 --- /dev/null +++ b/Documentation/memory-barriers.txt @@ -0,0 +1,1913 @@ + ============================ + LINUX KERNEL MEMORY BARRIERS + ============================ + +By: David Howells <dhowells@redhat.com> + +Contents: + + (*) Abstract memory access model. + + - Device operations. + - Guarantees. + + (*) What are memory barriers? + + - Varieties of memory barrier. + - What may not be assumed about memory barriers? + - Data dependency barriers. + - Control dependencies. + - SMP barrier pairing. + - Examples of memory barrier sequences. + + (*) Explicit kernel barriers. + + - Compiler barrier. + - The CPU memory barriers. + - MMIO write barrier. + + (*) Implicit kernel memory barriers. + + - Locking functions. + - Interrupt disabling functions. + - Miscellaneous functions. + + (*) Inter-CPU locking barrier effects. + + - Locks vs memory accesses. + - Locks vs I/O accesses. + + (*) Where are memory barriers needed? + + - Interprocessor interaction. + - Atomic operations. + - Accessing devices. + - Interrupts. + + (*) Kernel I/O barrier effects. + + (*) Assumed minimum execution ordering model. + + (*) The effects of the cpu cache. + + - Cache coherency. + - Cache coherency vs DMA. + - Cache coherency vs MMIO. + + (*) The things CPUs get up to. + + - And then there's the Alpha. + + (*) References. + + +============================ +ABSTRACT MEMORY ACCESS MODEL +============================ + +Consider the following abstract model of the system: + + : : + : : + : : + +-------+ : +--------+ : +-------+ + | | : | | : | | + | | : | | : | | + | CPU 1 |<----->| Memory |<----->| CPU 2 | + | | : | | : | | + | | : | | : | | + +-------+ : +--------+ : +-------+ + ^ : ^ : ^ + | : | : | + | : | : | + | : v : | + | : +--------+ : | + | : | | : | + | : | | : | + +---------->| Device |<----------+ + : | | : + : | | : + : +--------+ : + : : + +Each CPU executes a program that generates memory access operations. In the +abstract CPU, memory operation ordering is very relaxed, and a CPU may actually +perform the memory operations in any order it likes, provided program causality +appears to be maintained. Similarly, the compiler may also arrange the +instructions it emits in any order it likes, provided it doesn't affect the +apparent operation of the program. + +So in the above diagram, the effects of the memory operations performed by a +CPU are perceived by the rest of the system as the operations cross the +interface between the CPU and rest of the system (the dotted lines). + + +For example, consider the following sequence of events: + + CPU 1 CPU 2 + =============== =============== + { A == 1; B == 2 } + A = 3; x = A; + B = 4; y = B; + +The set of accesses as seen by the memory system in the middle can be arranged +in 24 different combinations: + + STORE A=3, STORE B=4, x=LOAD A->3, y=LOAD B->4 + STORE A=3, STORE B=4, y=LOAD B->4, x=LOAD A->3 + STORE A=3, x=LOAD A->3, STORE B=4, y=LOAD B->4 + STORE A=3, x=LOAD A->3, y=LOAD B->2, STORE B=4 + STORE A=3, y=LOAD B->2, STORE B=4, x=LOAD A->3 + STORE A=3, y=LOAD B->2, x=LOAD A->3, STORE B=4 + STORE B=4, STORE A=3, x=LOAD A->3, y=LOAD B->4 + STORE B=4, ... + ... + +and can thus result in four different combinations of values: + + x == 1, y == 2 + x == 1, y == 4 + x == 3, y == 2 + x == 3, y == 4 + + +Furthermore, the stores committed by a CPU to the memory system may not be +perceived by the loads made by another CPU in the same order as the stores were +committed. + + +As a further example, consider this sequence of events: + + CPU 1 CPU 2 + =============== =============== + { A == 1, B == 2, C = 3, P == &A, Q == &C } + B = 4; Q = P; + P = &B D = *Q; + +There is an obvious data dependency here, as the value loaded into D depends on +the address retrieved from P by CPU 2. At the end of the sequence, any of the +following results are possible: + + (Q == &A) and (D == 1) + (Q == &B) and (D == 2) + (Q == &B) and (D == 4) + +Note that CPU 2 will never try and load C into D because the CPU will load P +into Q before issuing the load of *Q. + + +DEVICE OPERATIONS +----------------- + +Some devices present their control interfaces as collections of memory +locations, but the order in which the control registers are accessed is very +important. For instance, imagine an ethernet card with a set of internal +registers that are accessed through an address port register (A) and a data +port register (D). To read internal register 5, the following code might then +be used: + + *A = 5; + x = *D; + +but this might show up as either of the following two sequences: + + STORE *A = 5, x = LOAD *D + x = LOAD *D, STORE *A = 5 + +the second of which will almost certainly result in a malfunction, since it set +the address _after_ attempting to read the register. + + +GUARANTEES +---------- + +There are some minimal guarantees that may be expected of a CPU: + + (*) On any given CPU, dependent memory accesses will be issued in order, with + respect to itself. This means that for: + + Q = P; D = *Q; + + the CPU will issue the following memory operations: + + Q = LOAD P, D = LOAD *Q + + and always in that order. + + (*) Overlapping loads and stores within a particular CPU will appear to be + ordered within that CPU. This means that for: + + a = *X; *X = b; + + the CPU will only issue the following sequence of memory operations: + + a = LOAD *X, STORE *X = b + + And for: + + *X = c; d = *X; + + the CPU will only issue: + + STORE *X = c, d = LOAD *X + + (Loads and stores overlap if they are targetted at overlapping pieces of + memory). + +And there are a number of things that _must_ or _must_not_ be assumed: + + (*) It _must_not_ be assumed that independent loads and stores will be issued + in the order given. This means that for: + + X = *A; Y = *B; *D = Z; + + we may get any of the following sequences: + + X = LOAD *A, Y = LOAD *B, STORE *D = Z + X = LOAD *A, STORE *D = Z, Y = LOAD *B + Y = LOAD *B, X = LOAD *A, STORE *D = Z + Y = LOAD *B, STORE *D = Z, X = LOAD *A + STORE *D = Z, X = LOAD *A, Y = LOAD *B + STORE *D = Z, Y = LOAD *B, X = LOAD *A + + (*) It _must_ be assumed that overlapping memory accesses may be merged or + discarded. This means that for: + + X = *A; Y = *(A + 4); + + we may get any one of the following sequences: + + X = LOAD *A; Y = LOAD *(A + 4); + Y = LOAD *(A + 4); X = LOAD *A; + {X, Y} = LOAD {*A, *(A + 4) }; + + And for: + + *A = X; Y = *A; + + we may get either of: + + STORE *A = X; Y = LOAD *A; + STORE *A = Y; + + +========================= +WHAT ARE MEMORY BARRIERS? +========================= + +As can be seen above, independent memory operations are effectively performed +in random order, but this can be a problem for CPU-CPU interaction and for I/O. +What is required is some way of intervening to instruct the compiler and the +CPU to restrict the order. + +Memory barriers are such interventions. They impose a perceived partial +ordering between the memory operations specified on either side of the barrier. +They request that the sequence of memory events generated appears to other +parts of the system as if the barrier is effective on that CPU. + + +VARIETIES OF MEMORY BARRIER +--------------------------- + +Memory barriers come in four basic varieties: + + (1) Write (or store) memory barriers. + + A write memory barrier gives a guarantee that all the STORE operations + specified before the barrier will appear to happen before all the STORE + operations specified after the barrier with respect to the other + components of the system. + + A write barrier is a partial ordering on stores only; it is not required + to have any effect on loads. + + A CPU can be viewed as as commiting a sequence of store operations to the + memory system as time progresses. All stores before a write barrier will + occur in the sequence _before_ all the stores after the write barrier. + + [!] Note that write barriers should normally be paired with read or data + dependency barriers; see the "SMP barrier pairing" subsection. + + + (2) Data dependency barriers. + + A data dependency barrier is a weaker form of read barrier. In the case + where two loads are performed such that the second depends on the result + of the first (eg: the first load retrieves the address to which the second + load will be directed), a data dependency barrier would be required to + make sure that the target of the second load is updated before the address + obtained by the first load is accessed. + + A data dependency barrier is a partial ordering on interdependent loads + only; it is not required to have any effect on stores, independent loads + or overlapping loads. + + As mentioned in (1), the other CPUs in the system can be viewed as + committing sequences of stores to the memory system that the CPU being + considered can then perceive. A data dependency barrier issued by the CPU + under consideration guarantees that for any load preceding it, if that + load touches one of a sequence of stores from another CPU, then by the + time the barrier completes, the effects of all the stores prior to that + touched by the load will be perceptible to any loads issued after the data + dependency barrier. + + See the "Examples of memory barrier sequences" subsection for diagrams + showing the ordering constraints. + + [!] Note that the first load really has to have a _data_ dependency and + not a control dependency. If the address for the second load is dependent + on the first load, but the dependency is through a conditional rather than + actually loading the address itself, then it's a _control_ dependency and + a full read barrier or better is required. See the "Control dependencies" + subsection for more information. + + [!] Note that data dependency barriers should normally be paired with + write barriers; see the "SMP barrier pairing" subsection. + + + (3) Read (or load) memory barriers. + + A read barrier is a data dependency barrier plus a guarantee that all the + LOAD operations specified before the barrier will appear to happen before + all the LOAD operations specified after the barrier with respect to the + other components of the system. + + A read barrier is a partial ordering on loads only; it is not required to + have any effect on stores. + + Read memory barriers imply data dependency barriers, and so can substitute + for them. + + [!] Note that read barriers should normally be paired with write barriers; + see the "SMP barrier pairing" subsection. + + + (4) General memory barriers. + + A general memory barrier is a combination of both a read memory barrier + and a write memory barrier. It is a partial ordering over both loads and + stores. + + General memory barriers imply both read and write memory barriers, and so + can substitute for either. + + +And a couple of implicit varieties: + + (5) LOCK operations. + + This acts as a one-way permeable barrier. It guarantees that all memory + operations after the LOCK operation will appear to happen after the LOCK + operation with respect to the other components of the system. + + Memory operations that occur before a LOCK operation may appear to happen + after it completes. + + A LOCK operation should almost always be paired with an UNLOCK operation. + + + (6) UNLOCK operations. + + This also acts as a one-way permeable barrier. It guarantees that all + memory operations before the UNLOCK operation will appear to happen before + the UNLOCK operation with respect to the other components of the system. + + Memory operations that occur after an UNLOCK operation may appear to + happen before it completes. + + LOCK and UNLOCK operations are guaranteed to appear with respect to each + other strictly in the order specified. + + The use of LOCK and UNLOCK operations generally precludes the need for + other sorts of memory barrier (but note the exceptions mentioned in the + subsection "MMIO write barrier"). + + +Memory barriers are only required where there's a possibility of interaction +between two CPUs or between a CPU and a device. If it can be guaranteed that +there won't be any such interaction in any particular piece of code, then +memory barriers are unnecessary in that piece of code. + + +Note that these are the _minimum_ guarantees. Different architectures may give +more substantial guarantees, but they may _not_ be relied upon outside of arch +specific code. + + +WHAT MAY NOT BE ASSUMED ABOUT MEMORY BARRIERS? +---------------------------------------------- + +There are certain things that the Linux kernel memory barriers do not guarantee: + + (*) 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 the 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/pci.txt + Documentation/DMA-mapping.txt + Documentation/DMA-API.txt + + +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> + P = &B + Q = 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) ???? + +Whilst 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 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> + P = &B + Q = 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 (1). + + +Another example of where data dependency barriers might by required is where a +number is read from memory and then used to calculate the index for an array +access: + + CPU 1 CPU 2 + =============== =============== + { M[0] == 1, M[1] == 2, M[3] = 3, P == 0, Q == 3 } + M[1] = 4; + <write barrier> + P = 1 + Q = P; + <data dependency barrier> + D = M[Q]; + + +The data dependency barrier is very important to the RCU system, for example. +See 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 +-------------------- + +A 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 = &a; + if (p) + q = &b; + <data dependency barrier> + x = *q; + +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. In such a case what's actually +required is: + + q = &a; + if (p) + q = &b; + <read barrier> + x = *q; + + +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. + +A write barrier should always be paired with a data dependency barrier or read +barrier, though a general barrier would also be viable. Similarly a read +barrier or a data dependency barrier should always be paired with at least an +write barrier, though, again, a general barrier is viable: + + CPU 1 CPU 2 + =============== =============== + a = 1; + <write barrier> + b = 2; x = a; + <read barrier> + y = b; + +Or: + + CPU 1 CPU 2 + =============== =============================== + a = 1; + <write barrier> + b = &a; x = b; + <data dependency barrier> + y = *x; + +Basically, the read barrier always has to be there, even though it can be of +the "weaker" type. + + +EXAMPLES OF MEMORY BARRIER SEQUENCES +------------------------------------ + +Firstly, write barriers act as a 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 occuring before the unordered set of { STORE D, STORE E +}: + + +-------+ : : + | | +------+ + | |------>| C=3 | } /\ + | | : +------+ }----- \ -----> Events perceptible + | | : | A=1 | } \/ to rest of 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 be take place. + | |------>| D=4 | } + | | +------+ + +-------+ : : + | + | Sequence in which stores committed to memory system + | by CPU 1 + V + + +Secondly, data dependency barriers act as a partial orderings on data-dependent +loads. Consider the following sequence of events: + + CPU 1 CPU 2 + ======================= ======================= + 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 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, 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 |------>| | + \ +-------+ | | + ----->| B->2 | | | + +-------+ | | + Makes sure all effects ---> ddddddddddddddddd | | + prior to the store of C +-------+ | | + are perceptible to | B->2 |------>| | + successive loads +-------+ | | + : : +-------+ + + +And thirdly, a read barrier acts as a partial order on loads. Consider the +following sequence of events: + + CPU 1 CPU 2 + ======================= ======================= + STORE A=1 + STORE B=2 + STORE C=3 + <write barrier> + STORE D=4 + STORE E=5 + LOAD A + LOAD B + LOAD C + LOAD D + LOAD E + +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: + + +-------+ : : + | | +------+ + | |------>| C=3 | } + | | : +------+ } + | | : | A=1 | } + | | : +------+ } + | CPU 1 | : | B=2 | }--- + | | +------+ } \ + | | wwwwwwwwwwwww} \ + | | +------+ } \ : : +-------+ + | | : | E=5 | } \ +-------+ | | + | | : +------+ } \ { | C->3 |------>| | + | |------>| D=4 | } \ { +-------+ : | | + | | +------+ \ { | E->5 | : | | + +-------+ : : \ { +-------+ : | | + Transfer -->{ | A->1 | : | CPU 2 | + from CPU 1 { +-------+ : | | + to CPU 2 { | D->4 | : | | + { +-------+ : | | + { | B->2 |------>| | + +-------+ | | + : : +-------+ + + +If, however, a read barrier were to be placed between the load of C and the +load of D on CPU 2, then the partial ordering imposed by CPU 1 will be +perceived correctly by CPU 2. + + +-------+ : : + | | +------+ + | |------>| C=3 | } + | | : +------+ } + | | : | A=1 | }--- + | | : +------+ } \ + | CPU 1 | : | B=2 | } \ + | | +------+ \ + | | wwwwwwwwwwwwwwww \ + | | +------+ \ : : +-------+ + | | : | E=5 | } \ +-------+ | | + | | : +------+ }--- \ { | C->3 |------>| | + | |------>| D=4 | } \ \ { +-------+ : | | + | | +------+ \ -->{ | B->2 | : | | + +-------+ : : \ { +-------+ : | | + \ { | A->1 | : | CPU 2 | + \ +-------+ | | + At this point the read ----> \ rrrrrrrrrrrrrrrrr | | + barrier causes all effects \ +-------+ | | + prior to the storage of C \ { | E->5 | : | | + to be perceptible to CPU 2 -->{ +-------+ : | | + { | D->4 |------>| | + +-------+ | | + : : +-------+ + + +======================== +EXPLICIT KERNEL BARRIERS +======================== + +The Linux kernel has a variety of different barriers that act at different +levels: + + (*) Compiler barrier. + + (*) CPU memory barriers. + + (*) MMIO write barrier. + + +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 a general barrier - lesser varieties of compiler barrier do not exist. + +The compiler barrier has 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_barrier_depends() smp_read_barrier_depends() + + +All CPU memory barriers unconditionally imply compiler barriers. + +SMP memory barriers are reduced to compiler barriers on uniprocessor compiled +systems because it is assumed that a CPU will be appear to be self-consistent, +and will order overlapping accesses correctly with respect to itself. + +[!] 