Memory consistency protection in a multiprocessor computing system

[0001] This application claims benefit of GB Patent Application No. 0623276.3 filed Nov. 22, 2006 and U.S. Provisional Patent Application Ser. No. 60/879,834, filed Jan. 11, 2007.

[0002] 1. Field of the Invention

[0003] The present invention relates generally to the field of computers and computer systems. More particularly, the present invention relates to the protection of memory consistency in a multiprocessor computing system.

[0004] 2. Description of the Related Art

[0005] Modern needs for high-powered computing systems have resulted in the development of multiprocessor computer architectures having two, four, eight or more separate processors. Such multiprocessor systems are able to execute multiple portions of program code simultaneously (i.e. execute multiple processes and/or multiple process threads simultaneously). Several different mechanisms exist to share data between the executing portions of program code including, for example, message passing for inter-process communication (IPC). However, most modern multiprocessor computing systems also support the use of shared memory that is accessible by two or more program code portions running on separate processors.

[0006] It is important that any changes to the data stored in the shared memory are made visible to each of the multiple code portions in an orderly and synchronised manner. Hence, each different type of multiprocessor system has its own corresponding memory consistency model that specifies the semantics of memory operations (particularly relating to load, store and atomic operations) that thereby defines the way in which changes to shared memory are made visible in each of the multiple processors. The program code and the hardware in the multiprocessor system should both adhere to the memory consistency model in order to achieve correct operation. Conversely, a memory consistency failure may lead to a fatal crash of the system.

[0007] A more detailed introduction to memory consistency models in multiprocessor computing systems is provided in “Shared Memory Consistency Models: A Tutorial” by Sarita V. Advey and Kourosh Gharachorlooz, published as Rice University ECE Technical Report 9512 and Western Research Laboratory Research Report 95/7 dated September 1995, the disclosure which is incorporated herein by reference.

[0008] In the simplest example, the memory consistency model specifies sequential consistency whereby the memory operations appear to take place strictly in program order as specified in the program code. However, the processors and memory subsystems in a multiprocessor architecture are often designed to reorder memory operations to achieve improved hardware performance. That is, many modern shared-memory multiprocessor systems such as Digital ALPHA, SPARC v8 & v9 and IBM POWER and others provide various forms of relaxed ordering and offer subtly different forms of non-sequential memory consistency. Here, further general background information in the field of memory consistency is provided in an article entitled “POWER4 and shared memory synchronisation” by B. Hay and G. Hook at http://www-128.ibm.com/developerworks/eserver/articles/power4_mem.html of 24 Apr. 2002, the disclosure of which is incorporated herein by reference.

[0009] According to the present invention there is provided a multiprocessor computer system and a method to protect memory consistency in a multiprocessor computer system, as set forth in the claims appended hereto. Other, optional, features of the invention will be apparent from the dependent claims and the description which follows.

[0010] The example embodiments of the present invention discussed herein concern the protection of memory consistency in a multiprocessor computing system. In particular, the exemplary embodiments of the present invention concern a mechanism to provide consistent and synchronised operations in relation to shared memory in a multiprocessor computer system.

[0011] The following is a summary of various aspects and advantages realizable according to embodiments of the invention. It is provided as an introduction to assist those skilled in the art to more rapidly assimilate the detailed design discussion that ensues and does not and is not intended in any way to limit the scope of the claims that are appended hereto.

[0012] In one exemplary aspect of the present invention there is provided a multiprocessor computing system, comprising a memory, a plurality of processors and a controller unit. The memory has a plurality of target code portions stored therein, including at least a first target code portion and a second target code portion. The plurality of processors are arranged to execute the plurality of target code portions stored in the memory, using parallel processing. The controller unit is arranged to control execution of the plurality of target code portions by the plurality of processors. The controller unit comprises an address space allocation unit arranged to divide a virtual address space used to address the memory into a plurality of virtual address space regions and to control execution of the plurality of target code portions to access the memory though the plurality of virtual address space regions initially according to a first memory consistency model. Also, the controller unit comprises a shared memory detection unit arranged to detect a memory access request made in execution of the first target code portion with respect to a shared memory area in the memory which is also accessible by at least the second target code portion and to identify at least one group of instructions in the first target code portion which access the shared memory area. Further, a memory protection unit is arranged to selectively apply a memory consistency protection to enforce a second memory consistency model in relation to accesses to the shared memory area in execution of the identified group of instructions in the first target code portion, responsive to the shared memory detection unit identifying the identified group of instructions.

[0013] In another exemplary aspect of the invention there is provided a method to protect memory consistency in a multiprocessor computing system, comprising the computer-implemented steps of: executing at least a first code portion under a first memory consistency model; detecting a shared memory area that is accessible by both the first code portion and at least a second code portion; identifying a group of instructions in the first code portion that access the shared memory area; and selectively applying a memory consistency protection such that the group of instructions in the first code portion execute under a second memory consistency model when accessing the detected shared memory area.

[0014] In another exemplary aspect of the invention there is provided a method to protect memory consistency in a multiprocessor computing system, comprising the computer-implemented steps of: executing at least a first program code portion and a second program code portion under a default memory consistency model of the multiprocessor computing system; detecting a request to initiate an explicitly shared memory area accessible to both the first program code portion and the second program code portion; and applying a memory consistency protection selectively such that the first and second program code portions execute under a different second memory consistency model when accessing the shared memory area.

[0015] In another exemplary aspect of the invention there is provided a method to protect memory consistency in a multiprocessor computing system having at least first and second processors, comprising the computer-implemented steps of: executing a first code portion on the first processor with respect to a first virtual address space region and executing a second code portion on the second processor with respect to a second virtual address space region, wherein the first and second virtual address space regions are separate and non-overlapping; mapping at least one mapped area within the first virtual address space region; detecting an access request by the second code portion to a memory area which is unmapped in the second virtual address space region but which is mapped area in the first virtual address space region and thereby determining that the second code portion is attempting to access a shared memory area; remapping the mapped area from the first virtual address space region instead into a third virtual address space region, wherein the third virtual address space is separate and non-overlapping with respect to the first and second address space regions; and modifying at least a selected block of instructions in the second code portion containing the access request to apply a memory consistency protection which causes the block to execute under predetermined memory ordering constraints and directing at least the selected block of code to access the shared memory area mapped in the third address space region.

[0016] The example embodiments provide a low-cost and effective mechanism to protect memory consistency in a multiprocessor computing system, in particular by identifying and applying memory consistency protection selectively only to those parts of a program which demand such protection. Further, the exemplary embodiments provide a low-cost and effective mechanism for detecting accesses to shared memory and thereby determining those parts of the program which require the memory consistency protection.

[0017] Some of the exemplary embodiments discussed herein provide improved memory consistency when undertaking program code conversion. Here, the controller unit is a translator unit. In particular, the inventors have developed mechanisms directed at program code conversion, which are useful in connection with a run-time translator that performs dynamic binary translation. That is, the controller unit is configured as a dynamic binary translator. For example, when undertaking such translation from a subject architecture with a first memory consistency model to a target architecture having another memory consistency model, the differences between the memory consistency models of the subject and target architectures may lead to unexpected or erroneous behaviour of the program code executed on the target architecture. In particular, memory consistency errors may arise when converting program code from a subject architecture having a strongly-ordered memory consistency model (such as SPARC and x86 architectures) to a target architecture having a memory consistency model with relatively weak ordering (such as in PowerPC and Itanium architectures). Hence, the example embodiments provide low-cost and effective mechanisms which protect memory consistency in a multiprocessor computer system performing these forms of program code conversion.

[0018] For further information regarding program code conversion as may be employed in the example embodiments discussed herein, attention is directed to PCT publications WO2000/22521 entitled “Program Code Conversion”, WO2004/095264 entitled “Method and Apparatus for Performing Interpreter Optimizations during Program Code Conversion”, WO2004/097631 entitled “Improved Architecture for Generating Intermediate Representations for Program Code Conversion”, WO2005/006106 entitled “Method and Apparatus for Performing Adjustable Precision Exception Handling”, and WO2006/103395 entitled “Method and Apparatus for Precise Handling of Exceptions During Program Code Conversion”, which are all incorporated herein by reference.

[0019] The present invention also extends to a controller apparatus or translator apparatus arranged to perform any of the embodiments of the invention discussed herein, when used in cooperation with a multiprocessor computing system. Also, the present invention extends to computer-readable storage medium having recorded thereon instructions which when implemented by a multiprocessor computer system perform any of the methods defined herein.

