Abstract:
The present invention relates generally to multiprocessor computer system, and particularly to a multiprocessor system designed to be highly scalable, using efficient cache coherence logic and methodologies. More specifically, the present invention is a system and method including a plurality of processor nodes configured to execute a cache coherence protocol that avoids the use of negative acknowledgment messages (NAKs) and ordering requirements on the underlying transaction-message interconnect/network and services most 3-hop transactions with only a single visit to the home node.

Description:
RELATED APPLICATIONS 
     This application is a provision of 60/210,675 filed Jun. 10, 2000. 
     This application is related to the following U.S. patent applications: 
     U.S. application Ser. No. 09/878,985 entitled System and Method for Daisy Chaining Cache Invalidation Requests in a Shared-memory Multiprocessor System, filed Jun. 11, 2001, and U.S. application Ser. No. 09/878,984. 
     Multiprocessor Cache Coherence System and Method in Which Processor Nodes and Input/Output Nodes Are Equal Participants, filed Jun. 11, 2001, and 
     U.S. application Ser. No. 09/878,983 entitled Cache Coherence Protocol Engine And Method For Processing Memory Transaction in Distinct Address Subsets During Interleaved Time Periods in a Multiprocessor System, filed Jun. 11, 2001. 
    
    
     The present invention relates generally to multiprocessor computer system, and particularly to a multiprocessor system designed to be highly scalable, using efficient cache coherence logic and methodologies. 
     BACKGROUND OF THE INVENTION 
     High-end microprocessor designs have become increasingly more complex during the past decade, with designers continuously pushing the limits of instruction-level parallelism and speculative out-of-order execution. While this trend has led to significant performance gains on target applications such as the SPEC benchmark, continuing along this path is becoming less viable due to substantial increases in development team sizes and design times. Such designs are especially ill suited for important commercial applications, such as on-line transaction processing (OLTP), which suffer from large memory stall times and exhibit little instruction-level parallelism. Given that commercial applications constitute by far the most important market for high-performance servers, the above trends emphasize the need to consider alternative processor designs that specifically target such workloads. Furthermore, more complex designs are yielding diminishing returns in performance even for applications such as SPEC. 
     Commercial workloads such as databases and Web applications have surpassed technical workloads to become the largest and fastest-growing market segment for high-performance servers. Commercial workloads, such as on-line transaction processing (OLTP), exhibit radically different computer resource usage and behavior than technical workloads. First, commercial workloads often lead to inefficient executions dominated by a large memory stall component. This behavior arises from large instruction and data footprints and high communication miss rates that are characteristic for such workloads. Second, multiple instruction issue and out-of-order execution provide only small gains for workloads such as OLTP due to the data-dependent nature of the computation and the lack of instruction-level parallelism. Third, commercial workloads do not have any use for the high-performance floating-point and multimedia functionality that is implemented in modern microprocessors. Therefore, it is not uncommon for a high-end microprocessor to stall most of the time while executing commercial workloads, which leads to a severe under-utilization of its parallel functional units and high-bandwidth memory system. Overall, the above trends further question the wisdom of pushing for more complex processor designs with wider issue and more speculative execution, especially if the server market is the target. 
     Fortunately, increasing chip densities and transistor counts provide architects with several alternatives for better tackling design complexities in general, and the needs of commercial workloads in particular. For example, the Alpha 21364 aggressively exploits semiconductor technology trends by including a scaled 1 GHz 21264 core, two levels of caches, memory controller, coherence hardware, and network router all on a single die. The tight coupling of these modules enables a more efficient and lower latency memory hierarchy that can substantially improve the performance of commercial workloads. Furthermore, the reuse of an existing high-performance processor core in designs such as the Alpha 21364 effectively addresses the design complexity issues and provides better time-to-market without sacrificing server performance. Higher transistor counts can also be used to exploit the inherent and explicit thread-level (or process-level) parallelism that is abundantly available in commercial workloads to better utilize on-chip resources. Such parallelism typically arises from relatively independent transactions or queries initiated by different clients, and has traditionally been used to hide I/O latency in such workloads. Previous studies have shown that techniques such as simultaneous multithreading (SMT) can provide a substantial performance boost for database workloads. In fact, the Alpha 21464 (the successor to the Alpha 21364) combines aggressive chip-level integration along with an eight-instruction-wide out-of-order processor with SMT support for four simultaneous threads. 
     Typical directory-based cache coherence protocols suffer from extra messages and protocol processing overheads for a number of protocol transactions. These problems are the result of various mechanisms used to deal with resolving races and deadlocks and the handling of “3-hop” transactions that involve a remote node in addition to the requester and the home node (where the directory resides). For example, negative-acknowledgment messages (NAKs) are common in several cache coherence protocols for dealing with races and resolving deadlocks, which occurs when two or more processors are unable to make progress because each requires a response from one or more of the others in order to do so. The use of NAKs also leads to non-elegant solutions for livelock, which occurs when two or more processors continuously change a state in response to changes in one or more of the others without making progress, and starvation, which occurs when a processor is unable to acquire resources. 
     Similarly, 3-hop transactions (e.g., requestor sends a request, home forwards request to owner, owner replies to requester) typically involve two visits to the home node (along with the corresponding extra messages to the home) in order to complete the transaction. At least one cache coherence protocol avoids the use of NAKs and services most 3-hop transactions with only a single visit to the home node. However, this cache coherence protocol places strict ordering requirements on the underlying transaction-message interconnect/network, which goes even beyond requiring point-to-point ordering. These strict ordering requirements are a problem because they make the design of the network more complex. It is much easier to design the routing layer if each packet can be treated independent of any other packet. Also, strict ordering leads to less than optimal use of the available network bandwidth. 
     The present invention also avoids the use of NAKs and services most 3-hop transactions with only a single visit to the home node. Exceptions include read transactions that require two visits to the home node because of a sharing write-back that is sent back to the home node. However, the present invention does not place ordering requirements on the underlying transaction-message interconnect/network. 
     SUMMARY OF THE INVENTION 
     In summary, the present invention is a system including a plurality of processor nodes configured to execute a cache coherence protocol that avoids the use of NAKs and ordering requirements on the underlying transaction-message interconnect/network and services most 3-hop transactions with only a single visit to the home node. Each node has access to a memory subsystem that stores a multiplicity of memory lines of information and a directory. Additionally, each node includes a memory cache for caching a multiplicity of memory lines of information stored in stored in a memory subsystem accessible to other nodes. Further, a protocol engine is included in each node to implement the negative acknowledgment free cache coherence protocol. The protocol engine itself includes a memory transaction array for storing an entry related to a memory transaction, which includes a memory transaction state. A memory transaction concerns a memory line of information and includes a series of protocol messages, which are routed both within a given node and to other nodes. Also included in the protocol engine is logic for processing memory transactions. This processing includes advancing the memory transaction when predefined criteria are satisfied (e.g., receipt of a protocol message) and storing an updated state of the memory transaction in the memory transaction array. 
    
    
     BRIEF DESCRIPTION OF THE DRAWINGS 
     Additional objects and features of the invention will be more readily apparent from the following detailed description and appended claims when taken in conjunction with the drawings, in which: 
     FIG. 1 is a block diagram of a multiprocessor system. 
     FIG. 2 is a block diagram of an input (I/O) node of the multiprocessor system of FIG.  1 . 
     FIG. 3 is a block diagram of a intra-chip switch and the module interfaces used to couple the modules of a system node to the intra-chip switch. 
     FIG. 4 depicts a directory data structure for keeping track of which nodes of the system have copies of each line of memory data. 
     FIG. 5 is a block diagram of a protocol engine. 
     FIG. 6A depicts the instruction format of the instructions executed in one embodiment of the protocol engine of FIG. 5; 
     FIG. 6B is a block diagram of a portion of the TSRF selection logic of the protocol engine of FIG. 5; and 
     FIG. 6C depicts a subset of the fields of each TSRF entry in the Transient State Register File (TSRF) of the protocol engine of FIG.  5 . 
     FIG. 7A is a table indicating operations performed during Even and Odd cycles of the execution unit of the protocol engine; 
     FIG. 7B depicts Even and Odd logical pipelines in the protocol engine that share use of many circuitry components; and 
     FIG. 7C depicts a state transition diagram for any single one of the TSRF entries in the Transient State Register File (TSRF) of the protocol engine of FIG.  5 . 
     FIG. 8 is a block diagram of a portion the execution logic of the protocol engine of FIG.  5 . 
     FIGS. 9A and 9B depict two embodiments of the Tag-State and Data arrays of an L1 cache. 
     FIG. 9C shows the architecture of the L1 cache in more detail. 
     FIGS. 10A and 10B depict the duplicate tag, tag-state and data arrays of an L2 cache. 
     FIG. 10C shows the architecture of the L2 cache in more detail. 
     FIGS. 11A,  11 B,  11 C,  11 D and  11 E illustrate the exchange of protocol messages in the course of a read request. 
     FIGS. 12A,  12 B,  12 C and  12 D illustrate the exchange of protocol messages in the course of a write request. 
     FIG. 13 illustrates the exchange of protocol messages in the course of completing a write-back request. 
     FIGS. 14A and 14B illustrate the exchange of protocol messages in the course of executing an invalidation request when nodes are represented in a limited-pointer format or a coarse-vector format. 
    