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 unnecessarily impose overhead on UP systems. They may, however, be +used to control MMIO effects on accesses through relaxed memory I/O windows. +These 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: + + (*) set_mb(var, value) + (*) set_wmb(var, value) + + These assign the value to the variable and then insert at least a write + barrier after it, depending on the function. They aren't guaranteed to + insert anything more than a compiler barrier in a UP compilation. + + + (*) smp_mb__before_atomic_dec(); + (*) smp_mb__after_atomic_dec(); + (*) smp_mb__before_atomic_inc(); + (*) smp_mb__after_atomic_inc(); + + These are for use with atomic add, subtract, increment and decrement + functions, especially when used for reference counting. These functions + do not imply memory barriers. + + 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_dec(); + 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_ops.txt for more information. See the "Atomic + operations" subsection for information on where to use these. + + + (*) smp_mb__before_clear_bit(void); + (*) smp_mb__after_clear_bit(void); + + These are for use similar to the atomic inc/dec barriers. These are + typically used for bitwise unlocking operations, so care must be taken as + there are no implicit memory barriers here either. + + Consider implementing an unlock operation of some nature by clearing a + locking bit. The clear_bit() would then need to be barriered like this: + + smp_mb__before_clear_bit(); + clear_bit( ... ); + + This prevents memory operations before the clear leaking to after it. See + the subsection on "Locking Functions" with reference to UNLOCK operation + implications. + + See Documentation/atomic_ops.txt for more information. See the "Atomic + operations" subsection for information on where to use these. + + +MMIO WRITE BARRIER +------------------ + +The Linux kernel also has a special barrier for use with memory-mapped I/O +writes: + + mmiowb(); + +This is a variation on the mandatory write barrier that causes writes to weakly +ordered I/O regions to be partially ordered. Its effects may go beyond the +CPU->Hardware interface and actually affect the hardware at some level. + +See the subsection "Locks vs I/O accesses" for more information. + + +=============================== +IMPLICIT KERNEL MEMORY BARRIERS +=============================== + +Some of the other functions in the linux kernel imply memory barriers, amongst +which are locking, scheduling and memory allocation 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. + + +LOCKING FUNCTIONS +----------------- + +The Linux kernel has a number of locking constructs: + + (*) spin locks + (*) R/W spin locks + (*) mutexes + (*) semaphores + (*) R/W semaphores + (*) RCU + +In all cases there are variants on "LOCK" operations and "UNLOCK" operations +for each construct. These operations all imply certain barriers: + + (1) LOCK operation implication: + + Memory operations issued after the LOCK will be completed after the LOCK + operation has completed. + + Memory operations issued before the LOCK may be completed after the LOCK + operation has completed. + + (2) UNLOCK operation implication: + + Memory operations issued before the UNLOCK will be completed before the + UNLOCK operation has completed. + + Memory operations issued after the UNLOCK may be completed before the + UNLOCK operation has completed. + + (3) LOCK vs LOCK implication: + + All LOCK operations issued before another LOCK operation will be completed + before that LOCK operation. + + (4) LOCK vs UNLOCK implication: + + All LOCK operations issued before an UNLOCK operation will be completed + before the UNLOCK operation. + + All UNLOCK operations issued before a LOCK operation will be completed + before the LOCK operation. + + (5) Failed conditional LOCK implication: + + Certain variants of the LOCK operation may fail, either due to being + unable to get the lock immediately, or due to receiving an unblocked + signal whilst asleep waiting for the lock to become available. Failed + locks do not imply any sort of barrier. + +Therefore, from (1), (2) and (4) an UNLOCK followed by an unconditional LOCK is +equivalent to a full barrier, but a LOCK followed by an UNLOCK is not. + +[!] Note: one of the consequence of LOCKs and UNLOCKs being only one-way + barriers is that the effects instructions outside of a critical section may + seep into the inside of the critical section. + +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 locking barrier effects". + + +As an example, consider the following: + + *A = a; + *B = b; + LOCK + *C = c; + *D = d; + UNLOCK + *E = e; + *F = f; + +The following sequence of events is acceptable: + + LOCK, {*F,*A}, *E, {*C,*D}, *B, UNLOCK + + [+] Note that {*F,*A} indicates a combined access. + +But none of the following are: + + {*F,*A}, *B, LOCK, *C, *D, UNLOCK, *E + *A, *B, *C, LOCK, *D, UNLOCK, *E, *F + *A, *B, LOCK, *C, UNLOCK, *D, *E, *F + *B, LOCK, *C, *D, UNLOCK, {*F,*A}, *E + + + +INTERRUPT DISABLING FUNCTIONS +----------------------------- + +Functions that disable interrupts (LOCK equivalent) and enable interrupts +(UNLOCK 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. + + +MISCELLANEOUS FUNCTIONS +----------------------- + +Other functions that imply barriers: + + (*) schedule() and similar imply full memory barriers. + + (*) Memory allocation and release functions imply full memory barriers. + + +================================= +INTER-CPU LOCKING 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. + + +LOCKS VS MEMORY ACCESSES +------------------------ + +Consider the following: the system has a pair of spinlocks (N) and (Q), and +three CPUs; then should the following sequence of events occur: + + CPU 1 CPU 2 + =============================== =============================== + *A = a; *E = e; + LOCK M LOCK Q + *B = b; *F = f; + *C = c; *G = g; + UNLOCK M UNLOCK Q + *D = d; *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, LOCK M, LOCK Q, *G, *C, *F, *A, *B, UNLOCK Q, *D, *H, UNLOCK M + +But it won't see any of: + + *B, *C or *D preceding LOCK M + *A, *B or *C following UNLOCK M + *F, *G or *H preceding LOCK Q + *E, *F or *G following UNLOCK Q + + +However, if the following occurs: + + CPU 1 CPU 2 + =============================== =============================== + *A = a; + LOCK M [1] + *B = b; + *C = c; + UNLOCK M [1] + *D = d; *E = e; + LOCK M [2] + *F = f; + *G = g; + UNLOCK M [2] + *H = h; + +CPU #3 might see: + + *E, LOCK M [1], *C, *B, *A, UNLOCK M [1], + LOCK M [2], *H, *F, *G, UNLOCK M [2], *D + +But assuming CPU #1 gets the lock first, it won't see any of: + + *B, *C, *D, *F, *G or *H preceding LOCK M [1] + *A, *B or *C following UNLOCK M [1] + *F, *G or *H preceding LOCK M [2] + *A, *B, *C, *E, *F or *G following UNLOCK M [2] + + +LOCKS VS I/O ACCESSES +--------------------- + +Under certain circumstances (especially involving NUMA), I/O accesses within +two spinlocked sections on two different CPUs may be seen as interleaved by the +PCI bridge, because the PCI bridge does not