1. Field of the Disclosure
The present disclosure relates to updating data warehouses.
2. Description of the Related Art
Data warehouses store data tables that contain data received from external sources. As an example, the data may relate to network performance parameters.
a is a graph of staleness values related to data tables in a data warehouse;
b is a further graph of staleness values related to data tables in a data warehouse; and
In a particular embodiment, a disclosed method updates data tables stored in a data warehouse. The data warehouse may be a real-time data warehouse. The method includes receiving data for updating the data tables, generating update requests responsive to the receiving, calculating a staleness for a portion of the data tables, and scheduling data table updates on a plurality of processors based at least in part on the calculated staleness and the update requests. The method further includes transforming the data tables based on the scheduled data table updates to include a portion of the received data.
Generally, the staleness is indicative of an amount of time elapsed since the previous update of the data tables. Update requests may be assumed non-preemptible and accumulated update requests are batched together. The method may further include determining a stretch value for the update request, wherein the stretch value is indicative of the maximum ratio between the duration of time an update waits until it is finished being processed and the length of the update.
Further embodiments relate to a server for managing a data warehouse. The server includes a memory for storing the data warehouse, which includes a plurality of data tables. An interface receives data for updating the data tables and a processor for calculating a staleness for a portion of the data tables responsive to receiving data on the interface. Further instructions are for weighting a portion of the calculated stalenesses and scheduling data table updates for completion by a plurality of processors based at least in part on the weighted stalenesses. Accumulated update requests from the generated update requests are batched together.
To provide further understanding of disclosed systems, data warehouses and aspects related to updating data tables are discussed. Data warehouses integrate information from multiple operational databases to enable complex business analyses. In traditional applications, warehouses are updated periodically (e.g., every night) and data analysis is done off-line. In contrast, real-time warehouses continually load incoming data feeds for applications that perform time-critical analyses. For instance, a large Internet Service Provider (ISP) may collect streams of network configuration, performance, and alarm data. New data must be loaded in a timely manner and correlated against historical data to quickly identify network anomalies, denial-of-service attacks, and inconsistencies among protocol layers. Similarly, on-line stock trading applications may discover profit opportunities by comparing recent transactions against historical trends. Finally, banks may be interested in analyzing streams of credit card transactions in real-time to protect customers against identity theft.
The effectiveness of a real-time warehouse depends on its ability to make newly arrived data available for querying. Disclosed embodiments relate to algorithms for scheduling updates in a real-time data warehouse in a way that 1) minimizes data staleness and 2) under certain conditions, ensures that the “stretch” (delay) of each update task is bounded. In some cases, disclosed systems seek to schedule the updating of data tables to occur within a constant factor of an optimal solution for minimizing staleness and stretch.
Data warehouses maintain sets of data tables that may receive updates in an online fashion. The number of external sources may be large. The arrival of a new set of data records may generate an update request to append the new data to the corresponding table(s). If multiple update requests have accumulated for a given table, the update requests are batched together before being loaded. Update requests may be long-running and are typically non-preemptible, which suggests that it may be difficult to suspend a data load, especially if it involves a complex extract transform-load process. There may be a number p processors available for performing update requests. At any time t, if a table has been updated with data up to time r (i.e., the most recent update request arrived at time r), its staleness is t−r.
Given the above constraints, some embodied systems solve the problem of non-preemptively scheduling the update requests on p processors in a way that minimizes the total staleness of all the tables over time. If some tables are more important than others, scheduling may occur to prioritize updates to important tables and thereby minimize “priority-weighted” staleness.
Some disclosed systems use scheduling algorithms to minimize staleness and weighted staleness of the data in a real-time warehouse, and to bound the maximum stretch that any individual update may experience.
On-line non-preemptive algorithms that are not voluntarily idle can achieve an almost optimal bound on total staleness. Total weighted staleness may be bounded in a semi-offline model if tables can be clustered into a “small” number of groups such that the update frequencies within each group vary by at most a constant factor.
In the following description, details are set forth by way of example to facilitate discussion of the disclosed subject matter. It should be apparent to a person of ordinary skill in the art, however, that the disclosed embodiments are exemplary and not exhaustive of all possible embodiments. Throughout this disclosure, a hyphenated form of a reference numeral refers to a specific instance of an element and the un-hyphenated form of the reference numeral refers to the element generically or collectively. Thus, for example, widget 12-1 refers to an instance of a widget class, which may be referred to collectively as widgets 12 and any one of which may be referred to generically as a widget 12.
With the addition of updates 102, portions of data tables 108 are transformed based on the scheduling and the update requests 116. Update requests 116, in some embodiments, are for appending new data 114 to data tables 108 as updates 102.