    
     DESCRIPTION OF THE PREFERRED EMBODIMENTS 
     All specific quantities (such as numbers of processors, number of nodes, memory sizes, bit sizes of data structures, operating speeds of components, number of interfaces, number of memory locations in buffers, numbers of cache lines), as well as the sizes and number of components in various data structures, disclosed in this document, are provided solely for purposes of explaining the operation of one particular embodiment. These quantities will typically vary, sometimes significantly, from one implementation of the invention to another. 
     The following is a list of abbreviations frequently used in the descriptions below: 
     CCP: cache coherence protocol; 
     FSM: finite state machine; 
     HPE: home protocol engine; 
     ICS: intra-chip switch; 
     I/O: input/output; 
     MC: memory controller; 
     PC: processor core; 
     RPE: remote protocol engine; and 
     TSRF: Transient State Register File. 
     Referring to FIG. 1, there is shown a multiprocessor system  100  including a multiplicity of processor nodes  102  and an I/O nodes  104 . Each processor node  102  is preferably implemented as a single chip multiprocessor. In a preferred embodiment, each processor node  102  has eight processor cores (PC)  106 ; however, other embodiments have two to sixteen PCs  106 . The PCs  106 , which may be comprised of a central processing unit, are processor cores since their caches, cache coherence logic and other supporting circuitry are shown separately. 
     Each processor core (PC)  106  is directly connected to dedicated instruction cache (iL1)  108  and data cache (dL1)  110  modules. These first-level caches (L1 cache modules)  108 ,  110  interface to other modules through an intra-chip switch (ICS)  112 . Also connected to the ICS  112  is a logically shared second level cache (L2)  114  that is interleaved into eight separate modules  116 , each with its own controller, on-chip tag, and data storage. Coupled to each L2 cache  116  is a memory controller (MC)  118  that preferably interfaces directly to a memory bank of DRAM (dynamic random access memory) chips (not shown) in a memory subsystem  123 . In a preferred embodiment, each memory bank provides a bandwidth of 1.6 GB/sec, leading to an aggregate bandwidth of 12.8 GB/sec. Also connected to the ICS  112  are two protocol engines, the Home Protocol Engine (HPE)  122  and the Remote Protocol Engine (RPE)  124 , which support shared memory across multiple nodes  102 ,  104  of the system. Multiple nodes are linked by a subsystem including a router (RT)  126 , an input queue (IQ)  128 , an output queue (OQ)  130 , a packet switch (PS)  132 , and a packet switched interconnect  134 . The router  136  sends and receives packets to and from other nodes via the interconnect  134 . The interconnect  134  physically links multiple nodes  102 ,  104 . In a preferred embodiment the total interconnect bandwidth (in/out) for each node is 32 GB/sec. Finally, a system control (SC) module  136  takes care of miscellaneous maintenance-related functions (e.g., system configuration, initialization, interrupt distribution, exception handling, performance monitoring). 
     In a preferred embodiment, the various modules communicate exclusively through the connections shown in FIG. 1, which also represent the actual signal connections. This modular approach leads to a strict hierarchical decomposition of the single chip used to implement each node of the system, which allows for the development of each module in relative isolation along with well defined transactional interfaces and clock domains. While each processor node  102  uses a complete multiprocessor system on a chip, the processor nodes  102  do not have any I/O capability in this embodiment. 
     Instead, I/O is performed by I/O nodes  104 , one of which is shown in FIG.  2 . Each I/O node  104  is preferably implemented as a single chip that is relatively small in area compared to the chip used to implement the processor nodes  102 . Each I/O node  104  is a stripped-down version of the processor node  102  having only one PC  106 , one L2 cache  116  and one memory controller module  118 . The router  140  on the I/O node  104  is a simplified version of router  126  having support for only two links instead of four, thus eliminating the need for a routing table. The I/O node  104  includes an I/O interface  142 , called the PCI/X interface in a preferred embodiment because it provides an interface between a PCI bus and an I/O bus  144 . 
     From the point of view of a programmer, the PC  106  on the I/O node  104  is indistinguishable from a PC  106  included on the processor node  102 . Similarly, memory at the I/O node  104  fully participates in the global cache coherence scheme of the multiprocessor system  100  (FIG.  1 ). The presence of a PC  106  on the I/O node  104  provides several benefits. For instance, it enables optimizations such as scheduling device drivers on this processor for lower latency access to I/O, or virtualization of the interface to various I/O devices (e.g., by having the PC  106  interpret accesses to virtual control registers). Except for the PCI/X interface  142 , most of the modules on the I/O node  104  are identical in design to those on the processor node  102 . For example, the same first-level data cache module (dL1)  110  that is used with the PCs  106  is also used to interface to the PCI/X module  142 . The dL1 module  110  also provides the PCI/X interface  142  with address translation, access to I/O space registers, and interrupt generation. The I/O node  104  may also be customized to support other I/O standards such as Fiber Channel and System I/O. 
     Referring back to FIG. 1, the multiprocessor system  100  in a preferred embodiment allows for glueless scaling up to 1023 nodes  102 ,  104 , with an arbitrary ratio of I/O nodes  104  to processing nodes  102 . The ratio of I/O nodes  104  to processor nodes  102  is adjustable to match the resource needs of any particular workload. Furthermore, the router  126 ,  140  in each of the nodes  102 ,  104  supports arbitrary network topologies and allows for dynamic reconfigurability. 
     The I/O nodes  104  of the system are treated the same as processor nodes  102 , that is, as full-fledged members of the multiprocessor system  100 . In part, this design decision is based on the observation that available inter-chip bandwidth is best invested in a single switching fabric that forms a global resource utilized for both memory and I/O traffic. 
     In an alternate embodiment, one or more of the I/O nodes  104  of the system have no processor cores and therefore no L1 caches other than the L1 cache for the interface  142  to an I/O bus or device. Furthermore, a first subset of the no-processor core versions of I/O nodes  104  may also lack a memory subsystem  123 , while other ones of the no-processor core versions of the I/O nodes do include a memory subsystem  123 . 
     Processor Core and First-Level Caches 
     In a preferred embodiment, the PC  106  uses a single-issue, in-order design capable of executing the Alpha instruction set. It consists of a 500 MHz pipelined datapath with hardware support for floating-point operations. The pipeline has 8 stages: instruction fetch, register-read, ALU 1 through 5, and write-back. The 5-stage ALU supports pipelined floating-point and multiply instructions. However, most instructions execute in a single cycle. The PC  106  includes several performance enhancing features including a branch target buffer, pre-compute logic for branch conditions, and a fully bypassed datapath. The PC  106  interfaces with separate first-level instruction and data caches designed for single-cycle latency. 
     As will be described in more detail below, the system uses 64 KB two-way set-associative, blocking caches with virtual indices and physical tags. The L1 cache modules  108 ,  110  include tag compare logic, instruction and data translation lookaside buffers (TLBs) (each storing 256 entries, in a 4-way associative caching arrangement), and a store buffer (data cache only). The L1 cache modules  108 ,  110  also maintains a 2-bit state field per cache line, corresponding to the four states in a typical MESI protocol. For simplicity, the L1 instruction cache modules  108  and L1 data cache modules  110  use virtually the same design. Therefore, unlike other Alpha implementations, the instruction cache is kept coherent by hardware. Treating all cache modules  108 ,  110  in the same way also simplifies the implementation of a no-inclusion policy at the L2 level. 
     Intra-Chip Switch 
     Referring to FIG. 3, conceptually, the ICS  112  is a crossbar that inter-connects most of the modules  150  on a processor node  102  or I/O node  104 . The ICS  112  includes a switch fabric  152  and an arbiter  154  for determining which data transfer(s) to handle during each available data transfer period. The length of the data period depends on the number of transfers required to send one cache line across the ICS  112 . In a preferred embodiment, each connection provided by the switch fabric  152  of the ICS  112  has a path width of 64 data bits, plus eight parity bits, for a total of 72 bits. Each cache line transported through the ICS  112  has 512 bits of data and sixty-four parity bits. Memory lines are transported along with the corresponding sixty-four parity bits when they are transported through the ICS  112 . Parity bits for memory lines are also sent to and used in the L1 cache arrays. However, parity bits are not used in the L2 cache and they are also not used in main memory. Instead, in the L2 cache, 20 ECC bits are associated with each memory line, and more specifically a 10-bit ECC is associated with each 256-bit half memory line. In the L2 cache and main memory, the 64 bits otherwise available for use as parity bits are used instead to store the 20 ECC bits, as well as a 44-bit directory entry, which will be described in more detail below. Data transfers generally are sent with a command or transaction type indicator, which is transferred in parallel with the first 64 bits of data of the cache line. Each cache line sized data transfer requires eight clock cycles, with 64 bits of data and a proportional share of the parity and ECC bits being transferred during each clock cycle. 
     Arbitration and flow control are handled by the arbiter  154 . To better understand the arbiter it is helpful to first review the interface  156  presented by each module  150  (i.e., L1 cache modules  108 ,  110 , L2 cache, protocol engine or system controller) to the ICS  112 . As shown in FIG. 3, the standard intra-chip interface  156  provided by each such module includes one or more input buffers  160 , one or more output buffers  162 , a first finite state machine (FSM)  164  for controlling use of the input buffer(s)  160 , and a second finite state machine (FSM)  166  for controlling use of the output buffer(s)  162 . The arbiter  154 , via the FSM  164 ,  166  of each module  150  keeps track of the availability of buffer space in the output buffers  162  of the modules  150  at all times, and exercises flow control by deferring requests to transfer data to modules with full input buffers  160 . The arbiter  154  also receives all intra-chip data transfer requests from the interfaces  156  of the modules  150 , and arbitrates between the requests whose destinations have input buffers  160  with sufficient room to receive a data transfer (i.e., a cache line of data). 
     In a preferred embodiment three parallel communication lanes, also called queues, are implemented in the input buffers  160  and output buffers  162  of the ICS interface  156 , as well as in the input and output buffers of interfaces (not shown) to the packet switch  126  and interconnect  134  (see FIG.  1 ). These lanes or queues are labeled I/O, low priority and high priority, respectively. The high priority queues in the input and output buffers are used to store messages sent from a home node to another node of the system, replies from third party nodes to the home node or the requester node for a particular transaction, and messages internal to a node. The low priority queues are used to store messages going to the home node for a particular transaction. The low priority message are thus messages for initiating new memory transactions, while the high priority messages are messages for completing previously initiated memory transactions. The I/O queues are used for handling requests being sent to I/O devices. The messages in the I/O queues are given the lowest priority by the intrachip switch  112  and also by the packet switch  126  and interconnect  134  (see FIG.  1 ). 
     The use of multiple communication lanes generally increases the size of the input and output buffers in the interfaces to the ICS  112 , packet switch  126  and interconnect  134 . However, the use of multiple communication lanes is important for avoid deadlock conditions in the network, and in particular for ensuring that active memory transactions make forward progress even when the system is experiencing high levels of protocol message traffic. In alternate embodiments, four or more communication lanes are used instead of three. In particular, in one alternate embodiment the high priority lane is replaced by two separate communication lanes, one for messages sent from the home node of a memory transaction and the other for replies sent by third parties to either the home node or any other node in the system. Providing the additional communication lane helps to ensure that messages sent by the home nodes of transactions are not blocked by reply messages being sent by the same node(s) for transactions in which those nodes are not the home node, and vice versa. 