necessarily participate in the +cache-coherence protocol, and is therefore incapable of issuing the required +read memory barriers. + +For example: + + CPU 1 CPU 2 + =============================== =============================== + spin_lock(Q) + writel(0, ADDR) + writel(1, DATA); + spin_unlock(Q); + spin_lock(Q); + writel(4, ADDR); + writel(5, DATA); + spin_unlock(Q); + +may be seen by the PCI bridge as follows: + + STORE *ADDR = 0, STORE *ADDR = 4, STORE *DATA = 1, STORE *DATA = 5 + +which would probably cause the hardware to malfunction. + + +What is necessary here is to intervene with an mmiowb() before dropping the +spinlock, for example: + + CPU 1 CPU 2 + =============================== =============================== + spin_lock(Q) + writel(0, ADDR) + writel(1, DATA); + mmiowb(); + spin_unlock(Q); + spin_lock(Q); + writel(4, ADDR); + writel(5, DATA); + mmiowb(); + spin_unlock(Q); + +this will ensure that the two stores issued on CPU #1 appear at the PCI bridge +before either of the stores issued on CPU #2. + + +Furthermore, following a store by a load to the same device obviates the need +for an mmiowb(), because the load forces the store to complete before the load +is performed: + + CPU 1 CPU 2 + =============================== =============================== + spin_lock(Q) + writel(0, ADDR) + a = readl(DATA); + spin_unlock(Q); + spin_lock(Q); + writel(4, ADDR); + b = readl(DATA); + spin_unlock(Q); + + +See Documentation/DocBook/deviceiobook.tmpl for more information. + + +================================= +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, three +circumstances in which reordering definitely _could_ be a problem: + + (*) Interprocessor interaction. + + (*) Atomic operations. + + (*) Accessing devices (I/O). + + (*) 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; + + (4) 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 otherwords, 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 there's only +one CPU, that CPU's dependency ordering logic will take care of everything +else. + + +ATOMIC OPERATIONS +----------------- + +Though they are technically interprocessor interaction considerations, atomic +operations are noted specially as they do _not_ generally imply memory +barriers. The possible offenders include: + + xchg(); + cmpxchg(); + test_and_set_bit(); + test_and_clear_bit(); + test_and_change_bit(); + atomic_cmpxchg(); + atomic_inc_return(); + atomic_dec_return(); + atomic_add_return(); + atomic_sub_return(); + atomic_inc_and_test(); + atomic_dec_and_test(); + atomic_sub_and_test(); + atomic_add_negative(); + atomic_add_unless(); + +These may be used for such things as implementing LOCK operations or controlling +the lifetime of objects by decreasing their reference counts. In such cases +they need preceding memory barriers. + +The following may also be possible offenders as they may be used as UNLOCK +operations. + + set_bit(); + clear_bit(); + change_bit(); + atomic_set(); + + +The following are a little tricky: + + atomic_add(); + atomic_sub(); + atomic_inc(); + atomic_dec(); + +If they're used for statistics generation, then they probably don't need memory +barriers, unless there's a coupling between statistical data. + +If they're used for reference counting on an object to control its lifetime, +they probably don't need memory barriers because either the reference count +will be adjusted inside a locked section, or the caller will already hold +sufficient references to make the lock, and thus a memory barrier unnecessary. + +If they're used for constructing a lock of some description, then they probably +do need memory barriers as a lock primitive generally has to do things in a +specific order. + + +Basically, each usage case has to be carefully considered as to whether memory +barriers are needed or not. The simplest rule is probably: if the atomic +operation is protected by a lock, then it does not require a barrier unless +there's another operation within the critical section with respect to which an +ordering must be maintained. + +See Documentation/atomic_ops.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. Whilst this, for the most part, renders the explicit +use of memory barriers unnecessary, there are a couple of situations where they +might be needed: + + (1) On some systems, I/O stores are not strongly ordered across all CPUs, and + so for _all_ general drivers locks should be used and mmiowb() must be + issued prior to unlocking the critical section. + + (2) 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/DocBook/deviceiobook.tmpl 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. Whilst 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. If this isn't sufficient then an +mmiowb() may need to be used explicitly. + + +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 +========================== + +When accessing I/O memory, drivers should use the appropriate accessor +functions: + + (*) inX(), outX(): + + These are intended to talk to I/O space rather than memory space, but + that's primarily a CPU-specific concept. The i386 and x86_64 processors do + indeed have special I/O space access cycles and instructions, but many + CPUs don't have such a concept. + + The PCI bus, amongst others, defines an I/O space concept - which on such + CPUs as i386 and x86_64 cpus readily maps to the CPU's concept of I/O + space. However, it may also mapped as a virtual I/O space in the CPU's + memory map, particularly on those CPUs that don't support alternate + I/O spaces. + + Accesses to this space may be fully synchronous (as on i386), but + intermediary bridges (such as the PCI host bridge) may not fully honour + that. + + They are guaranteed to be fully ordered with respect to each other. + + They are not guaranteed to be fully ordered with respect to other types of + memory and I/O operation. + + (*) readX(), writeX(): + + Whether these are guaranteed to be fully ordered and uncombined with + respect to each other on the issuing CPU depends on the characteristics + defined for the memory window through which they're accessing. On later + i386 architecture machines, for example, this is controlled by way of the + MTRR registers. + + Ordinarily, these will be guaranteed to be fully ordered and uncombined,, + provided they're not accessing a prefetchable device. + + However, intermediary