In some embodiments, a stretch value for one or more of data tables 108 is determined and ranking data tables 108 is based at least in part on the calculated stretch. In some embodiments, the calculated stretch value the maximum ratio between the duration of time an update (e.g., an update based on update request 116-3) waits until it is finished being processed and the length of the update. Data warehouse server 118 may batch together accumulated portions of the generated update requests 116. New data 114 is distributed to data tables 108 as updates 102 by the scheduling processors 109 to minimize staleness. As shown, there are a number p processors, which is indicated by scheduling processor 109-p.
In some embodiments, update requests 116 are assumed non-preemptible. A portion of data tables 108 may be analyzed for a staleness value indicative of an amount of time elapsed since a previous update of the portion of data tables. A first portion of data tables 108 (e.g., data table 108-1) may be weighted higher than a second portion of the data tables (e.g., data table 108-2), and the scheduling processors 109 can schedule updates to these data tables responsive to the weighting results.
Illustrated in
Referring now to
In disclosed methods including method 200, updating the data tables may include appending new data to corresponding data tables. The staleness can be indicative of an amount of time elapsed since a previous update of the data tables. Some embodied methods include scheduling data table updates on p processors based at least in part on the calculated staleness and the update requests. Updating the data tables with the new data transforms the data tables based on the scheduled data table updates to include a portion of received data. Disclosed methods may include weighting a first portion of the data tables higher than a second portion of the data tables, wherein the scheduling is at least in part based upon the weighting.
Further embodiments relate to computer instructions stored on a computer readable medium for managing a plurality of data tables in a data warehouse. The computer instructions enable a processor to maintain a plurality of data tables in the data warehouse, receive data for updating a portion of the plurality of data tables, and generate update requests corresponding to the received data. For individual data tables, a staleness for the data table is calculated and ranked. Updating the data tables is scheduled based on the ranked stalenesses, and the data table is transformed (e.g., appended with new data) based on the updating. A stretch value may be calculated for a portion of individual data tables. The stretch value is indicative the maximum ratio between the duration of time an update waits until it is finished being processed and the length of the update. Further instructions allow for accumulated update requests to be batched and processed together. A portion of update requests may be non-preemptible.
Calculations and other aspects of updating data warehouses are touched on for a better understanding of disclosed systems. Suppose a data warehouse consists of t tables and p processors, and that p≦t. Each table i receives update requests at times ri1<ri2< . . . <ri,ki, where ri0=0<ri1. An update request at time rij contains data generated between times ri,j−1 and rij. The length of this update is defined as rij−ri,j−1. Associated with each table i is a real αi≦1 such that processing an update of length L takes time at most αiL. The constants αi need not be the same. For example, some data feeds may produce more data records per unit time, meaning that updates will take longer to load. At any point in time, an idle processor may decide which table it wants to update, provided that at least one update request for this table is pending. At time τ, table i may be picked, and the most recently loaded update may arrive at time rij. A processor would need to non-preemptively perform all the update requests for table i that have arrived between time rij+1 and τ, and there could be one or more pending requests. These pending update requests may be referred to as a “batch” with its length defined as the sum of the lengths of the pending update requests. In practice, batching may be more efficient than separate execution of each pending update. When the entire batch has been loaded, the processor can choose the next table to update.
At any time τ, the staleness of table i is defined τ−r, where r is the arrival time of the most recent update request rij that has been loaded.
The flow time of a task can be defined as the difference between its completion time and release time, and the stretch of a task is the ratio of its processing time to the flow time. Stretch measures the delay of the task relative to its processing time. These definitions may be slightly modified for various update tasks disclosed herein. For example, the flow time of an update request arriving at time rij may be redefined as the “standard” flow time plus the length of the update, i.e., rsij−ri,j−1. Furthermore, the stretch of said update request may be redefined as the “standard” stretch plus the length of the update, i.e.:
Given the above definitions, the total staleness of table i in some time interval is upper-bounded by the sum of squares of the flow times (using the modified definition of flow time) of update requests in that interval. There are no known competitive algorithms for minimizing the L2 norm of “standard” flow time; however, any non-preemptive algorithm that is not voluntarily idle may nearly achieve the optimistic lower bound on total staleness.
There are competitive algorithms for minimizing the L2 norm of flow times of all the update requests, using the modified definition of flow times defined previously. An algorithm is so-called “opportunistic” if it leaves no processor idle while a performable batch exists.
For any fixed β and δ such that 0<β,δ<1, Cβ,δ=√δ(1−β)/√6 may be defined. Given a number p of processors and t of tables, α:=(p/t) min{Cβ,δ,¼} may be defined. Then, provided that each αi<α, the competitive ratio of any opportunistic algorithm is at most
Any choice of constant parameters β and δ results in a constant competitive ratio. Note that as β→1 and δ→0, α approaches 0 and hence the competitive ratio approaches 1.