     From a philosophical viewpoint, the ICS  112  is the primary facility for decomposing the processor node  102  and I/O node  104  into relatively independent, isolated modules  150 . For instance, the transactional nature of the ICS  112  and the uniformity of the interfaces  156  presented by the modules  150  to the ICS  112  together allow different types of modules  150  to have different numbers of internal pipeline stages for handling various type of memory transactions. 
     The ICS  112  uses a uni-directional, push-only data transfer technique. The initiator of a memory transaction always sources data. If the destination of a transaction is ready, the arbiter  154  schedules the data transfer according to datapath availability. A grant is issued by the arbiter  154  to the initiator of the transaction to commence the data transfer at a rate of one 64-bit word per cycle without any further flow control. Concurrently, the destination receives a signal from the arbiter  154  that identifies the initiator and the type of transfer. Transfers across the ICS  112  are atomic operations. 
     Each port to the ICS  112  consists of two independent 64-bit data paths (plus additional datapath bits for eight parity bits) for sending and receiving data. The ICS  112  supports back-to-back transfers without dead-cycles between transfers. In order to reduce latency, in a preferred embodiment the modules  150  are allowed to issue a “pre-request” indicating the target destination of a future request, ahead of the actual transfer request. The pre-request is used by the ICS  112  to pre-allocate data paths and to speculatively assert a grant signal to the requester. 
     Directory Used in Cache Coherence Protocol 
     Referring to FIG. 4, within each node of the system that has a memory subsystem  123 , a cache state directory  180  is maintained by the home protocol engine (HPE)  122 . The memory subsystem  123  of a node is also called the main memory array of the node. The directory  180  for a node&#39;s memory subsystem  123  includes one directory entry  182  for each “memory line”  184  in the memory system  123 . A “memory line” is the unit of memory that fits into one cache line of the L1 cache modules  108 ,  110  and L2 caches  114 . In a preferred embodiment, a memory line is 512 bits (64 bytes, or eight 64-bit words) of data; however, the size of the memory line will vary from one implementation to another. Each memory line  184  also includes two 10-bit ECC (error correction code) codes (one for each half memory line). The 20 bits of ECC codes and the 44-bit directory entry  182  occupy the same amount of memory, 64 bits, as would be required for one parity bit per byte. The ECC bits are used only in main memory and the L2 cache, to detect and correct errors in retrieved memory lines, while the directory entry is used by the home protocol engine (HPE)  122  to maintain cache coherence of the memory lines  184  corresponding to the directory entries  182 . 
     Each directory entry  182  includes a state field  186  for indicating the state of the corresponding memory line  184 , and a sharer-information field  188  for identifying nodes  102 ,  104  that have or might have a shared copy of the corresponding memory line  184 . A directory entry  182  in a preferred embodiment contains 44 bits, with the state field  186  comprising a 2-bit field that is repeated (i.e., stored twice in each directory entry  182 ) and the sharer-information field  188  comprising a 40-bit field that is split into two 20-bit fields  188 - 1 ,  188 - 2 . In a preferred embodiment there are two possible formats for the sharer-information field  188 , with the format of the sharer-information field  188  in a given directory entry  182  being determined by the number of nodes  102 ,  104  sharing the memory line  184  corresponding to the directory entry  182 . Generally, a node  102 ,  104  is said to “share” a memory line  184  if it maintains a read-only copy of the memory line  184 —typically stored in a cache array  108 ,  110 ,  114  within the respective node  102 ,  104 . 
     In a preferred embodiment (with a 40-bit sharer-information field and a maximum of 1023 nodes), when the number of nodes  102 ,  104  currently sharing a memory line  184  is four or less, a first sharer-information field  188  format called the “limited-pointer” format is used. In this format, the 40-bit sharer-information field  188  is divided into four 10-bit sub-fields, each of which is used to store a “direct node pointer” that identifies a node  102 ,  104  that is a sharer of the memory line  184 . A predefined null pointer value (e.g., 0×000 or 0×3FF) is stored in one or more of the 10-bit sub-fields to indicate that the respective 10-bit field does not identify a node  102 ,  104  (e.g., when fewer than four nodes  102 ,  104  share a memory line  184 ). More generally, the size of the sharer-information field  188  and the number of bits required for each direct node pointer determines the maximum number (DP) of direct node pointers that a sharer-information field  188  can store. Additionally, the node pointers (i.e., identifiers) included in the 10-bit sub-fields are obtained from requests to share a corresponding memory line of information  184 . Thus, each request to share a memory line of information  184  (described in detail below), includes a 10-bit identifier of the requesting node. 
     Also, in a preferred embodiment, when the number of nodes  102 ,  104  sharing a memory line  184  is more than four, a second sharer-information field  188  format called the “coarse vector” format is used. In this format, each bit in the sharer-information field  188  corresponds to one or more nodes  102 ,  104 . More specifically, when the number of nodes  102 ,  104  in the multiprocessor system  100  is more than four but less than forty-one, each bit of the sharer-information field  188  either corresponds to one node  102 ,  104  or does not correspond to any node  102 ,  104 . Thus, a set bit (zero or one depending on the specific implementation) in the sharer-information field  188  of a given directory entry  182  indicates that the one node  102 ,  104  corresponding to the set bit shares the memory line  184  corresponding to the directory entry  182 . And when the number of nodes  102 ,  104  in the multiprocessor system  100  is more than forty, one or more of the bits in the sharer-information field  188  correspond to a plurality of nodes  102 ,  104 . Thus, a set bit (zero or one depending on the specific implementation) in the sharer-information field  188  of a given directory entry  182  indicates that the one or more nodes  102 ,  104  corresponding to the set bit share the memory line  184  corresponding to the directory entry  182 . 
     Because only one bit is used to identify one or more nodes  102 ,  104  when the sharer-information field  188  is in the coarse-vector format, each node  102 ,  104  in the multiprocessor system  100  must be mapped to a bit in the sharer-information field  188 . The node to bit assignment table  189  of FIG. 4 illustrates a mapping of a plurality of nodes to a number of bits in a preferred embodiment (preferred embodiments of the invention do not actually utilize a table, which is included here merely for illustration). Specifically, table  189  shows 76 nodes  102 ,  104  mapped to respective bits in a 40-bit sharer-information field  188 . Each column in table  189  is associated with a bit in the sharer-information field  188 . Thus, according to table  189  the first bit in the sharer-information field  188  is associated with the node  102 ,  104  identified (and addressed) as 40. Since only 76 nodes  102 ,  104  are included in the multiprocessor system  100  of this example, table  189  includes only two rows. But if the number of nodes  102 ,  104  included in the multiprocessor system  100  in this example exceeded 79, 119, 159, etc., additional rows would be included in the table  189 . In other words, additional nodes  102 ,  104  would be associated with one or more of the bits in the sharer-information field  188 . 
     As indicated above, the numbers included in each entry of table  189  are node identifiers. The brackets around “0” is meant to indicate that 0 is not a valid node identifier in the embodiment illustrated in table  189 . In this embodiment, zero is used in the limited-pointer format to indicate that a particular sub-field of the sharer-information field  188  does not identify a node  102 ,  104 . To maintain consistency between the two formats, zero is not a valid node identifier in either format. 
     Determining the node identifiers for nodes  102 ,  104  associated with a given bit in sharer-information field  188  (which permits the home node  102 ,  104  to send out invalidation requests when a given sharer-information field  188  is in the coarse-vector format), is divided into two basic steps. Assuming that a given bit is set and associated with column 3 of table  189  (FIG.  4 ), the first node  102 ,  104  associated with this bit is simply the column number, i.e., 3. To calculate subsequent node identifiers of nodes  102 ,  104  associated with this bit, the system adds to the column number positive integer multiples of the number of bits included in the sharer-information field  188  to the column number. For example, for column three of the sharer-information field, the associated system nodes are  3 ,  43 ,  83  and so on. The second step (i.e., adding multiples of the number of bits in the sharer-information field  188 ) is continued until the calculated node identifier exceeds the total number of nodes  102 ,  104  in multiprocessor system  100 , in which case, the previously calculated node identifier is the identifier of the final node  102 ,  104  associated with a given bit. 
     As noted above, each directory entry  182  includes a state field  186 . In a preferred embodiment, the state field  186  is set to one of the following defined states: 
     invalid: indicates that the corresponding memory line  184  is not shared by another node  102 ,  104 ; 
     exclusive: indicates that a node  102 ,  104  has an exclusive copy of the corresponding memory line of information  184 , and thus may make changes to the memory line of information  184 ; 
     shared: indicates that the sharer-information field  188  is configured in the limited-pointer format described above and that the number of nodes having a non-exclusive (i.e., shared) copy of the corresponding memory line of information  184  is less than or equal to DP; 
     shared-cv: indicates that more than DP nodes  102 ,  104  have a non-exclusive (i.e., shared) copy of the corresponding memory line of information  184  and that the sharer-information field  188  is configured in the coarse vector format described above. 
     Protocol Engines 
     The basic architecture of each of the protocol engines  122 ,  124  (FIG. 1) is shown in FIG.  5 . The protocol engines are responsible for handling memory transactions, such as the sharing of cache lines, the exclusive assignment of a cache line to a processor in a particular node of the system, remote read and write operations. The protocol engines  122 ,  124  are responsible for maintaining cache coherence of cache lines among the nodes  102 ,  104  of the multiprocessor system  100 . 
     Each of the protocol engines  122 ,  124 , as shown in FIG. 5, includes an input controller  190 , preferably implemented as a finite state machine used in connection with a set of input buffers  192  for receiving data (inbound messages) from the ICS  112  and the PS  132 . Received messages, some of which include a full cache line of data and the associated parity bits, are stored in the input buffers  192 . In a preferred embodiment, sufficient input buffers  192  are provided to store inbound, received data for up to sixteen ongoing memory transactions. A test and execution unit  194  (herein called the execution unit) executes instructions obtained from an instruction memory  196 , also called the microcode array, so as to advance memory transactions, also called cache coherence transactions. The currently selected instruction, obtained from the instruction memory  196 , is held in a current instruction buffer  197  for decoding and execution by the execution unit  194 . Output messages generated by the execution unit  194  are stored in a output buffers  198 , the operation of which are controlled by an output controller  200 , preferably implemented as a finite state machine. The output messages are transferred from the output buffers  198  to specified destinations within the same node  102 ,  104  as a protocol engine  122 ,  124  via the ICS  112  or to specified destinations within other nodes  102 ,  104  of the multiprocessor system  100  via the PS  132 . 
     While the processor nodes  102  and I/O nodes  104  of a preferred embodiment use two protocol engines, including a home protocol engine (HPE)  122  (FIG. 1) for handling memory transactions where the node  102 ,  104  in which the protocol engine  122  resides is the home of the memory line that is the subject of the memory transaction, and a remote protocol engine (RPE) ( 124 , FIG. 1) for handling memory transactions where a remote node  102 ,  104  is the home of the memory line that is the subject of the memory transaction, for most purposes the two protocol engines  122 ,  124  maybe considered to be logically a single protocol engine. 
     FIG. 6A shows the format of each of the instructions stored in the instruction memory  196  and instruction buffer  197 . As shown, each instruction includes an operator, two operands, and a next program counter field. The operator indicates the type of operation to be performed by the execution unit  194  when executing the instruction, the two operands provide parameters that affect the execution of an instruction. 