hardware (such as a PCI bridge) may indulge in + deferral if it so wishes; to flush a store, a load from the same location + is preferred[*], but a load from the same device or from configuration + space should suffice for PCI. + + [*] NOTE! attempting to load from the same location as was written to may + cause a malfunction - consider the 16550 Rx/Tx serial registers for + example. + + Used with prefetchable I/O memory, an mmiowb() barrier may be required to + force stores to be ordered. + + Please refer to the PCI specification for more information on interactions + between PCI transactions. + + (*) readX_relaxed() + + These are similar to readX(), but are not guaranteed to be ordered in any + way. Be aware that there is no I/O read barrier available. + + (*) ioreadX(), iowriteX() + + These will perform as appropriate for the type of access they're actually + doing, be it inX()/outX() or readX()/writeX(). + + +======================================== +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 the 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 whilst 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 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; + + (*) whilst 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 effect 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 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 that 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 expected order, as the +cacheline holding p may get updated in one of the second CPU's caches whilst +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 = &b; 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, whilst 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. 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 = &b; 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. +Whilst 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 accesss. The semantics of the Alpha removes the +need for coordination in absence of memory barriers. + + +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 simply present in a CPUs 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/cachetlb.txt 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 have 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 a CPU is, for example, +given the following piece of code to execute: + + a = *A; + *B = b; + c = *C; + d = *D; + *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 being + 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 whilst 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 = *A; + *A = V; + *A = W; + X = *A; + *A = Y; + Z = *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. + +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 a write barrier, 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, 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 updating 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's kernel's memory barrier model. + +See the subsection on "Cache Coherency" above. + + +========== +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 + +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 + +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 + +Intel Itanium Architecture Software Developer's Manual: Volume 1: + Section 2.6: Speculation + Section 4.4: Memory Access diff --git a/Documentation/networking/packet_mmap.txt b/Documentation/networking/packet_mmap.txt index 4fc8e98..aaf99d5 100644 --- a/Documentation/networking/packet_mmap.txt +++ b/Documentation/networking/packet_mmap.txt @@ -254,7 +254,7 @@ and, the number of frames be <block number> * <block size> / <frame size> -Suposse the following parameters, which apply for 2.6 kernel and an +Suppose the following parameters, which apply for 2.6 kernel and an i386 architecture: <size-max> = 131072 bytes diff --git a/Documentation/networking/tuntap.txt b/Documentation/networking/tuntap.txt index ec3d109..76750fb 100644 --- a/Documentation/networking/tuntap.txt +++ b/Documentation/networking/tuntap.txt @@ -138,7 +138,7 @@ This means that you have to read/write IP packets when you are using tun and ethernet frames when using tap. 5. What is the difference between BPF and TUN/TAP driver? -BFP is an advanced packet filter. It can be attached to existing +BPF is an advanced packet filter. It can be attached to existing network interface. It does not provide a virtual network interface. A TUN/TAP driver does provide a virtual network interface and it is possible to attach BPF to this interface. diff --git a/Documentation/pcmcia/driver-changes.txt b/Documentation/pcmcia/driver-changes.txt index 97420f0..4739c5c 100644 --- a/Documentation/pcmcia/driver-changes.txt +++ b/Documentation/pcmcia/driver-changes.txt @@ -1,5 +1,11 @@ This file details changes in 2.6 which affect PCMCIA card driver authors: +* New release helper (as of 2.6.17) + Instead of calling pcmcia_release_{configuration,io,irq,win}, all that's + necessary now is calling pcmcia_disable_device. As there is no valid + reason left to call pcmcia_release_io and pcmcia_release_irq, the + exports for them were removed. + * Unify detach and REMOVAL event code, as well as attach and INSERTION code (as of 2.6.16) void (*remove) (struct pcmcia_device *dev); diff --git a/Documentation/video4linux/CARDLIST.saa7134 b/Documentation/video4linux/CARDLIST.saa7134 index 8c71954..bca5090 100644 --- a/Documentation/video4linux/CARDLIST.saa7134 +++ b/Documentation/video4linux/CARDLIST.saa7134 @@ -52,7 +52,7 @@ 51 -> ProVideo PV952 [1540:9524] 52 -> AverMedia AverTV/305 [1461:2108] 53 -> ASUS TV-FM 7135 [1043:4845] - 54 -> LifeView FlyTV Platinum FM [5168:0214,1489:0214] + 54 -> LifeView FlyTV Platinum FM / Gold [5168:0214,1489:0214,5168:0304] 55 -> LifeView FlyDVB-T DUO [5168:0306] 56 -> Avermedia AVerTV 307 [1461:a70a] 57 -> Avermedia AVerTV GO 007 FM [1461:f31f] @@ -84,7 +84,7 @@ 83 -> Terratec Cinergy 250 PCI TV [153b:1160] 84 -> LifeView FlyDVB Trio [5168:0319] 85 -> AverTV DVB-T 777 [1461:2c05] - 86 -> LifeView FlyDVB-T [5168:0301] + 86 -> LifeView FlyDVB-T / Genius VideoWonder DVB-T [5168:0301,1489:0301] 87 -> ADS Instant TV Duo Cardbus PTV331 [0331:1421] 88 -> Tevion/KWorld DVB-T 220RF [17de:7201] 89 -> ELSA EX-VISION 700TV [1048:226c] @@ -92,3 +92,4 @@ 91 -> AVerMedia A169 B [1461:7360] 92 -> AVerMedia A169 B1 [1461:6360] 93 -> Medion 7134 Bridge #2 [16be:0005] + 94 -> LifeView FlyDVB-T Hybrid