[0020] At least some embodiments of the invention may be constructed, partially or wholly, using dedicated special-purpose hardware. Terms such as ‘component’, ‘module’ or ‘unit’ used herein may include, but are not limited to, a hardware device, such as a Field Programmable Gate Array (FPGA) or Application Specific Integrated Circuit (ASIC), which performs certain tasks. Alternatively, elements of the invention may be configured to reside on an addressable storage medium and be configured to execute on one or more processors. Thus, functional elements of the invention may in some embodiments include, by way of example, components, such as software components, object-oriented software components, class components and task components, processes, functions, attributes, procedures, subroutines, segments of program code, drivers, firmware, microcode, circuitry, data, databases, data structures, tables, arrays, and variables. Further, although the preferred embodiments have been described with reference to the components, modules and units discussed below, such functional elements may be combined into fewer elements or separated into additional elements.

[0033] The following description is provided to enable a person skilled in the art to make and use the invention and sets forth the best modes contemplated by the inventors of carrying out their invention. Various modifications, however, will remain readily apparent to those skilled in the art, when considering the general principles of the present invention defined herein.

[0034] FIG. 1 gives an overview of a system and environment where the example embodiments of the present invention may find application, in order to introduce the components, modules and units that will be discussed in more detail below. Referring to FIG. 1, a subject program 17 is intended to execute on a subject computing system 1 having at least one subject processor 3. However, a target computing system 10 instead is used to execute the subject program 17, through a translator unit 19 which performs program code conversion. The translator unit 19 performs code conversion from the subject code 17 to target code 21, such that the target code 21 is executable on the target computing system 10.

[0035] As will be familiar to those skilled in the art, the subject processor 3 has a set of subject registers 5. A subject memory 8 holds, inter alia, the subject code 17 and a subject operating system 2. Similarly, the example target computing system 10 in FIG. 1 comprises at least one target processor 13 having a plurality of target registers 15, and a memory 18 to store a plurality of operational components including a target operating system 20, the subject code 17, the translator code 19, and the translated target code 21. The target computing system 10 is typically a microprocessor-based computer or other suitable computer.

[0036] In one embodiment, the translator code 19 is an emulator to translate subject code of a subject instruction set architecture (ISA) into translated target code of another ISA, with or without optimisations. In another embodiment, the translator 19 functions as an accelerator for translating subject code into target code, each of the same ISA, by performing program code optimisations.

[0037] The translator code 19 is suitably a compiled version of source code implementing the translator, and runs in conjunction with the operating system 20 on the target processor 13. It will be appreciated that the structure illustrated in FIG. 1 is exemplary only and that, for example, software, methods and processes according to embodiments of the invention may be implemented in code residing within or beneath an operating system 20. The subject code 17, translator code 19, operating system 20, and storage mechanisms of the memory 18 may be any of a wide variety of types, as known to those skilled in the art.

[0038] In the apparatus according to FIG. 1, program code conversion is performed dynamically, at run-time, to execute on the target architecture 10 while the target code 21 is running. That is, the translator 19 runs inline with the translated target code 21. Running the subject program 17 through the translator 19 involves two different types of code that execute in an interleaved manner: the translator code 19; and the target code 21. Hence, the target code 21 is generated by the translator code 19, throughout run-time, based on the stored subject code 17 of the program being translated.

[0039] In one embodiment, the translator unit 19 emulates relevant portions of the subject architecture 1 such as the subject processor 3 and particularly the subject registers 5, whilst actually executing the subject program 17 as target code 21 on the target processor 13. In the preferred embodiment, at least one global register store 27 is provided (also referred to as the subject register bank 27 or abstract register bank 27). In a multiprocessor environment, optionally more than one abstract register bank 27 is provided according to the architecture of the subject processor. A representation of a subject state is provided by components of the translator 19 and the target code 21. That is, the translator 19 stores the subject state in a variety of explicit programming language devices such as variables and/or objects. The translated target code 21, by comparison, provides subject processor state implicitly in the target registers 15 and in memory locations 18, which are manipulated by the target instructions of the target code 21. For example, a low-level representation of the global register store 27 is simply a region of allocated memory. In the source code of the translator 19, however, the global register store 27 is a data array or an object which can be accessed and manipulated at a higher level.

[0040] The term “basic block” will be familiar to those skilled in the art. A basic block is a section of code with exactly one entry point and exactly one exit point, which limits the block code to a single control path. For this reason, basic blocks are a useful fundamental unit of control flow. Suitably, the translator 19 divides the subject code 17 into a plurality of basic blocks, where each basic block is a sequential set of instructions between a first instruction at a single entry point and a last instruction at a single exit point (such as a jump, call or branch instruction). The translator 19 may select just one of these basic blocks (block mode) or select a group of the basic blocks (group block mode). A group block suitably comprises two or more basic blocks which are to be treated together as a single unit. Further, the translator may form iso-blocks representing the same basic block of subject code but under different entry conditions.

[0041] In the preferred embodiments, trees of Intermediate Representation (IR) are generated based on a subject instruction sequence, as part of the process of generating the target code 21 from the original subject program 17. IR trees are abstract representations of the expressions calculated and operations performed by the subject program. Later, the target code 21 is generated (“planted”) based on the IR trees. Collections of IR nodes are actually directed acyclic graphs (DAGs), but are referred to colloquially as “trees”.

[0042] As those skilled in the art may appreciate, in one embodiment the translator 19 is implemented using an object-oriented programming language such as C++. For example, an IR node is implemented as a C++ object, and references to other nodes are implemented as C++ references to the C++ objects corresponding to those other nodes. An IR tree is therefore implemented as a collection of IR node objects, containing various references to each other.

[0043] Further, in the embodiment under discussion, IR generation uses a set of register definitions which correspond to specific features of the subject architecture upon which the subject program 17 is intended to run. For example, there is a unique register definition for each physical register on the subject architecture (i.e., the subject registers 5 of FIG. 1). As such, register definitions in the translator may be implemented as a C++ object which contains a reference to an IR node object (i.e., an IR tree). The aggregate of all IR trees referred to by the set of register definitions is referred to as the working IR forest (“forest” because it contains multiple abstract register roots, each of which refers to an IR tree). These IR trees and other processes suitably form part of the translator 19.

[0044] FIG. 1 further shows native code 28 in the memory 18 of the target architecture 10. There is a distinction between the target code 21, which results from the run-time translation of the subject code 17, and the native code 28, which is written or compiled directly for the target architecture. In some embodiments, a native binding is implemented by the translator 19 when it detects that the subject program's flow of control enters a section of subject code 17, such as a subject library, for which a native version of the subject code exists. Rather than translating the subject code, the translator 19 instead causes the equivalent native code 28 to be executed on the target processor 13. In example embodiments, the translator 19 binds generated target code 21 to the native code 28 using a defined interface, such as native code or target code call stubs, as discussed in more detail in published PCT application WO2005/008478, the disclosure of which is incorporated herein by reference.

[0045] FIG. 2 illustrates the translator unit 19 in more detail when running on the target computing system 10. As discussed above, the front end of the translator 19 includes a decoder unit 191 which decodes a currently needed section of the subject program 17 to provide a plurality of subject code blocks 171a, 171b, 171c (which usually each contain one basic block of subject code), and may also provide decoder information 172 in relation to each subject block and the subject instructions contained therein which will assist the later operations of the translator 19. In some embodiments, an IR unit in the core 192 of the translator 19 produces an intermediate representation (IR) from the decoded subject instructions, and optimisations are opportunely performed in relation to the intermediate representation. An encoder 193 as part of the back end of the translator 19 generates (plants) target code 21 executable by the target processor 13. In this simplistic example, three target code blocks 211a-211c are generated to perform work on the target system 10 equivalent to executing the subject code blocks 171a-171c on the subject system 1. Also, the encoder 193 may generate control code 212 for some or all of the target code blocks 211a-211c which performs functions such as setting the environment in which the target block will operate and passing control back to the translator 19 where appropriate.