     The current state of multiple memory transactions is stored in a set of registers collectively called the Transient State Register File (TSRF)  202 . Each memory transaction has a memory line address (sometimes called the global memory address) that identifies the memory line that is the subject of the memory transaction. More specifically, the memory line address identifies the node  102 ,  104  that interfaces with the memory subsystem  123  that stores the memory line of information  184  (i.e., home node) and a specific position within the memory subsystem  123  of the memory line of information  184 . In a preferred embodiment, the top M (e.g., 10) bits of the memory line address identify the home node  102 ,  104  of the memory line of information  184 , while the remainder of the address bits identify the memory line  184  within the identified node. In a preferred embodiment, the memory line address for a memory line does not include any of the address bits used to identify sub-portions of the memory line, such as individual 64-bit words of individual bytes within the memory line of information  184 . However, in other embodiments that support transactions on sub-portions of memory lines, the memory line addresses used may include bits for identifying such memory line sub-portions. 
     Referring to FIG. 6B, each memory transaction has a respective entry  210  stored in the Transient State Register File (TSRF)  202  that indicates the state of the memory transaction. In a preferred embodiment, the TSRF  202  has registers for storing sixteen entries  210  as well as access circuitry for reading and updating the contents of the TSRF entries  210 . Obviously the number of entries in the TSRF  202  is a design choice that will vary from one implementation to another. Typically, the TSRF  202  will include at least as many entries as the number of PCs  106  included in a processor node  102 . 
     Referring to FIG. 6B, the entries  210  of the TSRF  202  are divided into two groups—“even” TSRF entries  210  and “odd” TSRF entries  210 . The “even” TSRF entries  210  are used for memory transactions associated with memory lines of information  184  that have “even” memory line addresses (i.e., memory line addresses ending in a “0” bit), while the “odd” TSRF entries  210  are used for memory transactions associated with memory lines of information  184  that have “odd” memory line addresses (i.e., memory line addresses ending in a “1” bit). 
     Referring to FIGS. 6B,  7 A- 7 C, and  8 , the sequence of operations required to execute an instruction so as to advance a memory transaction is: reading the TSRF entries, scheduling one of the transactions represented by the TSRF entries, retrieving from the instruction memory the instruction identified by the TSRF of the scheduled transaction, and executing the instruction. As shown in FIGS. 7A and 7B, this sequence of four operations is pipelined and is furthermore performed by two “logical pipelines” that are parallel but offset from each other by one clock cycle. One logical pipeline is for the odd TSRF entries and the other is for the even TSRF entries. However, the two logical pipelines are implemented using a shared scheduler  212 , a shared microcode array  196  and access circuitry (see FIG.  8 ), and shared execute logic  240 , which along with the scheduler  212  is part of the test and execution unit  194 . Only the TSRF registers and access circuitry  202  have distinct even and odd circuits. 
     Alternating clock cycles of the test and execution unit  194  are called Even and Odd clock cycles. As shown in FIG. 7A, during each even clock cycle the following operations are performed, simultaneously, by the circuitry modules identified in FIG.  7 B: 
     reading the Odd TSRF entries, including comparing the address in each of the Odd TSRF entries with the addresses of messages received from the packet switch and intra-chip switch; 
     scheduling a next Even transaction (by selecting an Even TSRF entry) to be advanced by executing an instruction identified by the “next PC” field of one of the Even TSRF entries; 
     reading the microcode instruction identified by (A) the Odd transaction scheduled in the immediately previous Odd clock cycle and the condition code (CC) bits stored in the TSRF entry for the scheduled Odd transaction; and 
     executing the instruction for the currently scheduled Even transaction, where the instruction is identified by the “next PC” field of the Even transaction selected by the scheduler two clock cycles ago as well as the condition code bits stored in the TSRF of the currently scheduled transaction. 
     Similarly, as shown in FIG. 7A, during each Odd clock cycle the following operations are performed, simultaneously, by the circuitry modules identified in FIG.  7 B: 
     reading the Even TSRF entries, including comparing the address in each of the Even TSRF entries with the addresses of messages received from the packet switch and intra-chip switch; 
     scheduling a next Odd transaction (by selecting an Odd TSRF entry) to be advanced by executing an instruction identified by the “next PC” field of one of the Odd TSRF entries; 
     reading the microcode instruction identified by (A) the Even transaction scheduled in the immediately previous Even clock cycle and the condition code (CC) bits stored in the TSRF entry for the scheduled Even transaction; and 
     executing the instruction for the currently scheduled Odd transaction, where the instruction is identified by the “next PC” field of the Odd transaction selected by the scheduler two clock cycles ago as well as the condition code bits stored in the TSRF of the currently scheduled transaction. 
     The scheduler  212  selects the next Even (or Odd) transaction at the same time that the current Even (or Odd) transaction is being executed. In some circumstances, it is important for the current transaction to remain active and to be executed during two or more successive even clock cycles. For example, this is the case when a transaction needs to send two or more messages to other nodes in the system. The scheduler is able to determine whether the current Even (or Odd) transaction should be scheduled to execute again during the next Even (or Odd) clock cycle by inspecting the state, counters and condition codes in the TSRF of the currently executing transaction to determine if they satisfy predefined criteria for continuing execution of the current transaction for an additional execution cycle. 
     By interleaving instruction fetch and instruction execute cycles, the bandwidth and computational resources of the test and execution unit  194  and the microcode memory  196  are fully utilized. 
     As shown in FIG. 6B, the test and execution unit  194  (FIG. 5) of the protocol engine includes a scheduler  212  that selects an even TSRF entry  210  and an odd TSRF entry  210 , corresponding to the next even memory transaction and the next odd memory transaction to be processed or advanced by the execution unit  194 . The selections by the scheduler  212  are conveyed to a pair of multiplexers  214 ,  215  that transfer information from selected even and odd TSRF entries  210  to a pair of latches  216 ,  217  for storing the state of the currently running memory transactions. The TSRF entries stored in latches  216 ,  217  are used by the execution logic  242  (FIG. 8) of the execute unit  194  (FIG.  5 ). 
     Referring to FIG. 6C, each TSRF entry  210  includes many fields, a small subset of which are identified and described below: 
     a state field  220 : indicates the state of the associated memory transaction if any; 
     an address field  222 : stores the memory line address associated with a memory transaction if any; 
     a next program counter field  224 : identifies the next instruction to be executed by the execution unit when certain preconditions required for continued execution of the memory transaction are satisfied; and 
     a set of counter fields  226 : are used to store count values that, for example, control repeated execution of an instruction (e.g., when a transaction needs to send out N identical protocol messages to other nodes  102 ,  104 , one of the counter fields  226  is initially to a value corresponding to N, and is then decremented or incremented after each execution of the instruction until a predefined terminal count value is reached, at which point the memory transaction is either complete or a next program counter for the transaction is determined). The counter fields  226  and the state field  220  together form an overall or more specific state of an associated memory transaction. 
     In a preferred embodiment, the set of defined states for the state field  220  include: 
     vacant (also called invalid): indicates that the TSRF entry  210  does not store information related to a memory transaction; 
     active: indicates that the associated memory transaction is available for scheduling/execution; 
     running: indicates that the associated memory transaction is currently running (i.e., is currently being executed by the execution unit  194 , or was the transaction for which an instruction was executed during the last available even or odd execution cycle); 
     waiting: indicates that the associated memory transaction is stalled/deferred, waiting for a protocol message from another node  102 ,  104  to be delivered via the PS  132 ; 
     local_waiting: indicates that the associated memory transaction is stalled, waiting for a protocol message from within the same node  102 ,  104  to be delivered via the ICS  112 ; and 
     suspended: indicates that the associated memory transaction is suspended because there is a memory address conflict with a previously allocated memory transaction having the same memory line address. 
     FIG. 7C shows all defined state transitions for each of the TSRF entries  210 . A Vacant TSRF entry  210  becomes Active when a message initiating a new memory transaction is received and there is no unfinished transaction having the same memory line address and that blocks activation of the new memory transaction. A Vacant TSRF entry  210  becomes Suspended when a message initiating a new memory transaction is received and there is unfinished memory transaction having the same memory line address that blocks activation of the new memory transaction. 
     When an Active transaction is scheduled for execution it enters the Running state. If the execution of the transaction completes the transaction, the TSRF returns to the Vacant state. The Running Transaction remains in the Running state until it was sent all the protocol messages required for handling a current portion of the transaction. If execution of the transaction does not complete the transaction, the state of the TSRF becomes Waiting if the transaction is waiting for one or more messages from one or more other nodes to be able to continue the transaction, and becomes Local_Waiting if the transaction is waiting only for one or more messages from the local node to be able to continue the transaction. 
     The scheduler  212  includes arbitration logic for selecting the next even TSRF entry and the next odd TSRF entry to be sent to the execution unit  194  in accordance with (A) the states of the TSRF entries, (B) the buffered received messages received via the PS  132  and the ICS  112  and which TSRF entry, if any, corresponds to each of the buffered received messages, and (C) a set of prioritization rules. Each TSRF entry and each buffered received message identifies the memory line associated therewith, and the arbitration logic of the scheduler includes an array of comparators for comparing the memory line addresses in the TSRF entries with the memory line addresses in the buffered received messages so as to produce a corresponding set of status update signals. The status update signals are used for “upgrading” TSRF entries from the Waiting and Local_Waiting state to the active state, as well as for downgrading the TSRF entry for the last running transaction to the waiting, local waiting or vacant state, depending on whether the transaction is finished, and if not finished, what type of message (i.e., from the local node or a remote note) the transaction needs to receive in order to ready to resume execution. 
     The status update signals are also used to determine when a buffered received message has the same address as a previously allocated TSRF, but is for a different memory transaction. When this condition is detected by the arbitration logic, one of three actions is performed: (A) a new TSRF entry is allocated for the transaction associated with the received message, and the new transaction is suspended, (B) the received message is merged into previously allocated transaction and modifies its state, or (C) the message is temporarily left in the input buffer because the previously allocated transaction is not currently in a state allowing the received message to be merged with it, and the received message is then either merged with the previously allocated transaction or, if that transaction completes, a new TSRF is allocated for the new message and that TSRF is placed in the Active state. When the received message is of the type that could potentially be merged with a previously allocated transaction, the previously allocated transaction must be in the Waiting or Local_Waiting state before the merger can be performed. When a Receive instruction is executed, the transaction enters a Waiting or Local_Waiting state. The transaction can not enter the Active state until either (A) one of the predefined messages required to advance the transaction, or (B) one of the predefined messages that can be merged with the transaction is received. 
