Patent Document

BACKGROUND OF THE INVENTION 
     The invention relates to a computer memory arrangement, comprising a first plurality of input ports that are collectively coupled through a first router facility to selectively feed a second plurality of memory modules. Present-day computing facilities such as Digital Signal Processors (DSP) require both a great processing power, and also much communication traffic between memory and processor(s). Furthermore, ideally, both of these two performance aspects associated to the number of memory modules and processors, respectively, should be scalable, and in particular, the number of parallel data moves should be allowed to exceed the value of 2. 
     As long as a scale of 2 were sufficient, a possible solution would be to have two fully separate and fully functional memories, but then the selecting for a storage location between the two memories represents a complex task. Obviously, the problem will be aggravated for scale factors that are higher than 2. Moreover, the programs that will handle such separate storage facilities will often fall short in portability, such as when they have been realized in the computer language C. Therefore, in general a solution with a “unified memory map ” will be preferred. In practice, each memory access is then allowed to refer to any arbitrary address. 
     A realization of the above with two-port memories is quite feasible per se, but extension of the number of ports above two is generally considered too expensive. Therefore, the providing of specific hardware configurations on the level of the memory proper is considered inappropriate. Now on the one hand, the providing of the stall signal represents a useful feature. On the other hand, the reducing of the stall penalty overhead associated with conflict resolving in such pseudo-multiport memories requires that an additional slack time is added to the read-write access latency of the pseudo-multiport memory. In many applications indeed, the resulting extended access latency has no severe impact on loop-pipeline parts of an application, where throughput is important, but latency is generally insignificant. However, for certain other applications, the impact of the slack on the performance of running control-dominated code can be very significant. Such applications may inter alia relate to compression and decompression of code strings. Therefore, the present inventor has recognized the need for having the access latency be programmable to allow for selecting the optimum setup for the actual application. 
     SUMMARY TO THE INVENTION 
     In consequence, amongst other things, it is an object of the present invention to provide a solution that is generally based on one-port memories which collectively use a unified memory map, and wherein conflicts between respective accesses are accepted, but wherein the adverse effects thereof are minimized through raising the latency of the various accesses. Therefore, the solution according to the present invention is based on specific facilities that are provided as peripheral to the memory banks proper. Now, it has been found that control-dominated parts of an application will generally be sequential in nature, so that in these parts the number of parallel memory accesses will be limited anyway. Therefore, also the number of memory bank conflicts and the number of stalls would be limited as well. Moreover, taken on its own, a short memory latency will improve the performance of these parts anyway. 
     Now therefore, according to one of its aspects the invention is characterized according to the characterizing part of claim  1 . 
     The invention also relates to a computer arrangement comprising a fourth plurality of load/store units interfaced to a memory arrangement as claimed in claim  1 . Further advantageous aspects of the invention are recited in dependent claims. 
    
    
     
       BRIEF DESCRIPTION OF THE DRAWING 
       These and further aspects and advantages of the invention will be discussed more in detail hereinafter with reference to the disclosure of preferred embodiments, and in particular with reference to the appended Figures that show: 
         FIG. 1 , a pseudo multiport data memory template or parametrizable embodiment; 
         FIG. 2 , a request queue embodiment; 
         FIG. 3 , a request queue stage; 
         FIG. 4 , a request queue bypass embodiment; 
         FIG. 5 , a request queue controller embodiment; 
         FIG. 6 , a request routing facility from the request queues to the memory bank arbiters; 
         FIG. 7 , an acknowledgement routing facility from the memory bank arbiters to the request queues; 
         FIG. 8 , a bank arbiter embodiment; 
         FIG. 9 , an intermediate queue embodiment; 
         FIG. 10 , an intermediate queue controller embodiment; 
         FIG. 11 , an intermediate queue stage; 
         FIG. 12 , a result queue embodiment; 
         FIG. 13 , a result queue stage; 
         FIG. 14 , a result queue controller embodiment; 
         FIG. 15 , a result router embodiment. 