[0046] In some exemplary embodiments, the translator 19 is further arranged to identify system calls in the subject code 17. As discussed above, the target system 10 may use a different target operating system 20 and a different target ISA, and hence have a different set of system calls compared to the subject ISA. Here, in the translation phase, the decoder 191 is arranged to detect system calls of the subject ISA, where the subject code 17 calls the subject operating system 2. Most modern operating systems provide a library that sits between normal user-level programs and the rest of the operating system, usually the C library (libc) such as glibc or MS LibC. This C library handles the low-level details of passing information to the kernel of the operating system 2 and switching to a more privileged supervisor mode, as well as any data processing and preparation which does not need to be done in the privileged mode. On POSIX and similar systems, some popular example system calls are open, read, write, close, wait, execve, fork, and kill. Many modern operating systems have hundreds of system calls. For example, Linux almost has 300 different system calls and FreeBSD has about 330. Further, in some cases it is desired to maintain control of the target code and not pass execution control directly from the target code 21 to the target OS 20. In the exemplary embodiments, at least some of the system calls identified in the subject code 17 cause the target code 21 to be generated including function calls which call back into the translator 19, which will be termed herein “x_calls”. These x_calls appear to the target code 21 as if a system call had been made to the target OS 20, but actually return execution control from the target code 21 back into the translator 19. In the example embodiment, the translator 19 includes a target OS interface unit (also termed a “FUSE”) 194 which is called from the target code 21 by such x_calls. The FUSE 194 responds to the x_call, including performing actual system calls to the target OS 20 where appropriate, and then returns to the target code 21. Thus, the translator 19 effectively intercepts system calls made by the target code 21 and has the opportunity to monitor and control the system calls required by the target code 21, whilst the target code 21 still acts as if a system call had been made to the target OS 20.

[0047] As also shown in FIG. 2, in some exemplary embodiments the translator 19 is arranged to selectively intercept exception signals raised during execution of the target code 21. The translator 19 includes one or more exception handlers 195 that are registered with the target OS to receive at least some types of exception signals raised by execution of the target code 21. The exception handler 195 is thus able to selectively intervene where appropriate in handling the exception and inform the translator 19 that a certain exception has been raised. Here, the exception handler 195 either handles the exception and resumes execution as appropriate (e.g. returning to the target code 21), or determines to pass the exception signal to an appropriate native exception handler such as in the target OS 20. In one embodiment, the translator 19 provides a proxy signal handler (not shown) that receives selected exception signals and passes certain of the received exception signals to be handled by the appropriate exception handler 195.

[0048] FIG. 3 is a schematic diagram showing a program code conversion system according to an exemplary embodiment of the present invention.

[0049] Firstly, for illustration and ease of explanation, FIG. 3 shows a multiprocessor subject computing system 1 having a plurality of processors 3a, 3b which execute separate portions of subject code 170a, 170b (SC1 & SC2) and access data stored in a memory subsystem (MS) 8.

[0050] Most commonly, the subject code portions 170a, 170b executing on the processors 3a, 3b access the physical memory 8 by referring to an address space (VAS) 81 which maps memory access addresses referred to in the subject code 170a, 170b to physical memory addresses in the memory subsystem 8. Hence, the term virtual address space is used in the art to distinguish the code's address space from the physical addressing.

[0051] In some circumstances, the first and second subject code portions 170a, 170b are both intended to access the same region of the physical memory 8. In the example situation illustrated in FIG. 3, an area such as a page of the memory 8 is mapped in the virtual address space 81 by both the subject code portions 170a,170b. In other cases, an explicitly shared memory area is mapped into two different virtual address spaces.

[0052] As discussed above, a memory consistency model of the subject computing architecture 1 defines the semantics of memory accesses and the extent to which the processors 3a, 3b and the memory subsystem 8 may reorder memory accesses with respect to the original program order of the subject code 17. In this example, the subject architecture 1 has relatively strong ordering constraints. That is, the subject memory consistency model may define that consecutive stores and consecutive loads are ordered, but that a store followed by a load or a load followed by a store may be reordered compared to the program order. The memory consistency model in this example subject architecture can be briefly summarised in the following Table 1.

	TABLE 1
	First Instruction	Second Instruction	Constraint
	Store	Store	Ordered
	Store	Load	Not ordered
	Load	Store	Not ordered
	Load	Load	Ordered

[0053] The subject code 17 relies on the memory consistency model in order to function correctly. In practice, subject code is often written and debugged to the point at which it works on the currently available versions of the subject hardware. However, implementing the subject code 17 on a target computing system 10 as a different version of the subject computing system 1, or converting the subject code 17 to run on a totally different target computing system 10, can reveal weaknesses in the subject code. Here, there are many practical examples of multiprocessor systems which employ various different forms of relaxed memory consistency, including Alpha, AMD64, IA64, PA-RISC, POWER, SPARC, x86 and zSeries (IBM 360, 370, 390) amongst others.

[0054] As shown in FIG. 3, the translator unit (TU) 19 on the target computing system 10 converts the subject code 17 into target code 21 for execution on multiple target processors 13a, 13b with reference to the physical memory 18 of the target system. In this example, the target computing system 10 has a memory consistency model with weaker, more relaxed constraints than those of the subject system 1. For example, the target memory consistency model may specify that there is no ordering whatsoever and the target memory consistency model allows loads and stores to be freely reordered whilst maintaining program semantics, as summarised in the following Table 2.

	TABLE 2
	First Instruction	Second Instruction	Constraint
	Store	Store	Not ordered
	Store	Load	Not ordered
	Load	Store	Not ordered
	Load	Load	Not ordered

[0055] As will be familiar to those skilled in the art, the memory subsystem 18 may include various cache structures (not shown) which are designed to increase memory access speeds. The memory subsystem 18 may comprise two or more layers of physical memory including cache lines provided by on-chip or off-chip static RAM, a main memory in dynamic RAM, and a large-capacity disc storage, amongst others, which are managed by the memory subsystem according to the architecture of the subject computing system. There are many mechanisms to protect cache consistency (also termed cache coherency) to ensure that the cache structures remain consistent, but these are not particularly relevant to the examples under consideration and are not discussed further herein.

[0056] A simplified example will now be provided to illustrate some of the ways in which memory consistency errors may arise in the target computing system 10. In this example, two memory locations (*area1, *area2) are accessed. These locations are assumed to be on different memory pages to ensure that they are not on the same cache line within the cache structure of the target memory subsystem 18, and to increase the possibility that accesses to the memory 18 will happen out of order. Initially, we define the values stored in these locations as *area1=0 and *area2=0. The first processor 13a is executing a first portion of target code 21a which monitors the values stored in *area2 and then sets a variable “a” according to the value of *area1, as illustrated in the following pseudocode:

	while (*area2 == 0) { }
	int a = *area1

[0057] The second processor 13b executes a second portion of target code 21b which contains instructions that modify the values stored in the two memory locations:

	*area1 = 1
	*area2 = 1

[0058] Intuitively, we expect that the variable “a” should now be set to the value “1”. Indeed, in a strongly ordered sequentially consistent system, this would be true. However, a memory consistency error may arise such that the variable “a” is instead set to “0”. The error may arise for two typical reasons. Firstly, relaxed store ordering may allow the second store (*area2=1) to reach the memory before the first store (*area1=1). The first processor 13a is then able to read the old value of *area1. Secondly, relaxed load ordering allows loads to be issued out of order in the instruction pipeline within the first processor 13a, including loads that a speculatively executed. In this case, while the first processor 13a is waiting for *area2 to change, the value in *area1 is already speculatively loaded and will not be reloaded once the test succeeds. This means that even though the stores from the second processor 13b are correctly ordered, the first processor 13a can still read the updated values in a different order.

[0059] Most multiprocessor systems provide a safety net which enables the program code to override the relaxed memory consistency model of the hardware and impose stronger ordering constraints, thereby providing a measure of protection against memory consistency errors. One such safety net mechanism uses serialisation instructions in the target code 21a, 21b to form appropriate synchronisation points, whilst another such safety net is to safeguard certain areas of memory by setting attributes in a page table. These and other memory consistency protection mechanisms can be employed alone, or in combination, as will now be discussed in more detail.

[0060] Looking firstly at the use of serialisation instructions, one commonly available form is a fence instruction. The fence instruction forms a memory barrier which divides the program instructions into those which precede the fence and those which follow. Memory accesses caused by instructions that precede the fence are performed prior to memory accesses which are caused by instructions which follow the fence. Hence, the fence is useful in obtaining memory consistency, but incurs a significant performance penalty. The instruction SYNC in the IBM POWER Instruction Set Architecture is a prime example of a fence instruction. Other specific variations of the fence instruction are also available in the POWER ISA, such as a lightweight synchronisation (LWSYNC) instruction or Enforce In-order Execution of I/O (EIEIO) instruction. Other examples include MB and MBW from the Alpha ISA, MFENCE from the x86 ISA and MEMBAR from the SPARC ISA.