     Referring to FIGS. 6B and 8, the scheduler  212  selects between continued execution of the currently Running transaction and any of the other Active transactions, if any. FIG. 6B shows a portion of the logic for selecting an Active transaction. FIG. 8 shows logic for continuing execution of a currently Running transaction. On the right side of FIG. 8 is shown a current instruction buffer  197  for holding the current instruction for Running transaction. 
     The operator and arguments of the current instruction are passed to the execute logic  242 , which also has access to all the fields of the TSRF of the Running transaction. The execute logic computes a set of condition codes, labeled “Curr_CC” in FIG. 8, as well as new State and Next PC for the TSRF of the running transaction. The Next PC, to be stored in the TSRF of the current Running transaction, is obtained from the current instruction stored in buffer  197 . The execute logic  242  may also update one or more counters in the TSRF of the current Running transaction as well as other fields of the TSRF. 
     When the scheduler  212  determines that the current Running transaction should continue to run, the next instruction for the transaction is determined as follows. The current instruction in buffer  197  includes a “Next PC” field that specifies the base address of a next instruction. Predefined bits (e.g., the four least significant bits) of the “Next PC” address are logically combined (by logic gate or gates  244 ) with the condition codes (Curr_CC) generated by the execute logic  242  so as to generate a microcode address that is stored in microcode address latch  246 . Multiplexers  248  and  250  are provided to facilitate selection between the current Running transaction and another Active transaction. Multiplexers  248  and  250  operate during both Even and Odd clock cycles so as to perform separate instruction retrieval operations during Even and Odd clock cycles (See FIG.  7 A). 
     When all the Even (or Odd) TSRF entries are in the Vacant state, meaning that there are no running, active or waiting Even (or Odd) memory transactions, there are no Even (or Odd) memory transactions for the scheduler to select for execution, and thus the corresponding logical pipeline is unused. More generally, when none of the Even (or Odd) TSRF entries are in the Running or Active state (see discussion of FIG.  6 C), meaning that there are no Even (or Odd) memory transactions that are ready to be processed by the execution unit of the protocol engine, the corresponding logical pipeline is unused. During the corresponding clock periods instructions are not fetched from the instruction memory and the test and execution unit remains dormant. 
     The operation of the protocol engine while handling various specific memory transactions will be described in more detail below. Additional aspects of the scheduler and execution logic will also be described in more detail below. 
     L1 Cache 
     Referring to FIG. 9A, for simplicity a direct mapped version of the L1 cache  260  will be explained before explaining the two-way set associative version, shown in FIG.  9 B. Each L1 cache  260 , whether it is a data or instruction cache (see FIG. 1) includes a data array  262  for storing cache lines, a tag array  264  and a state array  266 . Each entry  268  of the L1 cache  260  includes a cache line, a tag and a state value. The cache line consists of the data from one memory line, and in a preferred embodiment this consists of 64 bytes (512 bits) of data plus parity and ECC bits corresponding to the 64 bytes. 
     The tag of each entry  268  consists of the address bits required to uniquely identify the cache line, if any, stored in the entry. Each address used to access memory consists of a string of address bits, ABCD, where A, B, C and D each consist of different groups of the address bits. The D bits are used to identify specific words (or bits, or bytes, depending on the implementation) within the cache line. The B and C bits, herein called BC, identify the entry  268  into which the memory line at address ABC 0  is stored within the L1 cache. The BC bits are called the index or cache index of the address. The A bits comprise the tag of the cache line, which together with the cache index uniquely identify the memory line. The only reason for dividing the cache index bits, BC, into two groups is for purposes of explaining the embodiment shown in FIG.  9 B. 
     The state of each L1 cache entry  268  is represented by two bits, which for each cache line represent one of four predefined states: 
     invalid, which means that the cache entry  268  is empty, or that the data in it is invalid and should not be used; 
     shared, which means that other processors or other nodes in the system have non-exclusive copies of the same memory line as the one stored in the cache entry; 
     clean_exclusive, which means that this L1 cache has the only copy of the associated memory line, has been allocated exclusive use thereof, and that the value of the cache line has not been changed by the processor coupled to the L1 cache; and 
     dirty_exclusive, which means that this L1 cache has the only copy of the associated memory line, has been allocated exclusive use thereof, and that the value of the cache line has changed by the processor coupled to the L1 cache. 
     Referring to FIG. 9B, there is shown a two-way associative version of the L1 cache, which is a preferred implementation. Only the differences between the L1 caches of FIGS. 9B and 9A will be described. In particular, the set associative L1 cache  270  has the same number of entries  278  as the direct mapped L1 cache  260 , but in this version there are two cache lines mapped to each cache index instead of just one. As a result, there are only half as many cache index values, and therefore the cache index is represented by just the C bits of the ABCD address bits. In this embodiment of the L1 cache, the B address bit of each memory line address is included in the tag of the entry, and thus the tag array  274  is one bit wider in this embodiment than in the direct mapped L1 cache embodiment. If the L1 cache were a four-way associative cache, the tag array  274  would be two bits wider than in the direct mapped L1 cache embodiment. A two-way associative L1 cache is preferred over a direct mapped cache because it reduces cache evictions caused by cache index conflicts. 
     L1 Data Paths and Control Logic 
     FIG. 9C shows the data paths and primary components of the L1 cache  108 ,  110 . Some of the connections between the various finite state machines of the L1 cache and some of the connections between those finite state machines, the tag and state arrays  274 ,  266  and other components of the L1 cache  108 ,  110  that are described below are not shown in FIG. 9C in order to avoid undue cluttering of this figure. 
     The L1 cache receives data (PC_L1_data) and a virtual address (PC_vaddr) from the processor core coupled to the L1 cache. Other signals received by the L1 cache from the processor core are a read request signal (PC_RdRq), which signals that the processor core is requesting data from the L1 cache, and a write request (PC_WrRq), which signals that the processor is requesting to write data into the L1 cache. The signals sent by the L1 cache to the processor core include data output by the L1 cache (L1_PC_data), a replay signal (PC_replay) requiring the processor to retry the last request sent by the processor core to the L1 cache, and an inhibit signal (PC_inhibit) to inform the processor core to inhibit its memory accesses because the L1 cache is busy (e.g., servicing a cache miss). 
     The L1 cache receives data from and sends data to the L2 cache, main memory, and other devices via the intra-chip switch  112 . Received data is temporarily buffered by a data in buffer  310 , and data being sent elsewhere is output via an output finite state machine (Output FSM)  312 . The output buffer for sourcing data to the ICS  112  is called the Fwd/Evt buffer  366 . 
     Input logic  314  receives control signals sent via the ICS  112  and conveys those control signals to either a fill FSM  316  or a synonym FSM  318 . The fill FSM  316  controls the loading of a cache line received from the ICS  112  into the L1 cache data array  262 . The synonym FSM  318  controls the movement of a cache line from one L1 cache slot to another when the L2 cache instructs the L1 cache to do so. Multiplexer  320  routes cached data from a slot of the L1 cache data array  262  back to the data array input multiplexer  322  under the control of the synonym FSM  318 . Input and output staging buffers  321 ,  323  are preferably used in this data path, for instance to facilitate delivery of successive portions of the data in a cache line over the data path. 
     When the synonym FSM  318  is not active, multiplexer  320  sources data from the data input buffer  310  to the data array input multiplexer  322 . The movement of a cache line from one L1 cache slot to another is required when the cache line index derived from a virtual address does not match the physical location of a cache line in the L1 cache. A tag information input multiplexer  324  is also controlled by the synonym FSM  318  to enable tag information for the L1 tag array  274  to be sourced by synonym information from the synonym FSM  318  when the synonym FSM  318  is activated. When the synonym FSM  318  is not activated, the tag information input multiplexer  324  sources tag information for the L1 tag array  274  from the virtual address (PC_vaddr) provided by the processor core. 
     An inhibit FSM  330  receives signals from the fill FSM  316  and synonym FSM  318  when those finite state machines are activated and sources the PC_inhibit signal to the processor core while either of these finite state machines is servicing a cache fill or synonym cache line relocation operation. 
     When the processor core sends either a read or write request to the L1 cache, the processor core provides a virtual address, PC_vaddr. The virtual address and information derived from it, such as a valid tag match signal, are stored in a series of staging buffers  332 ,  334 ,  336 . Additional staging buffers, beyond those shown in FIG. 9C, may be required in some implementations. The virtual address is translated into a physical address (PA) by a translation lookaside buffer (TLB)  340  at the same time that a tag and state lookup is performed by the tag and state arrays  274 ,  266 . The resulting physical address and tag lookup results are stored in a second staging buffer  334  and are then conveyed to a tag checking circuit  342  that determines if there is a tag match for a valid cache line. The results of the tag check, which includes state information as well as tag match information and the virtual address being checked, are stored in yet another staging buffer  336 . The information in the staging buffer  336  is conveyed to a data write FSM  360  when a valid match is found, and is conveyed to the output FSM  312  when a cache miss is detected. The final staging buffer  336  also stores a “replay” signal, generated by the tag checking circuit  342 , and the replay signal is conveyed to the processor core to indicate whether the L1 read or write operation requested by the processor core must be resubmitted to the L1 cache after the PC_inhibit signal is deactivated. 
     When a data write is being performed, the write request signal (PC_WrRq) and the results of the tag lookup are used by a data write FSM  360  and a cache access Arbiter  362  to determine if (and when) the data sourced by the processor core is to be written into the L1 cache data array  262 . The data sourced by the processor core is buffered in a series of staging buffers  352 ,  354 ,  356  so that the data to be written is available at the data array input multiplexer  322  at the same time that the tag check results become available to the data write FSM  360 . The data write FSM  360  stalls the data pipeline  352 ,  354 ,  356  if the arbiter  362  determines that the L1 cache is not ready to store the sourced data into the L1 cache data array  262 . 
     When a data read is being performed, the read request signal (PC_RdRq) is received directly by the arbiter  362  and the virtual address is used to directly read a cache line in the data array  262  even before the results of the tag lookup and check are ready. The data read from the data array is temporarily buffered in staging buffer  321  and is discarded if a cache miss is detected. If the read data is being read in response to a processor core request and a cache hit is detected, the read data is sourced from the staging buffer  321  to the processor core via the data path labeled Array_Out Data (L1_PC_data). If the read data is being read in response to a request received via the ICS  112 , the read data is sourced from the staging buffer  321  to the Fwd/Evt buffer  366 , and from there it is conveyed to the output FSM  312  for transmission to the requesting device via the ICS  112 . 
     L2 Cache With Duplicate L1 Tags 
     Referring to FIG. 10A, the L2 cache includes a set of “duplicate L1 tag and state arrays”  280 . These “DTag” arrays  280  contain exact copies of the tag arrays of all the L1 caches in the same node as the L2 cache, and furthermore contain state information that is similar to, but not identical, to the state information in the L1 cache state arrays  266  (FIG.  9 A). Thus, each entry  288  of the DTag arrays  280  corresponds to exactly one of the L1 cache entries  268  in the L1 caches of the node. The relationship between the state information in the L1 cache, the state information in the DTag arrays  280  of the L2 cache, and the state information in the L2 cache (see FIG. 10B) is as follows: 
     