Cardbus [5168:3306,5168:3502] diff --git a/Documentation/usb/et61x251.txt b/Documentation/video4linux/et61x251.txt index 2934028..2934028 100644 --- a/Documentation/usb/et61x251.txt +++ b/Documentation/video4linux/et61x251.txt diff --git a/Documentation/usb/ibmcam.txt b/Documentation/video4linux/ibmcam.txt index c250036..4a40a2e 100644 --- a/Documentation/usb/ibmcam.txt +++ b/Documentation/video4linux/ibmcam.txt @@ -122,7 +122,7 @@ WHAT YOU NEED: - A Linux box with USB support (2.3/2.4; 2.2 w/backport may work) - A Video4Linux compatible frame grabber program such as xawtv. - + HOW TO COMPILE THE DRIVER: You need to compile the driver only if you are a developer diff --git a/Documentation/usb/ov511.txt b/Documentation/video4linux/ov511.txt index a7fc043..142741e 100644 --- a/Documentation/usb/ov511.txt +++ b/Documentation/video4linux/ov511.txt @@ -9,7 +9,7 @@ INTRODUCTION: This is a driver for the OV511, a USB-only chip used in many "webcam" devices. Any camera using the OV511/OV511+ and the OV6620/OV7610/20/20AE should work. -Video capture devices that use the Philips SAA7111A decoder also work. It +Video capture devices that use the Philips SAA7111A decoder also work. It supports streaming and capture of color or monochrome video via the Video4Linux API. Most V4L apps are compatible with it. Most resolutions with a width and height that are a multiple of 8 are supported. @@ -52,15 +52,15 @@ from it: chmod 666 /dev/video chmod 666 /dev/video0 (if necessary) - + Now you are ready to run a video app! Both vidcat and xawtv work well for me at 640x480. - + [Using vidcat:] vidcat -s 640x480 -p c > test.jpg xview test.jpg - + [Using xawtv:] From the main xawtv directory: @@ -70,7 +70,7 @@ From the main xawtv directory: make make install -Now you should be able to run xawtv. Right click for the options dialog. +Now you should be able to run xawtv. Right click for the options dialog. MODULE PARAMETERS: @@ -286,4 +286,3 @@ Randy Dunlap, and others. Big thanks to them for their pioneering work on that and the USB stack. Thanks to Bret Wallach for getting camera reg IO, ISOC, and image capture working. Thanks to Orion Sky Lawlor, Kevin Moore, and Claudio Matsuoka for their work as well. - diff --git a/Documentation/usb/se401.txt b/Documentation/video4linux/se401.txt index 7b9d1c9..7b9d1c9 100644 --- a/Documentation/usb/se401.txt +++ b/Documentation/video4linux/se401.txt diff --git a/Documentation/usb/sn9c102.txt b/Documentation/video4linux/sn9c102.txt index b957bea..142920b 100644 --- a/Documentation/usb/sn9c102.txt +++ b/Documentation/video4linux/sn9c102.txt @@ -174,7 +174,7 @@ Module parameters are listed below: ------------------------------------------------------------------------------- Name: video_nr Type: short array (min = 0, max = 64) -Syntax: <-1|n[,...]> +Syntax: <-1|n[,...]> Description: Specify V4L2 minor mode number: -1 = use next available n = use minor number n @@ -187,7 +187,7 @@ Default: -1 ------------------------------------------------------------------------------- Name: force_munmap Type: bool array (min = 0, max = 64) -Syntax: <0|1[,...]> +Syntax: <0|1[,...]> Description: Force the application to unmap previously mapped buffer memory before calling any VIDIOC_S_CROP or VIDIOC_S_FMT ioctl's. Not all the applications support this feature. This parameter is @@ -206,7 +206,7 @@ Default: 2 ------------------------------------------------------------------------------- Name: debug Type: ushort -Syntax: <n> +Syntax: <n> Description: Debugging information level, from 0 to 3: 0 = none (use carefully) 1 = critical errors @@ -267,7 +267,7 @@ The sysfs interface also provides the "frame_header" entry, which exports the frame header of the most recent requested and captured video frame. The header is always 18-bytes long and is appended to every video frame by the SN9C10x controllers. As an example, this additional information can be used by the user -application for implementing auto-exposure features via software. +application for implementing auto-exposure features via software. The following table describes the frame header: @@ -441,7 +441,7 @@ blue pixels in one video frame. Each pixel is associated with a 8-bit long value and is disposed in memory according to the pattern shown below: B[0] G[1] B[2] G[3] ... B[m-2] G[m-1] -G[m] R[m+1] G[m+2] R[m+2] ... G[2m-2] R[2m-1] +G[m] R[m+1] G[m+2] R[m+2] ... G[2m-2] R[2m-1] ... ... B[(n-1)(m-2)] G[(n-1)(m-1)] ... G[n(m-2)] R[n(m-1)] @@ -472,12 +472,12 @@ The pixel reference value is calculated as follows: The algorithm purely describes the conversion from compressed Bayer code used in the SN9C10x chips to uncompressed Bayer. Additional steps are required to convert this to a color image (i.e. a color interpolation algorithm). - + The following Huffman codes have been found: -0: +0 (relative to reference pixel value) +0: +0 (relative to reference pixel value) 100: +4 101: -4? -1110xxxx: set absolute value to xxxx.0000 +1110xxxx: set absolute value to xxxx.0000 1101: +11 1111: -11 11001: +20 diff --git a/Documentation/usb/stv680.txt b/Documentation/video4linux/stv680.txt index 6448041..4f8946f 100644 --- a/Documentation/usb/stv680.txt +++ b/Documentation/video4linux/stv680.txt @@ -5,15 +5,15 @@ Copyright, 2001, Kevin Sisson INTRODUCTION: -STMicroelectronics produces the STV0680B chip, which comes in two -types, -001 and -003. The -003 version allows the recording and downloading -of sound clips from the camera, and allows a flash attachment. Otherwise, -it uses the same commands as the -001 version. Both versions support a -variety of SDRAM sizes and sensors, allowing for a maximum of 26 VGA or 20 -CIF pictures. The STV0680 supports either a serial or a usb interface, and +STMicroelectronics produces the STV0680B chip, which comes in two +types, -001 and -003. The -003 version allows the recording and downloading +of sound clips from the camera, and allows a flash attachment. Otherwise, +it uses the same commands as the -001 version. Both versions support a +variety of SDRAM sizes and sensors, allowing for a maximum of 26 VGA or 20 +CIF pictures. The STV0680 supports either a serial or a usb interface, and video is possible through the usb interface. -The following cameras are known to work with this driver, although any +The following cameras are known to work with this driver, although any camera with Vendor/Product