[0061] Some ISAs also provide one or more serialisation instructions which synchronise execution of instructions within a particular processor. That is, instruction synchronisation causes the processor to complete execution of all instructions prior to the synchronisation, and to discard the results of any instructions following the synchronisation which may have already begun execution. After the instruction synchronisation is executed, the subsequent instructions in the program may then begin execution. Here, the instruction ISYNC in the IBM POWER Instruction Set Architecture is a prime example of an instruction to perform such an instruction synchronisation.

[0062] These serialisation instructions are inserted into the target code to assert a memory consistency model which differs from the default memory consistency model of the target machine. Inserting these serialisation instructions into the example pseudo code discussed above results in modified target code 21a and 21b as follows.

[0063] For the first processor 13a, the serialisation instruction ISYNC is inserted (because of the Load-Load ordering specified in Table 1) so that the target code 21a becomes:

	while (*area2 == 0) { }
	isync
	int a = *area1

[0064] For the second processor 13b, the serialisation instruction SYNC is inserted so that the target code 21b becomes:

	*area1 = 1
	sync
	*area2 = 1

[0065] Turning now to another mechanism to provide protection against memory consistency errors, some target computing systems allow the manipulation of page table attributes. As a specific example, the IBM POWER architecture allows certain areas of the memory 18 to be designated as both caching-inhibited and guarded (hereafter called store-ordered). If separate store instructions access such a protected area of memory, the stores are performed in the order specified by the program. Conveniently, some pages of the memory are marked as store-ordered, whilst other pages of the memory are not store-ordered. The store-ordered pages may be used to assert a memory consistency model which differs from the default memory consistency model of the target machine. However, access to such store-ordered pages usually incurs a performance penalty compared with accesses to non store-ordered pages.

[0066] Example embodiments of the present invention will now be described in more detail to address the memory consistency issues outlined above, whilst minimising or avoiding altogether the heavy performance penalties associated with the memory consistency protection mechanisms such as serialisation instructions and store-ordered pages.

[0067] FIG. 4 is a schematic flow diagram which provides a general overview of a method applied in the exemplary embodiments of the present invention to protect memory consistency in a multiprocessor computing architecture, such as in the target computing system 10 discussed above with reference to FIG. 3.

[0068] Step 401 comprises executing at least a first target code portion 21a under a first memory consistency model. That is, in step 401 at least a first portion of target code 21a is executed in the target computing system 10 under a first default memory consistency model applicable to the architecture of the target computing system.

[0069] Step 402 comprises detecting a request to access a shared memory area that is accessible by both the first target code portion 21a and at least one other second program code portion 21b. That is, at step 402 the translator unit 19 is arranged to detect a memory access request with respect to a shared memory area in the memory 18 which is accessible (or will become accessible) to at least both of the first and second target code portions 21a, 21b. Various mechanisms are available to access such a shared memory area and various detection mechanisms are considered herein as will be discussed in more detail below.

[0070] Step 403 comprises applying a memory consistency protection mechanism such that at least certain instructions or certain groups of instructions in the first target code portion 21a execute under a protected second memory consistency model when accessing the detected shared memory area. Here, the translator unit 19 selectively applies a memory consistency protection mechanism which causes selected instructions within the first target code portion 21a to access the identified shared memory area in a manner which enforces a second memory consistency model which is different to the first model. In particular, the protected second memory consistency model provides stronger ordering constraints than the first model, aimed at preventing memory consistency errors of the type noted herein. Later, when the second code portion 21b also attempts to access the shared memory area, the memory consistency protection mechanism is further selectively applied such that at least selected instructions in the second program code portion 21b also now execute under the protected second memory consistency model in relation to the detected shared memory area.

[0071] In this exemplary embodiment, the first and second target code portions 21a, 21b are not initially restricted according to the second memory consistency model and instead execute initially under the default first model. That is, the target code is initially created and executed according to the higher-speed default memory consistency model of the target system. By applying the memory consistency protection mechanism only to those identified target code instructions which access those areas of memory 18 which have been detected as shared memory areas, the performance penalty incurred due to the restrictions and constraints of the second memory consistency model is substantially reduced compared with applying the enhanced second memory consistency model more generally across all memory accesses by the target code 21.

[0072] FIG. 5 is a schematic diagram showing selected parts of the target computing system 10 to further illustrate the exemplary embodiments of the present invention. In FIG. 5, the subject code 17 is a multithreaded application program which when translated into target code 21 executes as a plurality of target code portions (e.g. a plurality of threads). Three such target code portions 21a-21c (T1, T2, T3) are shown for illustration.

[0073] The subject code 17 is suitably an application program which is converted into target code 21 to execute on the target system 10 with the support of the translator 19. As general examples, the subject code 17 is a complex program such as a web server, a digital content server (e.g. a streaming audio or streaming video server), a word processor, a spreadsheet editor, a graphics image editing tool, or a database application. The target computing system 10 is often required to run many such applications simultaneously, in addition to other tasks such as those associated with the operating system 20 and the translator 19.

[0074] In the context of dynamic binary translation, the subject code 17 may take the form of a binary executable which has been created (e.g. compiled) specific to the particular subject architecture 1. Hence, there is no opportunity for human intervention or review of the subject code 17 and instead it is desired to automatically convert the subject code 17 into target code 21 (i.e. target binary) for execution on the target computing system 10. The mechanisms discussed herein will, in at least some embodiments, allow such a conversion process to be implemented automatically, whilst also protecting memory consistency.

[0075] Many commercially-available application programs execute as a plurality of processes and/or as a plurality of process threads. Here, although the exact implementation differs depending upon the specific computing architecture, each process generally has a relatively large amount of state information (also commonly termed context information) and has its own virtual address space. By contrast, a parent process may spawn one or more threads which usually share the state information of their parent process, and two threads from the same process will usually share the virtual address space of the parent process. Switching between threads from the same parent process is typically faster than context switching between processes, and multithreading is a popular programming and execution model on modern multiprocessor systems. Although the terms “process” and “thread” are widely used and understood by those skilled in the art, the description herein refers instead generally to a “portion” of program code for clarity.

[0076] As shown in FIG. 5, in addition to the units already described, the translator 19 of the exemplary embodiment further includes an address space allocation unit (ASAU) 196, a shared memory detection unit (SMDU) 197, and a memory protection unit (MPU) 198.

[0077] The ASAU 196 is arranged to allocate a plurality of virtual address space regions (VASR) 181 to the plurality of target code portions 21a, 21b, 21c. Secondly, the ASAU 196 is arranged to direct the generated target code portions 21a-21c to access different ones of the plurality of allocated VASRs 181.

[0078] The SMDU 197 is arranged to detect a request by one of the target code portions 21a, 21b, 21c to access a shared memory area, for which specific embodiments are discussed below, and identifies one or more target code instructions within this target code portion for which memory consistency protection is required.

[0079] The MPU 198 is arranged to apply memory consistency protection to the selected target code instructions identified by the SMDU 197. This memory consistency protection causes the target code to enforce a different memory consistency model, in this case with stronger ordering constraints, to preserve memory consistency and thereby maintain the memory consistency model demanded by the subject code 17. Suitably, the MPU 198 selectively applies serialisation instructions to the target code and/or selectively asserts store-ordered pages, as will be discussed in detail later.

[0080] In the example of FIG. 5, three target code portions T1, T2, T3 (21a-21c) are shown each associated with a respective virtual address space region 181a-181c. Further, in this first embodiment the ASAU 196 allocates an additional VASR 181d which is used in relation to shared memory areas.

[0081] In one example embodiment of the ASAU 196, the target computing system 10 provides a number of different addressing modes. Most commonly available computing systems provide a 32-bit virtual addressing mode such that the virtual address space of a particular portion of program code is able to address 232 individual elements (i.e. bytes, words) of the physical memory 18. Hence, many commercially available application programs expect to run in 32-bit virtual address spaces. However, some computing systems also allow larger addressing modes, such as a 64-bit mode, which can be used instead of or alongside the smaller 32-bit addressing mode. Conveniently, the translator unit 19 is set to run in the 64-bit addressing mode and is thereby provided with a 64-bit virtual address space (referred to below as the translator virtual address space or translator VAS 180). The address space allocation unit 196 then allocates a plurality of separate 32-bit virtual address space regions (VASR) 181 within the larger 64-bit translator VAS 180. Other addressing options are also available and can be applied in appropriate combinations to achieve the same effect, such as a 32-bit translator VAS which is subdivided to provide a plurality of 24-bit virtual address space regions.