       
         
               
               
               
             
           
               
                   
               
               
                   
                   
                 Possible 
               
               
                 L1 state 
                 DTag-L1 state 
                 corresponding L2 states 
               
               
                   
               
             
             
               
                 invalid 
                 invalid 
                 invalid, clean, clean_nodex, 
               
               
                   
                   
                 dirty 
               
               
                 shared 
                 shared_clean 
                 invalid, clean, clean_nodex, 
               
               
                   
                   
                 dirty 
               
               
                   
                 shared_clean_owner 
                 invalid 
               
               
                   
                 shared_clean_owner_nodex 
                 invalid 
               
               
                   
                 shared_dirty 
                 invalid 
               
               
                 clean —   
                 exclusive 
                 invalid 
               
               
                 exclusive 
               
               
                 dirty —   
                   
                 invalid 
               
               
                 exclusive 
               
               
                   
               
             
          
         
       
     
     As shown in the above table, the L2 cache keeps additional information in the DTag arrays regarding the ownership of shared cache lines. For instance, the shared_clean_owner_nodex state for any particular cache line indicates that the cache line in the L1 cache has not been modified, and that this node is the exclusive owner of the cache line. The clean_nodex state in the L2 cache means the same thing. 
     An L1 cache line with a DTag state of exclusive, shared_dirty, shared_clean_owner or shared_clean_owner_nodex is the owner of the cache line. If the L2 cache has a valid copy of the cache line, it is the owner of the cache line, and the only possible DTag states for that cache line are invalid or shared_clean. An L1 cache always performs a write-back when it replaces a cache line of which it is the owner. The written back cache line is loaded into the L2 cache, possibly victimizing another L2 cache line. 
     The L1 cache owner of a cache line responds to other L1 misses on the same cache line. In this case the requester of the cache line become the new owner and the previous owner&#39;s DTag state for the cache line is changed to shared_clean. 
     If a cache line is present in a particular node, node-exclusive information is kept in either the L2 state of in the DTag state of the owner L1 cache. The L2 states clean_nodex and dirty, and the DTag states shared_clean_owner_nodex, shared_dirty and exclusive all indicate that the node is the only node in the system that is caching the identified memory line (i.e., identified by the tag and cache index of the cache line). In a preferred embodiment, dirty (i.e., modified) cache lines are never shared across nodes. Thus, if a node has cache line that has been modified with respect to the memory copy, no other node in the system can have a copy of the line. As a result, when a node requests a shared copy of a cache line that has been modified by another node, the memory transaction that satisfies the request will always write-back the modified data to memory. Within a single node, however, a preferred embodiment allows sharing of a modified cache line among the processor cores. In this case, the DTag state of the L1 owner is set to shared_dirty and any other sharers have their DTag state set to shared_clean. 
     Referring to FIG. 10B, the main L2 cache array  290  includes a data array  292  for storing cache lines, a tag array  294  and a state array  296 . The L2 cache array is preferably distributed across eight interleaved arrays, but for purposes of this explanation, the interleaved array structure is not shown, as it does not affect the logical organization and operation of the L2 cache. Each entry  298  of the L2 cache  260  includes a cache line, a tag and a state value. The cache line consists of the data from one memory line, and in a preferred embodiment this consists of 64 bytes (512 bits) of data plus parity and ECC bits corresponding to the 64 bytes. 
     The tag of each entry  268  consists of the address bits required to uniquely identify the cache line, if any, stored in the entry. Because the L2 cache is typically much larger than the L1 caches, a different subset of the address bits of a memory line address is used to identify the cache index and a different subset of the address bits is used as the tag compared with the address bits used for those purposes in the L1 caches. 
     The L2 cache line state value for each L2 cache entry is selected from among the following state values: 
     invalid, which means that the cache entry  268  is empty, or that the data in it is invalid and should not be used; 
     clean, which means that the value of the memory line has not been changed and is therefore the same as the copy in main memory, and furthermore means that copies of the cache line may be stored in (A) one or more of the L1 caches of the same node as the L2 cache and/or (B) the L1 or L2 caches in other nodes of the system, and that these copies are non-exclusive copies of the same memory line as the one stored in the L2 cache entry; 
     clean_nodex (clean node-exclusive), which means that the L2 cache has a clean copy of the associated memory line (i.e., the memory line has not been changed and is the same as the copy in main memory), and that there may be cached copies of this memory line in local L1 caches in the same node as the L2 cache, but there are no copies of the memory line in any other nodes of the system; and 
     dirty, which means that this L2 cache has the only copy of the associated memory line, and that the value of the cache line has been changed by one of the processor cores coupled to the L2 cache. 
     L2 Data Paths and Control Logic 
     FIG. 10C shows the data paths and primary components of the L2 cache  116 . As described earlier with respect to FIG. 3, the L2 cache has an interface to the intra-chip switch  112 . This interface includes one or more input buffers  160 , one or more output buffers  162 , an input finite state machine (In FSM)  164  for controlling use of the input buffer(s)  160 , and an output finite state machine (Out FSM)  166  for controlling use of the output buffer(s)  162 . Similarly, the L2 cache  116  has an interface to the memory controller  118  (see also FIG. 1) that includes one or more input buffers  400 , one or more output buffers  402  and a memory controller interface finite state machine (MC interface FSM)  404  for controlling the use of the MC interface input and output buffers  400 ,  402 . 
     A set of pending buffers  406  are used to store status information about memory transactions pending in the L2 cache. For instance, the pending buffers  406  keep track of requests made to the memory subsystem (see FIG. 1) via the memory controller  118 . A set of temporary data buffers  408  are used to temporarily store cache line data associated with pending memory transactions, including data being sourced to the L2 cache, data sourced from the L2 cache, and data transported through the L2 cache (i.e., from the memory subsystem  123  to the L1 cache). Data sent by the L2 cache in response to an L1 cache miss bypasses the temporary data buffers  408  and is sent via a bypass data path  410  so as to reduce latency when the L2 cache contains the data needed to satisfy a cache miss in an L1 cache (which is coupled to the L2 cache via the ICS  112 ). 
     The duplicate tag (DTag) arrays  280  and L2 tag and state arrays  294 ,  296  have been discussed above with reference to FIGS. 10A and 10B. Access to and updating of these arrays is handled by the main L2 finite state machine  412 . The main L2 FSM  412  includes DTag and tag lookup, DTag and tag checking, and DTag, tag and state updating logic. 
     When an L1 cache miss is serviced by the L2 cache  116 , and the L2 cache does not have a cached copy of the memory line required by the L1 cache, the request is forwarded to the memory subsystem  123  via the MC interface FSM  404 . The memory line of information provided by the reply from the memory subsystem  123  is not stored in the L2 cache  116 . Instead the memory line is sent directly to the L1 cache, bypassing the L2 data array  292 . More specifically, the reply from the memory subsystem is directed through multiplexer  414  to the Din2 input port of the temporary data buffers  408 . The reply is then output at the Dout1 port of the temporary data buffers  408  to the interface output buffer  162  via output multiplexer  416 . 
     When an L1 cache evicts a memory line from the L1 cache, the victim memory line is sent to the L2 cache for storage via the ICS  112  and the interface input buffer  160 . The victim memory line is received at the Din1 input port of the temporary data buffers  408  and temporarily stored therein. The victim memory line is then sent from the temporary data buffers  408  to the L2 data array  292 , via the Dout2 port of the temporary data buffers  408  and a staging buffer  418 , for storage in the L2 data array  292 . 
     When the L2 cache sources a memory line to an L1 cache, the memory line read from the L2 data array  292  is conveyed via bypass line  410  to output multiplexer  416 , and from there to the ICS interface output buffer  162 . The output FSM  166  handles the transfer of the memory line from the output buffer  162  to the ICS  112 , and from there it is sent to the L1 cache. 
     Duplicate tags (DTags) are used by the L2 cache to determine which L1 caches have cached copies of an identified memory line. The duplicate tags in the DTag arrays  280  are accessed by the main L2 FSM  412 , and information derived from the duplicate tags is used to send messages via the output FSM  166  to one or more of the L1 caches in the same node as the L2 cache, or to other components of the node. 
     Cache Coherence Protocol 
     The present invention includes a cache coherence protocol (CCP) that enables the sharing of memory lines of information  184  across multiple nodes  102 ,  104  without imposing protocol message ordering requirements or requiring negative acknowledgments (NAKs). Because invalidation NAKs are not used in this invention, the CCP includes an assumption that the various requests (e.g., read request) discussed below always succeed. Additionally, the CCP is invalidation based, so shared copies of a memory line of information  184  are invalidated when the memory line of information  184  is updated. 
     As noted above, memory transaction relates to a memory line of information. Completion of a memory transaction requires a plurality of protocol messages, which are generated in part by instructions. Preferred embodiments of the present invention use seven instruction types: SEND, RECEIVE, LSEND (to local node), LSEND_REC (combined send/receive to/from local node), TEST, SET, and MOVE. The actual protocol code is specified at a slightly higher level with symbolic arguments, and C-style code blocks. A sophisticated microcode assembler is used to do the appropriate translation and mapping to instruction memory  196 . 
     Typical memory transactions require only a few instructions at each node  102 ,  104  for completion. For example, a memory transaction including a read request of a memory line of information  184  stored in a memory subsystem interfaced with a remote node  102 ,  104  requires a total of four instructions at the requesting node  102 ,  104 : a SEND of the read request to the remote node  102 ,  104 ; a RECEIVE of the read reply; a TEST of the state of the memory transaction (e.g., state field  220  and counters field  226 ); and an LSEND that sends a protocol message based on the read reply to the PC  106  that initiated the memory transaction. The CCP supports read, read-exclusive, exclusive, and write-back request types. A number of other protocol messages are supported as well in order to implement the requests. 
     The request types are now discussed in greater detail. FIG. 11A illustrates steps executed to satisfy a read request for a memory line of information  184 . In a first step, a PC  106  issues the read request for the memory line of information  184  (step  1100 ). If the memory line of information  184  is stored locally (step  1102 -Yes), the state of the memory line of information  184  is checked by reference to a corresponding entry  182  in the directory  180  (step  1104 ). If the directory entry  182  does not indicate that a remote node  102 ,  104  has an exclusive copy of the memory line of information  184  (step  1106 -No), the memory line of information  184  is retrieved directly from the memory subsystem  123  (FIG. 11B, step  1108 ). 
     If the memory line of information  184  is not stored locally (step  1102 -No), the read request is routed to the RPE  124  (step  1110 ). The RPE  124  adds an entry  210  in the TSRF  202  (step  1112 ). The new entry  210  indicates that a read reply is required to advance the state of this memory transaction. The new entry  210  also indicates that until the read reply is received, incoming requests related to the memory line of information  184  are stalled, which means that a TSRF entry  210  is added to the TSRF  202  for the incoming requests. Once the read reply is received, the state of the TSRF entry  210  is updated by the RPE  124  so that these incoming requests are processed. 
     The RPE  124  then sends a read request to the home node (step  1114 ). The home node is the node  102 ,  104  to which the memory subsystem  123  storing the memory line of information  184  is interfaced. 
     The read request is received by the home node  102 ,  104 , and routed internally as described above to the HPE  122  (step  1116 ). The HPE  122  responds by adding an entry  210  in the TSRF  202  (step  1118 ) and checking the state of the memory line of information  184  in a corresponding entry  182  in the directory  180  (step  1120 ). If the entry  182  does not indicate that a node  102 ,  104  has an exclusive copy of the memory line of information  184  (FIG. 11C, step  1122 -No), the HPE  122  updates the entry  210  in the TSRF  202  so that it indicates that the memory transaction requires an internal response to a request for the memory line of information  184  in order to advance to another state (step  1124 ). The HPE  122  then submits an internal request for the memory line of information  184  from the memory subsystem  123  (step  1126 ). Upon receiving the memory line of information  184  (step  1128 ), the HPE  122  sends a read reply to the requesting node  102 ,  104  (step  1130 ), updates the state of the memory line of information (step  1131 ), and removes the TSRF entry  210  (step  1132 ). 
     As noted above, the state of the memory line of information  184  is embodied in a corresponding entry  182  in the directory  180 . Included in the entry  182  is a state field  186  and a sharer-information field  188 . If the state field  186  indicates that the state of the memory line of information is shared-cv, the HPE determines which bit in the bits of the sharer-information field  188  the requesting node  102 ,  104  is mapped to. If the bit is not already set to indicate that a node  102 ,  104  mapped to that bit is sharing a copy of the memory line of information  184 , the bit is so set. 
     If the state field  186  indicates that the state of the memory line of information is “shared”, the HPE  122  determines if the requesting node  102 ,  104  is already identified as sharing the memory line of information  184  in the sharer-information field  188 . If so, the sharer-information field  188  and state field  186  are not changed. If the requesting node  102 ,  104  is not already identified as sharing the memory line of information  184 , the HPE  122  determines if any of the sub-fields within the sharer-information field  188  is set to indicate that it does not identify a sharer node  102 ,  104  (e.g., set to zero). If such a field is found, the HPE  122  sets it to identify the requesting node  102 ,  104 . As noted above, the identity of the requesting node  102 ,  104  is included in the original request to share the memory line of information  184 . If no such sub-field within the sharer-information field  188  is set to indicate that it does not identify a sharer node  102 ,  104 , the HPE  122  must set the state field  186  to “shared-cv”. Additionally, the HPE  122  must identify and set the bits in the 40-bit sharer-information field associated with (A) the four nodes  102 ,  104  previously identified by the sharer-information field  188  and (B) the requesting node  102 ,  104 . The HPE  122  then removes the entry  210  from the TSRF  202  (step  1132 ). 
     If the entry  182  indicates that a node  102 ,  104  (i.e., owner node) has an exclusive copy of the memory line of information  184  (step  1122 -Yes), the HPE  122  updates the entry  210  in the TSRF  202  so that it indicates that the memory transaction requires a share write-back in order to advance to another state (FIG. 11D, step  1134 ). The state also indicates that any requests related to the memory line of information  184  received while the HPE  122  is waiting for the share write-back should be deferred (i.e., stalled) until after receipt of the share write-back. This is accomplished by adding a new entry  210  to the TSRF  202  for such requests, and setting the state of these new entries  210  to indicate that the associated memory transaction is eligible for processing once the share write-back is received. 