     
    
    
     DETAILED DESCRIPTION OF PREFERRED EMBODIMENTS 
       FIG. 1  illustrates a pseudo multiport data memory template or parametrizable embodiment. The template consists of an exemplary number of L building blocks (only the numbers  0 ,  1  , L−1 having been shown) that surround an array of B memory banks  20 – 24  (only the numbers  0  , b, and B−1 having been shown), each provided with an intermediate queue  26 – 30  that has been put in parallel between the input and the output of the memory module in question. The memory banks represent a unified address map. Generally, the values relate according to B≧L, but this is no restriction and in principle, the value of B may be as low as 1. 
     The external access signals may as shown at the top of the Figure, emanate from respective load/store unit facilites  17 – 19 . Each access signal comprises a  c hip  s elect cs, a  w rite  e na b le web, and each will furthermore present an address and write data. These signals will be sent to request/acknowledge router  32 . In general, each external access facility will also be able to receive external read data, as shown by the arrows such as arrow  35  from result router  34 . The purposes of the blocks surrounding the memory banks are threefold. First, the access requests to the various memory banks are routed by router facility  32  from the appropriate write port to the correct memory bank, and furthermore, the access results from the various memory banks are routed back by router facility  34  from the appropriate memory bank towards the correct read port. 
     Second, in the case of multiple access requests referring to the same memory bank in the same cycle, the obvious conflict will have to be resolved. To this effect, each memory bank  20 – 24  has a dedicated bank arbiter  36 – 40  located in front thereof. 
     Third, the number of bank conflicts is reduced by extending the latency of the bank accesses by an additional slack interval above the unavoidable access latency that is associated to the memory banks proper. The additional slack is obtained through delaying the accesses in L parallel request queues  42 – 46 , wherein L is the number of read/write ports, and furthermore through delaying the results in B parallel result queues  48 – 52 , wherein B is the number of memory banks. 
     The request queues  42 – 46  may delay input requests over a time as depending on the circumstances. A supplemental delay in the result queues  48 – 52  may produce an overall delay between the actual request and the instant that the result becomes available at the output port, wherein the overall delay has a uniform value. This latter feature implies that the compiler may take this uniform delay value into account when scheduling the program. The various building blocks may operate as discussed hereinafter more extensively. As shown in the embodiment, the request queues and result queues operate on the basis of serial-in-parallel-out, but this is no explicit restriction. Finally as shown, each memory bank  20 – 24  has a respectively associated intermediate queue  26 – 30  to be further discussed hereinafter. 
       FIG. 2  illustrates a request queue embodiment, wherein for clarity among a plurality of corresponding items only a single one has been labeled by a reference numeral. The arrangement as shown has a single controller  60  and a plurality of delay stages  64  equal to the length of the slack S measured in clock cycles. The controller  60  receives the master clock signal  59 , and is therefore continuously active, and furthermore the signals chip select and stall. It produces S stage valid flags  0  . . . S−1 that control respective clock gates  62 . A stage such as stage  64  is only active with valid data therein, thereby avoiding to spend power on invalid data. The signals cs, web, address, and data have been shown as in  FIG. 1 . A request that cannot be handled immediately is queued in the request queue. From every request queue stage  64 , a memory bank request signal  61  may be outputted in parallel. Request priority generally follows seniority (FIFO). Requests granted get a corresponding acknowledgement signal  63  and are thereupon removed from the queue. Such acknowledgement signals can arrive for each stage separately, inasmuch as such acknowledgements may in parallel originate from respective different memory banks. 
     If a request has traveled through all of the queue and arrives at the bottom of the queue (stage S−1), the flag signal full  65  is raised, which implies that the request cannot be handled in the standard interval recognized by the load/store latency. Such full state will then cause a stall cycle to the requesting facilities to allow resolving the bottleneck, whilst maintaining the existing processor cycle without further advancing. The bypass facility  66  will be discussed hereinafter. The signals full from the respective request controllers  60  are ORED in OR gate  65 A, the output thereof representing the stall signal for the complete arrangement of load/store facilities  17 – 19 . Although not shown in particular in the figure, this stall signal will then be sent to all of the relevant load/store units  17 – 19  in  FIG. 1 . 