codes of 0553/0202 should work: Aiptek Pencam (various models) @@ -34,15 +34,15 @@ http://www.linux-usb.org MODULE OPTIONS: When the driver is compiled as a module, you can set a "swapRGB=1" -option, if necessary, for those applications that require it -(such as xawtv). However, the driver should detect and set this +option, if necessary, for those applications that require it +(such as xawtv). However, the driver should detect and set this automatically, so this option should not normally be used. KNOWN PROBLEMS: -The driver seems to work better with the usb-ohci than the usb-uhci host -controller driver. +The driver seems to work better with the usb-ohci than the usb-uhci host +controller driver. HELP: @@ -50,6 +50,4 @@ The latest info on this driver can be found at: http://personal.clt.bellsouth.net/~kjsisson or at http://stv0680-usb.sourceforge.net -Any questions to me can be send to: kjsisson@bellsouth.net - - +Any questions to me can be send to: kjsisson@bellsouth.net
\ No newline at end of file diff --git a/Documentation/usb/w9968cf.txt b/Documentation/video4linux/w9968cf.txt index 9d46cd0..3b704f2 100644 --- a/Documentation/usb/w9968cf.txt +++ b/Documentation/video4linux/w9968cf.txt @@ -1,5 +1,5 @@ - W996[87]CF JPEG USB Dual Mode Camera Chip + W996[87]CF JPEG USB Dual Mode Camera Chip Driver for Linux 2.6 (basic version) ========================================= @@ -115,7 +115,7 @@ additional testing and full support, would be much appreciated. ====================== For it to work properly, the driver needs kernel support for Video4Linux, USB and I2C, and the "ovcamchip" module for the image sensor. Make sure you are not -actually using any external "ovcamchip" module, given that the W996[87]CF +actually using any external "ovcamchip" module, given that the W996[87]CF driver depends on the version of the module present in the official kernels. The following options of the kernel configuration file must be enabled and @@ -197,16 +197,16 @@ Note: The kernel must be compiled with the CONFIG_KMOD option enabled for the 'ovcamchip' module to be loaded and for this parameter to be present. ------------------------------------------------------------------------------- -Name: simcams -Type: int -Syntax: <n> +Name: simcams +Type: int +Syntax: <n> Description: Number of cameras allowed to stream simultaneously. n may vary from 0 to 32. Default: 32 ------------------------------------------------------------------------------- Name: video_nr Type: int array (min = 0, max = 32) -Syntax: <-1|n[,...]> +Syntax: <-1|n[,...]> Description: Specify V4L minor mode number. -1 = use next available n = use minor number n @@ -219,7 +219,7 @@ Default: -1 ------------------------------------------------------------------------------- Name: packet_size Type: int array (min = 0, max = 32) -Syntax: <n[,...]> +Syntax: <n[,...]> Description: Specify the maximum data payload size in bytes for alternate settings, for each device. n is scaled between 63 and 1023. Default: 1023 @@ -234,7 +234,7 @@ Default: 2 ------------------------------------------------------------------------------- Name: double_buffer Type: bool array (min = 0, max = 32) -Syntax: <0|1[,...]> +Syntax: <0|1[,...]> Description: Hardware double buffering: 0 disabled, 1 enabled. It should be enabled if you want smooth video output: if you obtain out of sync. video, disable it, or try to @@ -243,13 +243,13 @@ Default: 1 for every device. ------------------------------------------------------------------------------- Name: clamping Type: bool array (min = 0, max = 32) -Syntax: <0|1[,...]> +Syntax: <0|1[,...]> Description: Video data clamping: 0 disabled, 1 enabled. Default: 0 for every device. ------------------------------------------------------------------------------- Name: filter_type Type: int array (min = 0, max = 32) -Syntax: <0|1|2[,...]> +Syntax: <0|1|2[,...]> Description: Video filter type. 0 none, 1 (1-2-1) 3-tap filter, 2 (2-3-6-3-2) 5-tap filter. The filter is used to reduce noise and aliasing artifacts @@ -258,13 +258,13 @@ Default: 0 for every device. ------------------------------------------------------------------------------- Name: largeview Type: bool array (min = 0, max = 32) -Syntax: <0|1[,...]> +Syntax: <0|1[,...]> Description: Large view: 0 disabled, 1 enabled. Default: 1 for every device. ------------------------------------------------------------------------------- Name: upscaling Type: bool array (min = 0, max = 32) -Syntax: <0|1[,...]> +Syntax: <0|1[,...]> Description: Software scaling (for non-compressed video only): 0 disabled, 1 enabled. Disable it if you have a slow CPU or you don't have enough @@ -341,8 +341,8 @@ Default: 50 for every device. ------------------------------------------------------------------------------- Name: bandingfilter Type: bool array (min = 0, max = 32) -Syntax: <0|1[,...]> -Description: Banding filter to reduce effects of fluorescent +Syntax: <0|1[,...]> +Description: Banding filter to reduce effects of fluorescent lighting: 0 disabled, 1 enabled. This filter tries to reduce the pattern of horizontal @@ -374,7 +374,7 @@ Default: 0 for every device. ------------------------------------------------------------------------------- Name: monochrome Type: bool array (min = 0, max = 32) -Syntax: <0|1[,...]> +Syntax: <0|1[,...]> Description: The image sensor is monochrome: 0 = no, 1 = yes Default: 0 for every device. @@ -400,19 +400,19 @@ Default: 32768 for every device. ------------------------------------------------------------------------------- Name: contrast Type: long array (min = 0, max = 32) -Syntax: <n[,...]> +Syntax: <n[,...]> Description: Set picture contrast (0-65535). Default: 50000 for every device. ------------------------------------------------------------------------------- Name: whiteness Type: long array (min = 0, max = 32) -Syntax: <n[,...]> +Syntax: <n[,...]> Description: Set picture whiteness (0-65535). Default: 32768 for every device. ------------------------------------------------------------------------------- Name: debug Type: int -Syntax: <n> +Syntax: <n> Description: Debugging information level, from 0 to 6: 0 = none (use carefully) 1 = critical errors diff --git a/Documentation/usb/zc0301.txt b/Documentation/video4linux/zc0301.txt index f55262c..f55262c 100644 --- a/Documentation/usb/zc0301.txt +++ b/Documentation/video4linux/zc0301.txt |