     The HPE  122  then sends a read forward to the owner node  102 ,  104  (step  1136 ). The read forward is received by the owner node  102 ,  104 , and routed to the RPE  124  (step  1138 ). The RPE  124  responds by adding an entry  210  in the TSRF  202  indicating that the memory transaction requires an internal response to a request for the memory line of information  184  in order to advance to another state (step  1140 ). The RPE  124  then sends an internal request for the memory line of information  184  from L1 or L2 cache  110 ,  114  (step  1141 ). Upon receiving the memory line of information  184  (step  1142 ), the RPE  124  sends a share write-back to the home node  102 ,  104  (FIG. 11E, step  1144 ) and a read reply to the requesting node  102 ,  104  (step  1146 ), both of these protocol messages include an up-to-date copy of the memory line of information  184 . The RPE  124  also removes the entry  210  from the TSRF  202  (step  1148 ). 
     Upon receiving the share write-back (step  1150 ), the HPE  122  updates a copy of the memory line of information  184  (either in the memory subsystem  123  initially or a local cache initially and the memory subsystem  123  subsequently) (step  1152 ). HPE  122  then updates the state of the memory line of information  184  in the directory  180  to indicate that both the requesting node  102 ,  104  and the former owner node  102 ,  104  are both storing a shared copy of the memory line of information  184  (step  1154 ). The HPE  122  also updates the state of any entries  210  in the TSRF  202  for a request relating to the memory line of information  184  and received while waiting for the share write-back to indicate that the associated memory transaction may be executed. The HPE  122  then removes the entry  210  in the TSRF  202  related to this memory transaction (step  1155 ). 
     Upon receiving the read response (whether sent by the home node  102 ,  104  or an owner node  102 ,  104 ) (step  1156 ), the RPE  124  forwards the shared copy of the memory line of information  184  to the PC  106  that initiated the memory transaction (step  1158 ). The RPE also removes the entry  210  in the TSRF  202  related to the memory transaction (step  1160 ). 
     The read request steps described above with reference to FIGS. 11A-11E are subject to an optimization in preferred embodiments of the present invention. Specifically, if the memory line of information requested by the requesting node  102 ,  104  is not shared or owned by any nodes  102 ,  104 , the HPE  122  returns an exclusive copy of the memory line of information  184 . In other words, the response to a request for a shared copy of the memory line of information  184  is “upgraded” from a read reply to a read-exclusive reply. Thus, the requesting node  102 ,  104  is identified in the directory  180  as exclusive owner of the memory line of information. However, this optimization does not affect the home node&#39;s  102 ,  104  response to a request for a memory line of information that is comprised of an instruction since an instruction is never written to by a requesting node. Thus, there is no reason to provide an exclusive copy. 
     FIG. 12A illustrates steps executed to satisfy a request for an exclusive copy of a specified memory line of information  184 , which permits the node  102 ,  104  requesting the memory line of information  184  (i.e., requesting node) to modify the memory line of information  184 . In a first step, a PC  106  issues the request for an exclusive copy of the memory line of information  184  (step  1200 ). The request is routed to the RPE  124  (step  1210 ), which adds an entry  210  in the TSRF  202  (step  1212 ). The new entry  210  indicates that a read-exclusive reply and a number (zero or more) of invalidation acknowledgments are required to advance the state of this memory transaction. The RPE  124  then sends a read-exclusive request to the home node (step  1214 ). At this point the memory transaction in the RPE  124  enters the Waiting state, where it remains until it receives the aforementioned read-exclusive reply and (zero or more) invalidation acknowledgments. When these messages are received by the RPE  124 , the memory transaction it will made Active and then Running in order to receive and process these protocol messages so as to advance and complete the memory transaction. The new entry  210  also indicates that until the aforementioned replies are received, incoming requests related to the memory line of information  184  are stalled, which means that a TSRF entry  210  is added to the TSRF  202  for the incoming requests. Once the aforementioned replies are received, the state of the TSRF entry  210  is updated by the RPE  124  so that these incoming requests are processed. 
     The read-exclusive request is received by the home node  102 ,  104 , and routed to the HPE  122  (step  1216 ) of the home node, which adds an entry  210  in the TSRF  202  (step  1218 ). The HPE  122  then checks the state of the specified memory line  184  in a corresponding entry  182  in the directory  180  (step  1220 ). At this time, the HPE also sends a request to the L2 cache to locate and invalidate any copies of the specified memory line that may be present on the home node. The L2 cache uses the information in its L2 tag array and DTag arrays to determine if any copies of the specified memory line are present in the L2 cache or any of the L1 caches in the home node. If a copy of the specified memory line is found in the L2 cache, it is invalidated by the L2 cache, and if a search of the DTag arrays locates any copies of the specified memory line in the home node&#39;s L1 caches a command message is sent by the L2 cache to the identified local L1 cache or caches instructing those L1 caches to invalidate their copies of the specified memory line. Each L1 cache that receives the invalidate command respond to this command by setting the state of the corresponding cache line to “invalid”. It should be noted that when the requester for exclusive ownership of the specified memory line is a processor core in the home node of the memory line, L2 cache invalidates all cached copies of the specified memory line except for the copy (if any) held by the L1 cache of the requesting processor. 
     If the directory entry  182  for the specified memory line does not indicate that a node  102 ,  104  has an exclusive copy of the memory line of information  184  (FIG. 12B, step  1222 -No), the HPE  122  updates the entry  210  in the TSRF  202  to indicate that the memory transaction requires an internal response to a request for the memory line of information  184  in order to advance to another state (step  1224 ). The HPE  122  then sends a request for the memory line of information  184  from the memory subsystem  123  (step  1226 ). Upon receiving the memory line of information  184  (step  1228 ), the HPE  122  determines the number of nodes  102 ,  104  that have a shared copy of the memory line of information by reference to an entry  182  in the directory  180  corresponding to the memory line of information  184  (step  1230 ). The HPE  122  then sends a read-exclusive reply to the requesting node  102 ,  104  (step  1232 ). The read-exclusive reply includes a copy of the memory line of information and indicates the number of invalidation acknowledgments to expect. HPE  122  then sends an invalidation request to each node  102 ,  104 , if any, that has a shared copy of the memory line of information  184  (step  1233 ). The HPE uses the information in the directory entry for the memory line to identify the nodes having a shared copy of the memory line. HPE  122  then updates the state of the memory line of information  184  in the directory  180  to indicate that the requesting node  102 ,  104  is an exclusive owner of the memory line of information (step  1234 ) and removes the TSRF entry  210  in the TSRF  202  related to this memory transaction (step  1235 ). Thus, from the perspective of the home node  102 ,  104 , the entire memory transaction (including activity at other nodes  102 ,  104 ) is now complete, though other nodes  102 ,  104  must process protocol messages relating to this memory transaction. 
     The invalidation request is received by the sharer node(s)  102 ,  104 , and routed to the RPE  124  (step  1236 ) in each of those nodes, which respond by adding an entry  210  to the TSRF  202  (step  1237 ). The RPE  124  responds initially by sending an invalidation acknowledgment to the requesting node  102 ,  104  (step  1238 ). Additional steps taken by the RPE  124  depend upon whether the RPE is waiting on any requests related to the same memory line of information  184  (step  1239 ). See the discussion below, in the section entitled “Limited Fanout Daisy-Chaining Invalidation Requests,” for a description of another methodology of sending and handling invalidation requests and acknowledgments. 
     If the RPE  124  is waiting for a response to a read request, the invalidation request is merged with the outstanding read request transaction. To do this the RPE updates the TSRF entry  210  corresponding to the outstanding read request to indicate that an invalidation request related to the same memory line of information  184  has been received. Once the response to the read request is received, the PC  106  that initiated the read request/memory transaction is given a read-once copy of the memory line of information. In other words, the PC  106  is not permitted to cache a copy of the memory line of information  184 . This situation (receiving an invalidation request while waiting for a response to a read request) occurs because the CCP does not order protocol messages. More specifically, the home node  102 ,  104  received the read request and sent a response to the read request before receiving the read-exclusive request and sending the invalidation request, but the invalidation request is received before the response. 
     If the RPE  124  is waiting for a response to a read-exclusive request or an exclusive request, the invalidation request is acknowledged as noted above and no additional steps are taken (e.g., there is no limitation to a read-once copy). 
     Once these additional steps are complete, the RPE  124  removes the TSRF entry  210  related to this memory transaction (step  1240 ). 
     If the directory entry  182  indicates that a node  102 ,  104  has an exclusive copy of the memory line of information  184  (step  1222 -Yes), the HPE  122  sends a “read-exclusive forward” message to the owner node  102 ,  104  (step  1241 ), updates the state of the memory line of information  184  in the directory  180  to indicate that the requesting node  102 ,  104  is exclusive owner of the memory line of information  184  (step  1242 ), and removes the TSRF entry  210  in the TSRF  202  related to this memory transaction (step  1243 ). Thus, from the perspective of the home node  102 ,  104 , the entire memory transaction (which includes activity at other nodes  102 ,  104 ) is now complete, though other nodes  102 ,  104  continue to process this memory transaction. 
     The read-exclusive forward is received by the owner node  102 ,  104 , and routed to the RPE  124  (step  1244 ). The RPE  124  responds by adding an entry  210  in the TSRF  202  indicating that the memory transaction requires an internal response to a request for the memory line of information  184  in order to advance to another state (step  1245 ). The RPE  124  then sends a request for the memory line of information  184  from the L1 or L2 cache  110 ,  114  in which the memory line is locally stored (step  1246 ). Upon receiving the memory line of information  184  (step  1247 ), the RPE  124  sends a read-exclusive reply to the requesting node  102 ,  104  (step  1248 ). This protocol messages includes an up-to-date copy of the memory line of information  184 . The RPE  124  then invalidates the local copy of the memory line of information  184  (step  1249 ) and removes the entry  210  from the TSRF  202  (step  1250 ). 
     When the home node is the owner node, there is no need for the HPE of the owner node to send a read-exclusive forward to the owner node. Instead, the HPE sends a message to the L2 cache requesting that it forward a copy of the specified memory line and that it furthermore invalidate all cached copies of the memory line in the L2 cache and/or the L1 caches in the home node. The HPE would then send the read-exclusive reply message to the requesting node (i.e., steps  1246  through  1250  would be performed by the home node, since it is also the owner node in this example). 
     Upon receiving the read-exclusive response (step  1252 ), the steps taken depend upon the content of the response. As noted above, a read-exclusive request can result in a number of invalidation acknowledgments from nodes  102 ,  104  that have or had a shared copy of the memory line of information  184 . Additionally, the CCP does not requires protocol message ordering, so invalidation acknowledgments can arrive at the requesting node before a read-exclusive reply. If the response is an invalidation acknowledgment (step  1253 -Yes), RPE  124  updates the TSRF entry  210  in the TSRF  202  associated with this memory transaction to reflect that the invalidation acknowledgment was received (step  1256 ). More specifically, RPE  124  increments or decrements a counter in the counter fields  226  of the TSRF entry  210 . 
     If the response is not an invalidation acknowledgment (step  1253 -No), it is a read-exclusive reply, in which case the RPE  124  forwards the memory line of information  184  included in the reply to the PC  106  that requested the memory line of information (step  1254 ). If the read-exclusive reply indicates that a number of invalidation acknowledgment are to be received, the reply to the PC  106  also indicates that the memory transaction is not complete (unless the number of invalidation acknowledgments have already been received). RPE  124  then updates the TSRF entry  210  to reflect that the read-exclusive reply has been received and to indicate the number of invalidation acknowledgments, if any, to be received as well (step  1256 ). 
     Whether an invalidation acknowledgment or a read-exclusive reply is received, RPE  124  then determines if another protocol message is due (e.g., an invalidation acknowledgment or a read-exclusive reply). If no additional protocol messages are due, (step  1258 -Yes), RPE  124  removes the TSRF entry  210  from the TSRF  202  (step  1260 ). Otherwise, the entry  210  is not removed immediately, but is updated and eventually removed as additional, related protocol messages are received. Additionally, the RPE  124  sends an additional message to the PC  106  to indicate that the memory transaction is complete if the RPE  124  indicated to the PC  106  in its earlier reply that the memory transaction was not complete. 
     Until the TSRF entry  210  in the TSRF  202  is removed, incoming requests (read, read-exclusive, exclusive protocol messages) related to the memory line of information  184  are merged with the existing TSRF entry  210  related to this memory line of information  184  and put in the Suspended state. Once the read-exclusive reply and all invalidation acknowledgments, if any, are received, the state of the TSRF entry  210  is updated to the Active state so that it will be selected by the scheduler and the merged requests will be processed by the test and execution unit  194 . 
     Additionally, the write request steps described above with reference to FIGS. 12A-12D are subject to an optimization in preferred embodiments of the present invention. Specifically, if the requesting node  102 ,  104  already has a copy of the memory line of information, the RPE  124  of the requesting node sends an “exclusive request” to the home node  102 ,  104  instead of a “read-exclusive request.” If the requesting node  102 ,  104  is unambiguously listed as a sharer node  102 ,  104  in the entry  182  of the directory  180 , the steps are the same as those described above with reference to FIGS. 12A-12D, with the exception that the home node  102 ,  104  does not include the memory line of information  184  with the exclusive reply (a protocol message sent instead of a read-exclusive reply). 
     A given node is unambiguously listed as a sharer node if the sharer-information field  188  is in the limited-pointer format and includes the identifier of the given node or in coarse-vector format and only the requesting node is associated with a particular set bit. Thus, a given node is not unambiguously listed as a sharer node  102 ,  104  if (1) the sharer-information field  188  is in the limited-pointer format but does not include the identifier of the given node, or (2) the sharer-information field  188  is in the course-vector format and the bit associated with the given node  102 ,  104  is also associated with other nodes. 
     If the requesting node  102 ,  104  is not unambiguously listed as a sharer node  102 ,  104  in the entry  182  of the directory  180 , the HPE  122  converts the exclusive request to a read-exclusive request, which is then processed as described above. Alternatively, the HPE  122  sends a protocol message to the RPE  124  at the requesting node  102 ,  104  directing it to send a read-exclusive request to the home node. In another alternate embodiment, the RPE of the requesting node is configured to recognize when the number of nodes in the system is sufficiently great that the coarse vector bit used to represent the requesting node in the sharer information field  188  of directory entries also represents at least one other node. In this alternate embodiment, the RPE of the requesting node is further configured to not send exclusive requests when it recognizes, detects or knows this of this system status, and to instead send a read-exclusive request. In other words, in this situation the “exclusive request” optimization is suppressed or not used. 
     FIG. 13 illustrates steps taken to support a write-back request protocol message. A write-back request is initiated by a PC  106  when, for example, space is needed in the caches  110 ,  114  (step  1300 ). As an exception to the general rule described above, the write-back request is a high-priority protocol message. This exception is required because of a potential for the race condition described below. 