       FIG. 3  illustrates a request queue stage which generally corresponds to a shift register stage, allowing to store the quantities web  43 , address  45 , and data  47 . Load/store operations have not been shown in particular in this Figure. In general, they can be effected in either of two modes. In the short latency mode, load/store operations will only experience the intrinsic memory latency L 1  of the memory that may, for example, be an SRAM. Then, each memory bank conflict will result in a stall cycle. In the long latency mode however, a slack interval S is added to memory latency L 1 , so that load/store operations will experience an overall latency of (S+L 1 ). The request queue embodiment selectively supports both latency modes by using a special bypass block  66 . In the short latency mode, this block will be used to bypass all stages  64 , thereby assigning to the actually incoming request the highest priority, whilst disabling all others. The associated latency mode signal has been labeled  67 . 
       FIG. 4  illustrates a request queue bypass embodiment. It has been constructed from a set of multiplexers  66 A,  66 B,  66 C, that will collectively select either the latest incoming request, or the request that is leaving the final queue stage S−1. The actually selected request will have the highest priority. 
       FIG. 5  illustrates a request queue controller embodiment, for realizing block  60  in  FIG. 2 . For brevity, only the various logic elements pertaining to the uppermost stage have been labeled. Acknowledge signal ack  0  arrives at top left and feeds AND gate  78 , as well as after inversion in element  73 , AND gate  72 . The chip select cs signal feeds AND gates  72 ,  78 , and  80 , the latter furthermore receiving acknowledge signal ack S. The latter two AND gates feed selector  86 , that is controlled by the long/short latency mode signal to select the input signal as indicated. The transmitted signal enters OR gate  88 , and is selectively transmitted through selector  90 , that is controlled by the stall signal as shown. The signal transferred is latched in latch  92 , to operate as mask signal. The mask signal is furthermore retrocoupled to OR gate  88 . Next, the mask signal is inverted in inverter  96  and then fed to clocked AND gate  98 . The output of AND gate  98  is fed to selector  100  which is controlled by the latency mode signal, so that the output of selector  100  will be either  0  , or equal to the output signal of clocked AND gate  98 . 
     On the left hand side of the arrangement, the inverted acknowledge ack  0  signal will be clocked to selector  84  that is controlled by the stall signal. The signal selected is fed to selector  94  that on its other input receives a  0  signal and is itself controlled by the stall signal, to on its output generating the stage valid signal  76 , cf.  FIG. 2 . Furthermore, the output of selector  84  will be fed to latch  70 . The latch content represents the request signal req  1 , and is furthermore retrocoupled to selector  82  which on its other input receives a zero ( 0 ), and which selector is controlled by the signal ack  1  from the next lower stage. 
     For the other stages, generally the items corresponding to items  70 ,  72 ,  73 ,  82 ,  84 , and  94  will be present. Now, the chip select value cs travels through a dedicated one-bit shift register with stages like stage  70 , which register thus contains all pending requests. Furthermore, a received acknowledgement signal like signal  63 A will clear the chip select signal cs at the stage in question through inverting an input to an AND gate like  72 . Furthermore, from every chip select pipeline stage a valid flag like flag  76  is derived that drives the associated clock gate  62  in  FIG. 2 . The register will keep shifting as long as no memory bank conflicts will occur in the memory, that is, as long as no request queue gets full. ( 65 ) A conflict will however automatically cause a stall cycle, that stops the shifting of the queue. While the queue remains stalled, memory bank conflicts may get resolved, which means that actual requests will still be acknowledged. Hence, the clearing of acknowledged requests will continue during the stall interval. 
     Note that the final stage S has the request controlled in an inverse manner with respect to the other stages. Furthermore, the final stage S comprises AND gate  102  that corresponds to AND gates  72  of earlier stages, and also a second AND gate  104  that receives the inverted value of acknowledge signal ack S, and furthermore the inverted mask signal from latch  92 . The two AND gates feed a selector  106  that is controlled by the latency mode control signal and transmits the full signal. When a request that had caused a queue to raise its full flag is acknowledged in this manner (without occurrence of a further full signal), the stalling signal from OR  65 A is automatically terminated. 
     Furthermore, in the request queue facility, a bypass element  66  controlled by the latency mode signal  67  is also visible. As can be seen in  FIG. 2 , in the short latency mode, the entire queue will be bypassed, to assign the highest priority to the latest incoming request, whilst coincidently therewith, blocking all other requests. In the long latency mode, the seniority among the requests will generally prevail. The latency mode signal  67  may be given by an operator and/or by the system, such as being based on statistical and/or dynamic data. A longer latency, even caused by only a single stage, will dramatically decrease the number of conflicts, and thereby, the number of delaying stalls. However, a longer latency will also present a longer delay. The system should be controlled by the best trade-off that were relevant for the application or application interval in question. 