[0082] The ASAU 196 is further arranged to direct each portion of target code 21 to a selected one or more of the VASR 181. As noted above with respect to FIG. 2, each portion of target code 21a is subdivided into a plurality of blocks 211 comprising a short sequence of individual instructions as a minimum unit handled by the translator 19. Some of these instructions make memory accesses such loads or stores and most of the instructions within a particular target code portion 21a access private memory with respect to the VASR 181a allocated to that portion. However, certain instructions or groups of instructions make memory accesses with respect to shared memory and are directed to access the VASR 181d for shared memory areas.

[0083] In one embodiment, the target code 21 is generated to refer to a base register BR 15a when performing memory operations. The base register 15a is a fast and readily available storage location for most architectures and can be used efficiently in “base plus offset” type memory accesses, but other suitable storage can be employed if appropriate. The base register BR is conveniently provided as part of the context information for this portion of target code (i.e. this thread or process). The base register BR 15a is used to store a base address giving a start address in the 64-bit translator VAS 180 as the start address of one of the 32-bit VASRs 181 to be used by the generated portion of target code 21. Each portion of target code 21a, 21b, 21c is then generated by the translator 19 to make memory accesses with reference to the start address in the base register BR 15a.

[0084] In the illustrated example of FIG. 5, for the target code portion 21a the base register BR contains the 64-bit value “1<<32,2³²” whereby the thread T1 makes memory accesses referring to its allocated first (32-bit) VASR 181a as an offset from this 64-bit base value. Similarly, for the second target code portion 21b the base register BR contains the value “2<<32,2³²” as the 64-bit start address of the second 32-bit VASR 181b.

[0085] Here, the example subject code 17 has been created to run in a 32-bit VAS and hence is concerned only with 32-bit addresses. The translator 19 accordingly generates the relevant portions of target code 21a-21b referring to 32-bit VASRs 181. However, since these 32-bit VASRs 181 are allocated from the larger 64-bit translator VAS 180, the target code uses the full 64-bit address when making memory accesses. This is achieved conveniently by concatenating a lower 32-bit address referring to the 32-bit VASR 181 with a full 64-bit base address specified in the base register BR 15a. For example, a target register r31 acts as the base register to hold the 64-bit base address and a target register r6 is used in the target code to hold a desired 32-bit address. The addresses are combined, as illustrated by the following pseudo code:

r6=0x00003210	;a 32-bit address in the target code VASR
r31=0x00000001 00000000	;a 64-bit base address for this VASR
add r3, r31, r6	;combine the addresses into r3
lwz r5, 0(r3)	;access memory using the combined
	address in r3

[0086] Further, the ASAU 196 is arranged to direct certain instructions within the target code portion 21a to refer to a different one of the allocated VASRs 181. In particular, certain instructions which concern accesses to shared memory are directed to the VASR 181d reserved for shared memory areas.

[0087] In one example implementation, the start address given in the base register BR 15a is modified, such that subsequent instructions in the target code 21 then refer to a different one of the allocated VASRs 181. That is, the base address stored in the base register BR 15a is modified and the modified base address is then employed by the one or more subsequent instructions in a particular block of the target code, until the base register is reset to the previous value. Here, as in the example above, the value originally given in the BR 15a is “1<<32,2³²” as the 64-bit start address of the VASR 181a allocated to the first target code portion 21a. Temporarily changing the base address to “0” would, in the illustrated example, now cause the target code instructions to instead refer to the fourth VASR 181d reserved for shared memory areas. Returning BR15a to the value “1<<32,2³²” again causes the target code 21a to refer to the allocated first VASR 181a.

[0088] Conveniently, the default base address in the base register 15a is set as part of the context/state for this portion of target code 21a. Thus, the default value is readily available from the context and can be quickly set to the default value when needed, such as at the beginning of each target code block 211.

[0089] In another example implementation, the ASAU 196 is arranged to selectively generate target code instructions referring to at least two base registers 15a, 15b as also shown in FIG. 5. Conveniently the first base register BR1 holds a base address of the VASR 181a-181c allocated to the current portion of target code 21a-21c. Meanwhile, the second base register BR2 holds a base address of the VASR 181d allocated for shared memory areas. Here, target code instructions are generated to perform memory accesses relating to the first base register BR1 or the second base register BR2, or a combination of both. Thus, generating the first portion of target code 21a to refer only to the first base register BR1 throughout causes this portion of target code to operate solely with respect to the respective allocated VASR 181a. However, where the target code instructions instead refer to the base address in register BR2, then the target code is directed to access the VASR 181d for shared memory areas. By selectively planting references to the first and second base registers BR1, BR2, the ASAU 196 is arranged to control which VASR is accessed by the target code.

[0090] The SMDU 197 is arranged to detect a request by one of the portions of target code 21a, 21b, 21c to access a shared memory area. Firstly, this request may take the form of a request to initialise an explicit shared memory area that is to be shared with other threads or processes. Secondly, the request may take the form of an implicit request relating to shared memory, such as a request to access a memory area which is already mapped in the virtual address space of another thread. The detection of explicit shared memory will be discussed first, referring to FIG. 6. Then, the detection of implicit shared memory will be discussed in more detail referring also to FIG. 7.

[0091] As discussed above, the translator 19 is arranged to monitor and intercept the system calls made by the executing target code 21. In particular, x_calls are provided to pass execution control to the FUSE 194 in the translator 19 and thereby emulate the behaviour of memory mapping system calls such as mmap( ).

[0092] If the x_call does not relate to shared memory, then suitably a system call is made to the target OS to take action as required, such as loading a private non-shared page into the VASR 181 allocated to the executing portion of target code. Execution control then returns to the target code via the FUSE 194, and the target code receives context as if returning from the target system call.

[0093] However, where the x_call relates to shared memory, then action is taken by the shared memory detection unit 197. Here, the x_call, or at least information derived from the x_call, is passed to the SMDU 197. As a specific example, the target operating system 20 supports memory mapping system calls such as shmget or mmap( ). As a particular example in UNIX and LINUX type operating systems, the mmap( ) system call typically takes the form mmap (start, length, prot, flags, fd, offset) to request a mapping of length bytes starting at offset offset from the file or other object specified by the file descriptor fd into virtual memory at address start. For an anonymous file the argument fd is null. The argument prot describes the desired memory protection that sets read and write protections. The parameter flags includes, amongst others, the flag MAP_SHARED which explicitly shares this mapping with all other processes that map this object. Alternatively, the parameter flags includes the flag MAP_PRIVATE which creates a private copy-on-write mapping. Hence, the mmap( ) system call is planted in the target code as an equivalent x_call (e.g. x_mmap( )) and is able to explicitly request a private memory area, in which case a corresponding mmap( ) system call is passed to target OS 20 as noted above, or explicitly request a shared memory area, whereby action is taken by the SMDU 197.

[0094] FIG. 6 is a more detailed schematic view of the target computing system shown in FIG. 5, to illustrate the actions taken by the SMDU 197 in relation to a request to map explicit shared memory. In particular, FIG. 6 is a schematic representation of part of the translator VAS 180.

[0095] In this example shown in FIG. 6, the currently executing portion of target code 21a is a thread T1 which contains an x_mmap( ) system-like function call to request an explicitly shared memory area 182a. However, the requested shared memory area 182a is not mapped into the virtual address space region 181a associated with this particular thread T1 21a. Rather, a memory area 182d of the same size and offset as the requested shared memory area 182a is mapped instead into the virtual address space region 181d reserved for shared memory. A pointer PTR to the requested shared memory area is returned to the T1 target code 21a by the FUSE 194 as expected behaviour following a mmap( ) system call. In this exemplary embodiment, a 32-bit pointer is returned as a start address in the 32-bit. VASR 181a. Execution of target thread T1 21a then continues as if a pointer had been given to a newly mapped shared memory area.

[0096] Optionally, the SMDU 197 records details of the requested shared memory area 182a derived from the arguments of the x_mmap( ) call. That is, the SMDU forms a mapping of each requested shared memory area 182, which conveniently includes the size and location of each shared memory area and may also identify a particular portion of target code as the owner or originator of this area. Also, the FUSE 194 and/or the SMDU 197 updates the subject state held in the translator 19 to reflect the manner in which this newly allocated shared memory region appears to the subject code 17.