     The request is routed to the RPE  124 , which responds by adding an entry  210  in the TSRF  202  (step  1302 ) and sending a write-back request to the home node  102 ,  104  (step  1304 ). The entry  210  indicates that a write-back acknowledgment is required to advance the memory transaction to a next state. Additionally, the RPE  124  maintains the memory line of information  184  until the write-back acknowledgment is received and, if necessary, a forwarded request is received. If a forwarded request is received (e.g., read forward), it is handled as described above; however, the RPE  124  updates the state of the TSRF entry  210  to indicate that the forwarded request was received. 
     Upon being received at the home node  102 ,  104 , the write-back request is routed to the HPE  122  (step  1306 ) of the home node, which responds by adding an entry  210  in the TSRF  202  (step  1308 ). HPE  122  responds by checking the state of the memory line (step  1310 ). In particular, the HPE  122  determines if the directory entry  182  corresponding to the memory line of information still indicates that the “owner” node  102 ,  104  is the owner of the memory line of information  184 . If so (step  1312 -Yes), the HPE  122  updates the memory line of information  184  in the memory subsystem  123  (step  1314 ) and the state of the associated directory entry to indicate that the memory line of information  184  is no longer shared or owned by the former owner node  102 ,  104  (step  1316 ). HPE  122  then sends a write-back acknowledgment to the former owner node  102 ,  104  indicating that the memory transaction was successful (step  1318 ). The HPE then removes the TSRF entry  210  related to this memory transaction (step  1320 ). 
     If the directory entry  182  corresponding to the memory line of information does not indicate that the “owner” node  102 ,  104  is the owner of the memory line of information  184  (step  1312 -No), HPE  122  sends a write-back acknowledgment to the former owner node  102 ,  104  indicating that the write-back request was stale (i.e., that the memory transaction was not successful) (step  1318 ). More specifically, the write-back acknowledgment indicates that the home node  102 ,  104  forwarded a request related to the memory line of information  184  to the former owner node  102 ,  104  before receiving the write-back request. The HPE then removes the TSRF entry  210  related to this memory transaction (step  1320 ). 
     Upon receiving the write-back acknowledgment (step  1324 ), the RPE  124  of the former owner node determines if a race condition exists and whether it has been satisfied. As noted above, the write-back acknowledgment will indicate whether a race condition exists (i.e., whether the home node has forwarded a request related to the memory line that is the subject of the write-back request). The TSRF entry  210  in the RPE of the former owner node will indicate if the forwarded request has already been received and processed by the former owner node  102 ,  104 . If so, the RPE  124  removes the TSRF entry  210  for the memory transaction (step  1326 ). If not, the RPE  124  updates the state of the TSRF entry  210  to indicate that the forwarded request is required in order to advance the state of the memory transaction to a final state, and thus remove the TSRF entry  210 . 
     Limited Fanout Daisy-Chaining Invalidation Requests 
     In the above described embodiments, the home node  102 ,  104  always sends invalidation requests to sharer nodes  102 ,  104  individually. Each sharer node  102 ,  104  then sends an invalidation acknowledgment to the requesting node  102 ,  104 . Accordingly, the maximum number of invalidation requests and invalidation acknowledgments is entirely dependent upon the number of nodes  102 ,  104  sharing a given memory line of information  184  and bound only by the number of nodes  102 ,  104  in the multiprocessor system  100 . To reduce the number of protocol messages (e.g., invalidation requests and invalidation acknowledgments) active at any given moment, the invention configures directory entries (see FIG.  4  and the above discussion of the directory data structure  180 ) using the above described limited-pointer format and coarse-vector format, and furthermore employs a limited fanout, daisy-chaining invalidation methodology that ensures that no more than a specified number of invalidation requests and invalidation acknowledgments are active at any given moment, which avoids deadlocks. 
     The maximum number of invalidation requests and acknowledgments, resulting from a request for exclusive ownership of a particular memory line, that are active at any given moment is herein called the maximum fanout. In the preferred embodiments, the maximum fanout is a number between four and ten. The protocol engines of the present invention are configured to ensure that the number of invalidation requests and/or acknowledgments simultaneously active in a system as a resulting of a single a request for exclusive ownership of a particular memory line never exceeds the maximum fanout. 
     In preferred embodiments, the maximum number of invalidation requests and invalidation acknowledgments is set to four. Thus, the sharer-information field  188  of each directory entry  182  (FIG. 4) is configured to identify a maximum of DP (e.g. four) nodes when using the limited-pointer format. Similarly, the bits (e.g., 40-bits) of the sharer-information field  188  are grouped into DP (e.g., four) groups (e.g., 10-bit groups) when in the coarse-vector format. While the operation of the invention will be described with respect to an embodiment in which the sharer-information field  188  contains four groups of 10-bits for a total of 40 bits, in other embodiments the total number of bits in the sharer-information field, the number of groups of bits, and the number of bits per group, may vary substantially from those used in the preferred embodiment. 
     As described in more detail below, the home node  102 ,  104  sends at most one invalidation request for each of the four 10 bit groups. In particular, the home node sends an invalidation request to the first node, if any, identified as being a potential sharer by each 10-bit group within the sharer-information field. Thus, a home node  102 ,  104  sends at most four invalidation request messages to other nodes. Further, a subsequent set of invalidation request messages, if needed, are sent by the nodes that receive the initial invalidation request messages, this time to the second node, if any, identified as being a potential sharer by each respective 10-bit group within the sharer-information field. This process is repeated by each node receiving an invalidation request until the last node identified as being a potential sharer by each respective 10-bit group within the sharer-information field has received an invalidation request. Only the last identified node for each respective 10-bit group sends an invalidation acknowledgment to the requesting node  102 ,  104 . Using this limited fanout, daisy chaining-like methodology, the maximum number of invalidation request messages and invalidation acknowledgment messages that are active at any one time as the result of a request for exclusive ownership of a particular memory line never exceeds four, which is the maximum fanout in a preferred embodiment. In other preferred embodiment, the maximum fanout varies from four to ten. 
     In some embodiments of the present invention, the bits are grouped, for example, as follows: the first 10-bits, the second 10-bits, the third 10-bits, and the fourth 10-bits of a 40-bit sharer-information field  188  are groups 1-4 respectively. But in preferred embodiments of the invention, the bits within each group are interleaved. Specifically, in the preferred embodiment, the bits (and table  189  columns)  0 ,  4 ,  8 ,  12 ,  16 ,  20 ,  24 ,  28 ,  32 , and  36  form one group; bits (and table  189  columns)  1 ,  5 ,  9 ,  13 ,  17 ,  21 ,  25 ,  29 ,  33 , and  37  form a second group; bits (and table  189  columns)  2 ,  6 ,  10 ,  14 ,  18 ,  22 ,  26 ,  30 ,  34 , and  38  form a third group; bits (and table  189  columns)  3 ,  7 ,  11 ,  15 ,  19 ,  23 ,  27 ,  31 ,  35 , and  39  form a fourth group. 
     Though group identifiers (e.g., first group, second group, etc.) are not required for a node  102 ,  104  to determine which group it is in (since each node  102 ,  104  has access to its identifier) the number of bit groups and the number of bits in the sharer-information field  188  are required to establish the bit membership of each group (i.e., to determine the position of the bits of a given group within the sharer-information field  188 ) or equivalently, to establish the identity of a first node  102 ,  104  associated with each bit and additional nodes  102 ,  104  associated with each bit of a given group. 
     This aspect of the invention is now described in greater detail with reference to FIGS. 14A and 14B. The steps taken by the home node  102 ,  104  before and after an invalidation request is sent to a sharer node  102 ,  104  as described above are not changed in this embodiment of the invention. 
     In a first step, the home node  102 ,  104  determines the state of a given memory line of information  184  by reference to a corresponding directory entry  180  (step  1402 ). As described above, each directory entry  180  includes a state field  186 , which is preferably set to one of four values—including invalid, exclusive, shared, and shared-cv. Accordingly, this determination is made by reference to the state field  186 . If the state field  186  is set to shared, the format of the sharer-information field  188  is the limited-pointer format. If, however, the state field is set to shared-cv, the format of the sharer-information field  188  is the coarse-vector format. 
     If the state field  186  indicates that the sharer-information field  188  is in the limited-pointer format (step  1406 -Yes), the home protocol engine  122  extracts the node identifiers directly from each of the four sub-fields of the sharer-information field  188  (step  1410 ). The node identifier in each sub-field is valid if it is not the predefined null identifier. As noted above, in preferred embodiments the null identifier value is zero. The home protocol engine  122  then sends an invalidation request to each node  102 ,  104  identified in the sharer-information field  188  as a sharer node  102 ,  104  (step  1414 ). 
     If, however, the state field  186  indicates that the sharer-information field  188  is in the coarse-vector format (step  1406 -No), the home protocol engine  122  identifies for each group of bits within the sharer-information field  188  the first set bit (step  1418 ). Note that it is possible that one or more the groups may have no bits that are set. 
     Once the first set bit, if any, in each group of bits is identified, the home protocol engine  122  identifies the first node  102 ,  104  that corresponds to each of the identified first-set-bits using the techniques described above (step  1422 ). The above described techniques are extended somewhat in preferred embodiments however. If the first node  102 ,  104  that corresponds to a given identified first-set-bit is the requesting node or the home node, the home protocol engine  122  identifies the second node  102 ,  104  that corresponds to the identified first-set-bit. This step is repeated until a node  102 ,  104  that is neither the home node nor the requesting node is identified. If it is determined that none of the set bits in the group correspond to a node other than the home node and requesting node, an invalidation request is not sent by the home node for this particular group of bits in the sharer-information field  188 . In alternative embodiments, this step is not taken by the home node  102 ,  104 . Instead, the HPE  122  of the home node and the RPE  124  of the requesting node are configured to process these messages as described above without ever responsively invalidating the memory line of information  184 . 
     Once one or more nodes  102 ,  104  are identified (i.e., up to one node per group of bits in the sharer-information field of the directory entry), the home protocol engine  122  sends an invalidation request to each of the identified nodes  102 ,  104  (step  1426 ). Included each invalidation request is a sharer group field containing the 10-bit group of bits associated with the designated recipient of a given invalidation request and possibly an identifier of the 10-bit group. (The sharer group field is not included in an invalidation request if the sharer-information field  188  is not in the coarse-vector format.) This sharer group field is required because the sharer nodes do not maintain information about the nodes  102 ,  104  that share a given memory line of information  184 . The 10-bit group of sharer information that is sent along with the invalidation request permits each node that receives the invalidation request to identify the next node  102 ,  104  to receive an invalidation request as described above or to determine that there is no next node  102 ,  104  (i.e., that an invalidation acknowledgment should be sent to the requesting node  102 ,  104 ). 
     Additionally, the group identifier of the 10-bit group permits the sharer node  102 ,  104  to identify the position of each bit within the 10-bit group in the sharer-information field  188 , which also permits the sharer node  102 ,  104  to identify the next node  102 ,  104  (if any) to receive the invalidation request, as described above, or to determine that there is no next node  102 ,  104 . In an alternate embodiment, the group identifier is not included in the invalidation request and instead the protocol engines in each node are programmed to know the sharer group in which each such node resides. Since all the invalidation requests received by any particular node would always have the same sharer group identifier, the sharer group identifier is not strictly needed. 
     Upon receiving an invalidation request (step  1430 ) and adding a related entry  210  in the TSRF  202  (step  1432 ), a sharer node  102 ,  104  determines a next node, if any, by analyzing the sharer group field of the invalidation request. If all of the bits of the sharer group field are set to zero, there is no sharer information in the request ( 1434 -No) and therefore there is no next node to which to send the invalidation request. Instead, the remote protocol engine  124  in the sharer node  102 ,  104  sends an invalidation acknowledgment to the requesting node (step  1438 ). The sharer-node then processes the invalidation request as described above with reference to step  1238  (step  1458 ). 
     If the sharer group field in the received invalidation request includes any set bits (i.e., includes sharer information) (step  1434 -Yes), the remote protocol engine  124  in the sharer node  102 ,  104  determines the next node, if any, to receive an invalidation request (step  1442 ). The remote protocol engine in the sharer node identifies the next node by first determining the bit in the sharer group field that corresponds to the node identifier of the sharer node, and then determining if there is a next node (e.g., with a higher node identifier) that (A) also corresponds to that same bit of the sharer group field, and (B) is neither the home node (which is identified by the address of the memory line to be invalidated) nor the requesting node (which is identified by a requesting node field in the invalidation request). If not, the remote protocol engine looks for a next set bit (if any) in the sharer group field and determines if that next set bit corresponds to a node  102 ,  104  that is neither the home node  102 ,  104  nor the requesting node  102 ,  104 . This process continues, processing the bits of the sharer group field in a predetermined order (e.g., from left to right) until the remote protocol engine either identifies a next node, or determines that there is no next node. 
     If a valid next node  102 ,  104  is identified (step  1446 -Yes), the sharer node  102 ,  104  sends an invalidation request to the next node (step  1450 ). The sharer node  102 ,  104  includes in this invalidation request the same 10-bit sharer group field (and possibly a group identifier) that was included in the invalidation request received by the sharer node  102 ,  104 . The sharer node  102 ,  104  then processes the invalidation request as described above with reference to step  1238  (step  1458 ). The sharer node  102 ,  104  then removes the related entry  210  from the TSRF  202  (step  1460 ). 
     If, a valid next node is not identified (step  1446 -No), this means that the sharer node is the last node in the invalidation request daisy chain. In this case the sharer node sends an invalidation acknowledgment to the requesting node (step  1454 ). The sharer node then processes the invalidation request as described above with reference to step  1238  (step  1458 ). The sharer node  102 ,  104  then removes the related entry  210  from the TSRF  202  (step  1460 ). 
     Because each of the bits of the sharer group field may be associated with more than one nodes, the remote protocol engines in the nodes of the system are unable to determine which of the associated nodes (other than itself) are actually sharer nodes. When a node receives an invalidation request for a memory line of information  184  that it does not share, the node nevertheless sends an invalidation request (step  1450 ) or acknowledgment (step  1454 ) as described above. However, the processing of the received invalidation request at step  1458  comprises determining that the node is not a sharer of the specified memory line, and therefore no cache lines in the node are invalidated in response to the received invalidation request. 
     In other preferred embodiments, the bits of the sharer information field of the directory entries are divided into a larger number of groups of bits (e.g., four to ten groups). The number of such groups of bits corresponds to the maximum fanout of the daisy chained invalidation messages in these embodiments. 
     Alternate Embodiments 
     While the present invention has been described with reference to a few specific embodiments, the description is illustrative of the invention and is not to be construed as limiting the invention. Various modifications may occur to those skilled in the art without departing from the true spirit and scope of the invention as defined by the appended claims.