       FIG. 6  illustrates a request routing facility from the request queues to the memory bank arbiters. This combinatory network routes all requests from all request queues to the appropriate memory bank arbiter ( 36 – 40  in  FIG. 1 ), and would route acknowledgements pertaining to these requests back to the associated request queue, the latter not having been shown in  FIG. 6 . Since the memory map is uniformly interleaved over the various memory banks, the specific bank in question is determined as based on examining the least significant bits associated with the relevant access request signals. The latter is effected by bit select items like  69 . The result of this bit select operation controls a demultiplexer-like item  70  that will in consequence route a one-bit request flag to the intended memory bank arbiter such as items  36 – 40 . The components web, address, and data of the request are directly forwarded to all bank arbiters in parallel on an interconnection shown in bold representation. The total number of request lines arriving at every bank arbiter then equals the number of request queues times the maximum number of requests generated by each request queue. With a slack S, this number is (S+1)*L. 
     Furthermore, since a request is always directed to one single bank, each request may only be acknowledged by one single bank arbiter. Therefore, for each particular request, an ORING in elements  37 ,  39 ,  41  of all corresponding acknowledge flags from the respective arbiters  36 – 40  will yield the associated acknowledge value.  FIG. 7  illustrates this acknowledgement routing facility from the various memory bank arbiters to the pertinent request queues. 
       FIG. 8  illustrates a bank arbiter embodiment. The bank arbiter is operative for selecting the highest priority request presented to its input, for acknowledging this request, and for forwarding the information associated with the request to the memory bank in question. For this purpose, the incoming request flags are considered as a sorted bit vector with the relatively higher priority request flags as most significant bits, and the relatively lower priority request flags as least significant bits. The arbiter will search for the bit with the highest significance level in this vector that is “1”. This bit corresponds to the valid request with the highest priority. The index of the bit in question is used to select the request that must be acknowledged. It also indicates the load/store unit from which the request is coming. This latter information will send load data for a read access back to the proper load/store unit. To have this information available at the instant on which the loaded data is available for reading, the index is sent to the intermediate queue, through which it travels to stay in synchronism with the data to be read from the addressed memory bank. 
     To this effect, the arbiter facility comprises selector facilities like  120  and  122 . Selector facility  120  is controlled at the left side by the various requests ranging from LSU L−1, row S), to (LSU  0 , row  0 ). Furthermore, the selector facility  120  receives at the upper side the web, address, data, abd Isu id signals, and furthermore a signal def, or  0 . As shown, the priorities have a leading string of  0 , . . . zeroes, then a first “1” at mutually exclusive positions, followed by string that may have any appropriate value. The selector will output the selected request, web, address, data, isu id, and remaining slack signals. 
     A second selector facility  122  is controlled by the same control signals at the left hand side as earlier, and receives at the various top bit strings with a single “1” at mutually exclusive positions, and furthermore exclusively zeroes. The selector will output acknowledge signals ranging from (LSU L–1, row S), to (LSU  0 , row  0 ). 
     The selecting operation of the most significant bit from a bit vector can be viewed as a large multiplexer which takes the bit vector as control input, and which selects ports on the basis of the value of the Most Significant Bits MSB. In this context,  FIG. 9  illustrates an intermediate queue embodiment The intermediate queue is therefore used as a delay line for synchronizing the information pertaining to a load access during the interval in which the actual SRAM in the data memory is accessed. It consists of a single controller  130  and a number of stages such as stage  134 , the number thereof being equal to the latency L 1  of the SRAM. The controller is clocked by the master clock  134  and will therefore always be active: it will produce a number of L 1  stage valid flags that control one clock gate such as gate  132  for each stage. As a result, a particular stage will only be active when valid data are present in that stage. No power consumption is wasted on the storing of invalid data. All of the signals chip select (cs), write enable (web), address, and data will enter the intermediate queue at its top end. Only information pertaining to load requests is stored in the queue. The final output result includes as shown a signal load valid, a signal remaining delay, and a signal load origin. 