[0097] Since the requested shared memory area 182a has not actually been mapped within the VASR 181a of the first target code thread T1 21a, when thread T1 attempts to access a page within the unmapped shared memory area 182a, an exception (i.e. a page fault) occurs. The exception is intercepted by the exception handler 195 as shown in FIG. 2 and passed to the SMDU 197, which thus is able to identify the block of target code that is attempting to access the explicit shared memory region 182a.

[0098] In response to this exception signal, the identified target code instruction is firstly directed to the VASR 181d reserved for shared memory and secondly the memory consistency protection mechanism is applied.

[0099] As discussed above, the ASAU 196 redirects at least certain instructions in the block of target code to the shared memory area 182d in the shared VASR 181d, by altering the code to amend the value in the base register BR 15a or by amending the code to refer instead to the second base register BR2 15b. The shared memory area 182d in the VASR 181d is mapped to the physical memory and thus the relevant instructions in the target code now obtain access to the shared memory area 182.

[0100] This exemplary embodiment readily enables the detection of an attempt to access the shared memory area 182 because the explicit shared memory area is not mapped within the virtual address space region 181 associated with the executing thread T1. However, by providing the additional virtual address space region 181d and redirecting selected target code instructions thereto, the desired shared memory region 182 is still accessible by the portion of target code 21.

[0101] Also, as will be discussed in more detail below, the MPU 198 applies the memory consistency protection mechanism to the identified target code instructions. That is, the memory consistency protection mechanism is applied selectively only for those blocks of target code 21 which attempt to access a shared memory region, to preserve memory consistency. Thus, relatively few instructions are affected. Notably, this mechanism does not need to apply the expensive memory protection mechanism to the whole program or even the whole thread.

[0102] Referring again to FIG. 5, it will be noted that the VASR 181d for shared memory areas does not overlap with the virtual address space region of any of the executing portions of target code T1, T2 or T3. Thus, any attempt by the second or third target code portions T2, T3 to access the explicitly shared memory area 182 will fail initially because the explicitly shared memory area is not mapped within the respective VASR 181b or 181c associated with that thread. Again, the resultant exception signal is handled by exception handler 195 and passed to the SMDU 197 which causes the relevant instructions to access the VASR 181d reserved for shared memory and have the memory consistency protection mechanism applied thereto. Hence, any target code instructions which attempt to access the explicit shared memory area are detected through the exception handler 195 and SMDU 197 and appropriate action is taken.

[0103] FIG. 7 is a more detailed schematic view of the target computing system shown in FIG. 5, to illustrate the actions taken by the SMDU 197 in relation to implicit shared memory. In particular, FIG. 7 is a schematic representation of part of the translator VAS 180 during the initiation of a new portion of target code, such as a new thread, to illustrate mechanisms to protect memory consistency when an implicit shared memory area is initiated at the beginning of a new portion of target code. In particular, FIG. 7 concerns a system call such as clone( ) in LINUX-type operating systems. Here, the normal system response is to create a child thread which runs concurrently with the parent process in the same shared virtual address space, where the child thread contains a subset of the context information from the parent process. Hence, a new thread created by a clone( ) system call will by default occupy the same virtual address space and thus share memory with a parent process. However, the response of the exemplary embodiments differs from this normal response as will now be described.

[0104] As shown in FIG. 7A, in this example a first thread T1 is executing in a first VASR 181a and has mapped in at least one memory area 182a as private to this process. Here, the mapped area 182a typically contains global data, initial heap memory and optionally additional heap memory. When the first thread T1 performs a clone( ) system call (conveniently planted as an x_call), a new thread T2 is allocated a separate VASR 181b using the ASAU 196 of FIG. 5. In this example, the base register 15a referenced by the new thread T2 21b contains the value “2<<32” such that the thread T2 is directed to the second VASR 181b. Since the two threads T1 and T2 are now allocated separate VASRs, the areas of memory 182a previously mapped in by thread T1 will not be mapped into the virtual address space region 181b associated with thread T2, as shown in FIG. 7B. Thus, an equivalent area 182b with a size and offset corresponding to the private mapped area 182a in VASR 181a remains unmapped in the second VASR 181b associated with thread T2.

[0105] As illustrated in FIG. 7B, thread T1 continues to access the private memory area 182a without, at this point, any changes to the portion of target code 21a of thread T1. This differs from the mechanism to handle explicit shared memory discussed above referring to FIGS. 5 and 6. Whilst thread T1 21a can still access the potentially shared memory area 182a, if thread T2 21b attempts to access the corresponding area 182b within its own VASR 181b the relevant pages are not mapped in and an exception will occur.

[0106] The exception signal is passed to the SMDU 197, which cooperates with the exception handler 195 to handle the exception. Firstly, the thread T1 is interrupted, because T1 owns the mapped in memory area 182a which the second thread T2 21b is attempting to access. Here, all pending accesses to the relevant memory area 182a from thread T1 are completed. Secondly, as shown in FIG. 7C, a corresponding memory area 182d of the same size and offset is now mapped in the shared VASR 181d such that the data in physical memory as referred to by the first thread T1 at area 182a is now available instead at the shared area 182d. A single page that faulted may be copied to the shared memory area 182d, or the entire relevant memory area 182a may now be copied. The shared area 182a which has been copied is now unmapped in the virtual address space region 181 of the first thread T1 21a such that the thread T1 can no longer access the area 182a, e.g. by using munmap( ) or by marking the area as protected.

[0107] T1 then notifies T2 that it is safe to retry the access in the newly created memory area 182d in the shared region 181d. T1 resumes normal execution. T2 now retries the memory access that faulted, this time by accessing the shared memory region 181d and with appropriate memory consistency protection applied, and then resumes execution.

[0108] As shown in FIG. 7C, if the target code portions T1 or T2 subsequently access the shared area 182 again (which is now non-accessible/unmapped in their private VASRs 181a, 181b), an exception will occur and the memory access will be completed instead through the exception handler 195 to access the shared address region 182d under the appropriate memory consistency protection applied by the MPU 198.

[0109] As a result of this mechanism, the appropriate instructions in the target code portions T1 and T2 are directed to the shared virtual address space region 181d to obtain access to the shared data area 182d and the stronger constraints of the second memory consistency model are applied to only for those parts of the target code which attempt to access the shared data area 182d.

[0110] The process now continues with threads T1 and T2 executing in parallel. Each time one of the threads, e.g. the second thread T2, attempts to access an area of memory which has already been mapped in by another thread, e.g. the first thread T1, an exception occurs which is handled to move the relevant area or page from the owner thread T1 into the shared VASR 181d and apply the memory consistency protection mechanism selectively to that area of target code. Any other thread which then attempts to access the now-shared memory area likewise causes an exception and the relevant code in that thread is likewise directed and subject to memory consistency protection. Thus, the mechanism applies to any number of portions of program code (threads T1, T2, T3 etc).

[0111] An alternative mechanism is to use a remapping system call as available in many Linux and UNIX type operating systems. Here, a MREMAP system call allows changes to a page table used by the target system 10 to control access to the memory 18. By changing the page table, a page of memory is mapped to a new position in the virtual address space 180 and is thus moved directly from the first VASR 181a to the second VASR 181b. The remapping occurs atomically from the point of view of the executing user-space threads and thus the first thread T1 does not need to be interrupted or notified.

[0112] FIG. 7D is an another view of the translator VAS 180 showing the plurality of address space regions 181, but here the VASR 181 are shown aligned at their respective base addresses for ease of illustration. Also, FIG. 7D shows a VASR map 199 held by the SMDU 197 which records the mapped areas within each of the VASRs 181. In this exemplary embodiment, the VASRs are all of equal 32-bit size and a single 32-bit map conveniently records the mapped memory areas within each VASR. Hence, even though privately mapped areas initially reside in the VASR for one of the target code portions, implicit shared memory is readily detected by consulting the map 199 to determine that the requested 32-bit address in a particular VASR is already mapped at the corresponding position in another VASR. In response, the actions illustrated in FIGS. 7B and 7C are performed only for the target code instructions which access the detected shared memory areas.