       FIG. 10  illustrates an intermediate queue controller embodiment. It consists of a single-bit-wide delay line with stages like latch  136 , which is serially fed by the signals cs and web, and which holds the valid flags for the various stages. Each such flag signifies a load operation and is created by ANDING the chip select and low-active write enable input signals. The serial output signal is load valid. 
       FIG. 11  illustrates an intermediate queue stage. Each stage has two registers, first one ( 140 ) for holding an identifier to identify the load/store unit that issued the load request, and another one ( 138 ) to hold the remaining delay value, thereby indicating how much slack the request still will have to undergo in the data memory to meet the intended load/store latency interval. If a conflict occurs in the memory that leads to a processor stall, processor cycle time will be halted. To keep the remaining delay value consistent with processor cycle time, the value in this case will be incremented in incrementing element  142 ; the pertinent selection is executed through selector  144  that is controlled by the stall signal. 
       FIG. 12  illustrates a result queue embodiment, here consisting of a single controller  146  and a plurality of stages like stage  150 , the number of these stages is a function of the slack, the memory latency times and the number of load/store units minus 1. The number of stages is approximately equal to MAX(S,L*(LSU−1). The controller is clocked by the master clock and therefore, always active. It produces S stage valid flags that control S clock gates like clock gate  148 , one for every stage. As a result, the stage is only active when it stores valid data, again for diminishing the level of power consumption. 
     A result queue will collect and buffer the data and other informations coming out of its corresponding SRAM and out of the intermediate queue to perform the final synchronization with the processor core. The signals load valid, load data, load/store unit identifier, and remaining delay will successively enter the queue in stage  150  at its top as shown. Once in the queue, the result as loaded undergoes its final delay. Once the remaining delay of a result reaches the value zero, the result is issued as a valid result at one of the stage outputs of the relevant queue. Multiple results that are intended for respective different load/store units can leave the queue in parallel. 
       FIG. 13  illustrates a result queue stage embodiment that consists of three registers that respectively hold the remaining delay value (register  156 ), the data (register  160 ), and the load/store unit identifier of a load operation (register  158 ). All stages taken together constitute a shift register. In normal operation, the remaining delay value is decremented in each stage (element  152 ), so that with the traversing of successive stages the remaining delay will eventually reach zero, to indicate that the total load/store latency has been attained. At that instant, the result may be sent back to the load/store unit. If a conflict occurs in memory that leads to a processor stall, the processor cycle time will be standing still. To keep the remaining delay value synchronized with the processor cycle time count, the remaining delay value is kept constant in this case through circumventing decrementer stage  152  and appropriate control of selector  154 . 
       FIG. 14  illustrates a result queue controller embodiment, thereby realizing a delay line for producing stage valid flags for the clock gates that control the various stages. At every stage in the delay line, the remaining delay value is examined. Once this value attains zero, the stage valid flag in this stage is cleared. Note also that if the processor is stalled by a memory conflict, no result can be sent back, since no load/store unit will be able to receive it. Therefore, in this case the result valid flags are cleared. Each stage in this controller comprises the following items, that are referred by number only in the first stage. First, the remaining delay value is entered into decrementing element  162 . The output thereof, which equals the ORed value of all bits of the delay value, is fed to AND gate  164 , together with the load valid signal in a serial arrangement across all stages. The output of the AND gate is fed to a selector  172  that is controlled by the stall signal. The output of the selector is latched in element  174 , and thereupon fed to the next serial stage. Furthermore, the reducing element output is inverted in item  166 , and likewise ANDED with the load valid signal in AND gate  168 . The output value of this gate is fed to selector  170  which furthermore receives a “0” signal and which is controlled by the stall signal. The output signal of the selector  170  may yield the result valid 0 signal. The difference in the ultimate stage relates to the leaving out of items  164 ,  172  and associated wiring. 
       FIG. 15  illustrates a result router embodiment. This router executes sending the valid results that leave the result queue, back to the load/store units  162 – 164  from which the relevant requests did originate. The load/store unit id is used to determine the target load/store unit. A demultiplexer  180 – 184  selects the proper load/store unit whereto a valid flag should be sent. Since it is certain that in each cycle at most one valid request is sent back to a particular load/store unit, the result data to be sent back are determined by first bit-wise ANDING all result data with their corresponding result flag in two-input AND gates like  174 – 178 , and next ORING them in OR gates  168 – 172  for each particular router.

Technology Category: g