[0113] The exemplary embodiments discussed herein exactly one VASR 181 for each of the target code portions 21a-21c. However, other embodiments are also possible and are contemplated as variations on the described exemplary embodiments. For example, more than one shared area may be provided. In one alternate embodiment, each target code portion 21a-21c is associated with a corresponding private VASR holding only private memory areas, and a respective shared memory area to hold shared memory areas and also one or more private memory areas. Here, the use of multiple VASRs for the plurality of target code portions still allows shared memory, and particularly implicit shared memory, to be detected easily by the SMDU 197.

[0114] FIG. 8 shows the exemplary embodiment of the memory consistency protection mechanism in more detail.

[0115] The example of FIG. 8 shows a subject code block 171 and a corresponding target code block 211. At some point during execution of the target code block 211, an exception occurs in relation to a shared memory area and, as discussed above, action is taken by the exception handler 195 in cooperation with the ASAU 196, the SMDU 197 and the MPU 198 to protect memory consistency. In the example of FIG. 8, the exception arises in relation to instructions part way through execution of this block and hence the block 211 has been divided into two halves for illustration, where the top half represents the instructions that have already been executed whilst a remainder in the bottom half have not yet begun execution. Here, the memory protection mechanism firstly attempts to complete execution of the current block 211 and measures are taken on the fly to protect memory consistency. Afterwards, when an appropriate settled state has been achieved, longer-term changes are made to the target code such as regenerating the entire block 211 with the aim of avoiding exceptions in future executions of this block of target code.

[0116] Looking firstly at the immediate measures taken by the memory consistency protection mechanism, various example embodiments will be described.

[0117] In one example embodiment (marked by {circle around (1)} in FIG. 8), the target code 21 is generated to include null operations at appropriate synchronisation points, e.g. between each pair of stores. These null operations, such as the NOP instruction in the IBM POWER ISA, have no effects other than to cause the processor to do nothing for a particular number of clock cycles and are hence convenient to use as placeholders. The null operations are now replaced with active serialisation instructions (e.g. SYNC and ISYNC) to apply the memory consistency safety net to the target code. Also, the code is modified to refer to the shared VASR 181d as discussed above. This embodiment thus at least partially modifies the non-executed part of the block 211 ready for future executions.

[0118] In another embodiment (marked by {circle around (2)} in FIG. 8), execution of the block of target code is completed through a subject-to-target interpreter STInt 200 which resides within or is associated with the MPU 198. That is, execution is completed by interpreting the remaining instructions of the corresponding subject code block 171b instruction by instruction through the STInt 200 into equivalent target code instructions. Here, the MPU 198 causes the interpreter to apply serialisation instructions to form appropriate synchronisation points (e.g. inserting SYNC and ISYNC following loads or stores). However, this embodiment assumes that an appropriate subject state is available, in order to begin execution through the STInt 200.

[0119] In yet another embodiment, at least the unexecuted part of the target block is immediately regenerated to insert the serialisation instructions. That is, the remaining part of the target code block 211 is replaced by a modified version wherein serialisation instructions are inserted at the determined synchronisation points. Again, this embodiment assumes that a suitable subject state is available such that the regenerated target code may again move forward from a known state.

[0120] Where an appropriate subject state is not available at point where the exception occurred, the MPU 198 suitably rolls back in the target code to reach a checkpoint or recovery point at which the required subject state is achievable. An example mechanism to achieve subject state in relation to an exception is discussed in detail in WO2005/006106 cited above. Here, checkpoints are provided such as the beginning or end of a block or at selected points within a block. The MPU seeks the last reached checkpoint and is thus able to recover the subject state at that checkpoint. Execution of the block is now completed by going forward from the checkpoint with reference to the recovered subject state.

[0121] In a further refinement, the MPU 198 rolls forward to a next checkpoint subsequent to the point at which the exception occurred. Here, the MPU is assisted by a target-to-target interpreter TTInt 201 which interprets the already generated target code in the block 211 whilst inserting appropriate serialisation instructions to protect memory consistency, until the target code rolls forward to the next checkpoint. This forward rolling mechanism to recover subject state is discussed in detail in WO2006/103395. As a further refinement, the target-to-target interpreter TTInt 201 gathers translation hints during the roll-forward operation, such as recording those memory accesses which faulted and those which did not, in order to improve a later regeneration of that block of target code. Conveniently, these translation hints are implanted into the target code by initially generating the target code with NOP null operations and then selectively replacing the NOPs with translation hint flags.

[0122] Having dealt with the immediate needs of this target code block 211, the translator 19 may now devote further attention to the block 211. For example, all or part of the entire target block 211 is regenerated, such as to include the serialisation instructions (e.g. SYNCs and ISYNCs) throughout the block or to protect selected groups of instructions within the block. Thus, the regenerated target code block 211b is now subject to memory consistency protection in relation to shared memory accesses when that block is executed in future. The regeneration of the target code may employ translation hints gathered from execution of the previous incarnation of the block of target code. The regeneration can be performed immediately or can be deferred until a later point, such as when the block 211b is next needed for execution, by marking the block as requiring regeneration using a regeneration flag 211f as shown schematically in FIG. 8. The regeneration process may be iterative and take several passes. That is, the memory consistency protection is applied selectively to a first group of instructions after a first regeneration, and then is also applied to a second group of instructions in a second regeneration. Here, the translation hints gathered from the previous one or more incarnations may be used to assist the latest iteration of the regeneration. Further, the regeneration process may include the combination of two or more basic blocks of target code to from a group block having more than one unique entry point and/or more than one unique exit point and/or having internal jumps. Here, the translation hints embedded in the target code are helpful in allowing the translator to form an efficient group block which already takes account of the previous regenerations of the relevant basic blocks and so reduces regenerations of the group block.

[0123] In practical implementations, a particular section of code may be used to access both shared and private memory. As discussed above, the target code is originally generated appropriate to private memory in the relevant private VASR 181a-c. If the code is then retranslated appropriate to shared memory, it will now instead cause an exception when attempting to access private memory because the private memory is not mapped within the shared VASR 181d. One option is therefore to translate the code again back to the original format appropriate to private memory. The mutually exclusive nature of the memory pages being mapped either to the shared VASR 181d or the private VASR 181a-c ensures that this change of case is always detected.

[0124] There is an overhead in handling the exception and retranslating the relevant block or blocks of code. In some programs, the retranslation overhead is encountered relatively infrequently and hence is the most appropriate overall solution. However, it has also been found that some instances involve frequent retranslations, such as when a section of code is called from many different sites within a program. One particular example is the memory copy function memcpy( ). Here, the mechanism has been further developed and refined to address this issue.

[0125] As shown in FIG. 8, the translator 19 may retain at least two different versions of the target block 211. A first version 211A is the original translation without memory consistency protection, which executes quickly according to the reordering and other optimisations performed by the target system. The second version 211B is subject to the memory consistency protection, in this example referring to the shared VASR 181d with serialisation instructions, and hence executes more slowly. The translator may now selectively execute either the first or second version 211A or 211B when this block is next encountered during execution of the program. On entry to a function, a dynamic test is applied to determine the type of memory being accessed, i.e. either private or shared, and the appropriate version then selected. Whilst this solution reduces translation overhead, there is an execution penalty in performing the dynamic test.

[0126] In another refinement, the translator performs a loop optimisation. Here, a loop is executed for the first time and causes a memory exception because a memory access within the loop refers to shared memory. The translator may now retranslate the code in the loop to refer to shared memory, such that future executions referring to shared memory are less likely to fault. Providing a dynamic check specialises the code in the loop to either access private or shared memory. Also, the translator may attempt to hoist the dynamic check out of the loop and place it before the loop, thus further reducing execution workload.

[0127] As an alternative to dynamically checking the called code, another option is to inline the specialised code at the caller site. Another option is to specialise callers to a particular function. That is, a caller is specialised to call either private-type or shared-type accessor function to access private or shared memory respectively. For example:

	Caller > memcopy > memory

[0128] Becomes:

	Caller1(private) > memcopy_private > private memory
	Caller2(shared) > memcopy_shared > shared memory

[0129] These specialised callers may also involve further layer of indirection (i.e. wrapper functions as extra items on a call stack). Here, the memory address to be accessed is determined by the caller, and the memory address is only used by the accessor function (e.g. memcopy). The wrapper functions are initially set to call a private version of their successor. Hence, inspecting the call stack determines the wrapper functions which need to be specialised in order to allow future calls from this caller site to succeed. Suitably, progressive specialisation adapts one wrapper layer at a time, starting closest to the accessor function, until each layer has been specialised into private and shared versions.

[0130] FIG. 9 is a schematic flow diagram to provide a general overview of the memory consistency protection method as a summary of the various detailed embodiments discussed herein.

[0131] At step 901, first and second code portions are executed in separate virtual address space regions. For example, the first and second target code portions 21a, 21b execute with respect to distinct and non-overlapping first and second virtual address space regions 181a, 181b respectively.

[0132] Step 902 optionally comprises recording mapped areas 182a, 182b within each of the virtual address space regions 181a, 181b. Here, the address offset and size (address range) of each mapped memory area is recorded in a memory map 199 in response to a memory mapping action, such as a mmap( ) system call amongst others.

[0133] At step 903, the method comprises detecting an access request to a memory area which is unmapped in the address space associated with the currently executing code portion, but which is mapped in another of the plurality of address spaces. Here, the corresponding memory area is mapped either in the address space associated with another executing code portion (i.e. another thread) or in a separate address space reserved for shared memory. In either case, the access request by the currently executing code portion causes a memory exception and, in response to the memory exception, it is determined that the currently executing code portion is attempting to access a shared memory area.

[0134] At step 904, the method comprises amending the currently executing code to apply a memory consistency protection mechanism which causes the code to execute under a memory consistency model having predetermined constraints. Also, the currently executed code is amended to be directed to the predetermined shared memory area in the address space reserved for shared memory.

[0135] Finally, at step 905, where the shared memory area is not already residing within the address space reserved for shared memory, the shared memory area is moved into such address space and is unmapped or otherwise protected at least in the address space associated with the current code portion.

[0136] Considering mechanisms to initiate a new executing code portion such as the clone( ) system call discussed above, it will be appreciated that the step 901 may further include the steps of detecting such an attempt to initiate a newly executing code portion, allocating a separate address space for the new executing code portion and then executing the new code portion in the newly allocated separate address space.

[0137] It will also be appreciated that the steps illustrated in FIG. 9 need not be performed in the sequential order shown. As a particular example, it will be appreciated that the step 902 of recording the mapped areas in each address space may be performed dynamically as each new area of memory is mapped in to a particular address space, which will occur before, in parallel with, or after, the step 901 of executing the plurality of code portions each in separate address spaces. Further, the steps 904 and 905 may optionally be performed predictively, such that target code is first generated having the memory consistency protection mechanism applied thereto. These alternative implementations may depend upon settings within the translator 19. Where the translator predicts that, as a result of converting the subject code 17, such optional implementations would be beneficial for a particular section of the program, then the memory consistency protection mechanism is applied to the generated target code 21.

[0138] It will further be appreciated that the mechanisms discussed above are not limited to the processes and threads operating within a single application program, but may also be applied to a set or suite of programs operating simultaneously on the target computing system. That is, two or more separate programs (tasks) may operate together in a manner which shares memory under the mechanisms discussed above.

[0139] FIG. 10 is a schematic flow diagram of a method to implement the memory consistency protection mechanism in the MPU 198 according to another embodiment of the present invention. The memory consistency protection mechanism discussed above in detail applied serialisation instructions to the generated target code. In an alternative arrangement, a page flag modification is employed on certain architectures of the target computing system to create store-ordered pages in the memory 18.

[0140] In step 1001, the plurality of target code portions each execute in separate virtual address space regions, similar to the embodiment discussed above. At step 1002, the method comprises recording the memory areas mapped into each of the plurality of address spaces such as by using the VASR map 199 of FIG. 7D. These steps are suitably performed by the ASAU 196 of FIG. 5 in the manner discussed above.

[0141] At step 1003, the method comprises detecting a request to initiate a shared memory area. In one particular embodiment this request is a memory mapping system call such as mmap( ) which explicitly requests shared memory. In another example, an exception is raised when a child thread attempts to access a region which is unmapped in its own address space but which is mapped within the address space of a parent thread, where the child thread has been generated such as by a clone( ) system call. Suitably, the detection mechanisms of the SMDU 197 are employed as discussed above.

[0142] At step 1004, the page or pages of the detected shared memory area are marked by the MPU 198 by manipulating page table attributes such that accesses to these pages are forced to adhere to the second, non-default memory consistency model. As a specific example, an implementation of system hardware based on a PowerPC architecture is adapted to allow the relevant pages to be marked as requiring sequential consistency.

[0143] This embodiment advantageously does not require the shared memory area 182 to be moved to a separate address space region 181. Instead, the shared memory area 182 is mapped into the VASR 181a, 181b, 181c of each target code portion 21a, 21b, 21c which requires access to the shared memory area 182. Any code accessing the shared area will do so in a store-ordered manner and thus the desired memory consistency model is applied. Further, the target code will access the shared memory area 182 without a page fault and modification of the target code is avoided.

[0144] FIG. 11 is a schematic view of parts of the target computing system including the translator VAS 180 to further illustrate this example embodiment relating to store-ordered pages, together with a page table PT 183 which maps the virtual address space 180 to the physical memory subsystem 18.

[0145] In FIG. 11A, the first code portion T1 21a induces a mmap( ) type system call which explicitly requests shared memory, e.g. file-backed mmap_shared memory. The FUSE 194 in the translator unit 19 intercepts the system call and, if the page is not already marked as store ordered, invalidates cache lines for the region and marks the page as store-ordered in the page table PT 183. The file is then mapped into the VASR 181a of the first code portion T1 21a as a shared memory area 182a.

[0146] As shown in FIG. 11B, where a second target code portion 21b now attempts to access the shared memory area 182a, an exception will be raised because the shared memory area is not currently mapped in the relevant VASR 181b. In response, the SMDU 197 now maps the shared memory area 182b also into the second VASR 181b and, where not already so marked, marks the relevant memory pages as store-ordered by manipulating the page table attributes.

[0147] FIG. 11B also illustrates the response of the system if a clone( ) system call occurs. The new thread in code portion 21b is allocated a separate and distinct VASR 181b which does not overlap with the VASR 181a of the parent process 21a. In this case, a previously private memory region 182a in the first VASR 181a of the first code portion 21a may now become shared. Even though certain regions of memory 182a will be already mapped within the VASR 181a parent process, these remain unmapped for the newly cloned thread. If the second code portion 21b now attempts to access a memory region 182b which is unmapped in its own VASR 181b but which is mapped at a corresponding area 182a in the VASR 181a of the parent process 21a, then the child thread T2 21b will cause an exception. The SMDU 197 maps the desired file into the VASR of the child thread to map in the shared memory area 182b to the same relative position in both of these VASRs 181a, 181b to provide both portions of target code 21a, 21b access to the same page of the physical memory. In this case, the previously private but now implicitly shared memory area 182 is marked as store ordered in the page table PT 183.

[0148] The example embodiments have been discussed above mainly in relation to a program code conversion system for acceleration, emulation or translation of program code. Also, the mechanisms discussed herein are applicable to a debugging tool which detects, and optionally automatically corrects, program code that is vulnerable to memory consistency errors. Design problems or bugs are difficult to find, isolate and correct in shared memory multiprocessor architectures. Undetected bugs result in improper operations that often lead to system failures and that delay new software releases or even require post-release software updates. To this end, the controller/translator unit here is configured to run as a debugging tool to detect shared memory areas and apply appropriate code modifications to the subject code such as inserting serialisation instructions or modifying page table attributes, such that the generated target code is debugged.

[0149] Although a few example embodiments have been shown and described, it will be appreciated by those skilled in the art that various changes and modifications might be made without departing from the scope of the invention, as defined in the appended claims.

[0150] Attention is directed to all papers and documents which are filed concurrently with or previous to this specification in connection with this application and which are open to public inspection with this specification, and the contents of all such papers and documents are incorporated herein by reference.

[0151] All of the features disclosed in this specification (including any accompanying claims, abstract and drawings), and/or all of the steps of any method or process so disclosed, may be combined in any combination, except combinations where at least some of such features and/or steps are mutually exclusive.

[0152] Each feature disclosed in this specification (including any accompanying claims, abstract and drawings) may be replaced by alternative features serving the same, equivalent or similar purpose, unless expressly stated otherwise. Thus, unless expressly stated otherwise, each feature disclosed is one example only of a generic series of equivalent or similar features.

[0153] The invention is not restricted to the details of the foregoing embodiment(s). The invention extends to any novel one, or any novel combination, of the features disclosed in this specification (including any accompanying claims, abstract and drawings), or to any novel one, or any novel combination, of the steps of any method or process so disclosed.