The penalty (i.e., sum of squares of flow times) of a given algorithm may be compared to a simple lower bound. Let A be the set of all updates. Independent of their batching and scheduling, each update i of length ai needs to pay a minimum penalty of ai2. This follows from the convexity of the square function. If a set of updates are batched together, it may be necessary to pay no less than the sum of the squares of the updates' lengths. Therefore,
Looking at the cost a particular solution is paying, let B be the set of batches in the solution. For a batch BiεB, let Ji be the first update in the batch, having length ci. This batch is not applied until, for example, di time units have passed since the release of Ji. The interval of size di starting from the release of update Ji is called the “delay interval” of the batch Bi, and di is called the “delay” for this batch. As for the length of the batch, denoted by bi, the following applies:
ci≦bi≦ci+di. (2)
For the penalty of this batch, denoted by ρi, the following applies:
Considering the case of one processor and two tables, if the updates of one table receive a large delay, it may indicate that the processor was busy applying updates from the other table (because it may be desirable for an applied algorithm to avoid remaining idle if it can perform something). Therefore, these jobs which are blocking the updates from the first table can pay (using their own sum-of-squares budget originally coming from the lower bound on OPT) for the penalty of a disclosed solution. It may be problematic if updates from the other tables (which are responsible for the payment) might be very small pieces whose sum of squares is not large enough to make up for the delay experienced (consider that the sum of their unsquared values is comparable to the delay amount). There may only be two batches, one of which is preferably large, occurring while a job of table 1 is being delayed. Another caveat is that the budget—the lower bound (LOW)—is ΣiεAai2, rather than ΣiεBbi2, which may be much larger. If each of these batches is not much larger than its first piece (i.e., bi=Θ(ci)), then by losing a constant factor, the length of the batch can be ignored, and job sizes may be adjusted. Otherwise, this batch has a large delay and some other batches should be responsible for this large delay. It may be preferable to have those other batches pay for the current batch's delay. This indirection might have more than one level, but may not be circular.
Each job iεA has a budget of ai2 units. A batch BiεB by (4), may need to secure a budget which is proportional to (ci+di)2. These conditions may be relaxed slightly in the following: A so-called “charging scheme” specifies what fraction of its budget each job pays to a certain batch. Let a batch Bi be “tardy” if ci<β(ci+di) (where β comes from the statement above); otherwise it is “punctual.” Let us denote these sets by Bt and Bp respectively. More formally, a charging scheme is a matrix (vij) of nonnegative values, where vij shows the extent of dependence of batch i on the budget available to batch j, with the following two properties.
1. For any batch BiεB,
2. There exists a constant λ>0 such that, for any punctual batch Bj,
The existence of a charging scheme with parameters β and λ gives a competitive ratio of at most
for an opportunistic algorithm.
Let:
Hence, the total cost of a solution is
The “execution interval of a batch B” may be defined as the time interval during which the batch B is being processed. Accordingly, its length is αi times the length of B, if the update belongs to table i.
Batch B blocks batch B′ if B's execution interval has intersection of positive length with the delay interval of B′. Note that many batches can block a given batch B′. A charging scheme with the desired properties may be introduced. This is done by defining how vij values are computed. If a batch Bi is punctual, this may be relatively straightforward: all vij values are zero except for vii=1/β2. Take a tardy batch Bi. In this case di is large compared to ci. During the interval of length di, during which Ji is waiting (i.e., the delay interval of batch Bi), all p processors should be busy. Let [r, r′] denote this interval. The total time is pdi. A relaxed version of this bound may be used to draw the conclusion (e.g., established by equation (6)).
A weighted directed graph with one node for each batch may be built. Punctual batches may be denoted as sinks (i.e., denoted as having no outarcs). Any tardy batch has arcs to all the batches blocking it, and there is at least one, since it has positive di. Even though punctual batches may be blocked by other batches, they have no outarcs.
The result is a so-called “directed acyclic graph” (DAG), because along any directed path in the graph, the starting (execution) times of batches are decreasing. The weight we on any such arc e is the fraction, between 0 and 1, of the blocking batch which is inside the delay interval [r, r′] of the blocked batch). Also, there is a parameter γ,
Then, for any two batches i and j, vij is defined as
where Pij denotes the set of directed paths from i to j. The dependence along any path is the square of the product of weights on the path multiplied by γ to the power of the length of the path. This definition includes as a special case the definition of the vij's for punctual batches i, since there is a path of length zero between any batch i and itself (giving vii=1/β2) and no path from batch i to batch j for any j≠i (giving vij=0 if j≠i).
Such a charging scheme may satisfy the desired properties as shown: The cost paid for each batch should be accounted for using the budget it secures, as required above in (5).
For any batch BiεB,
If B1, . . . , Bk are the children of B0, having weights w1, . . . , wk, respectively, in a run of any opportunistic algorithm, then
By the definition of the we's, the construction of the graph, the fact that B1, B2, . . . , Bk are all the batches blocking B0, and the fact that the k blocking batches are run on p processors in a delay interval of length d0 (so that their actual lengths must sum to at least 1/α times as much), it can be expressed that:
All but 3t of the batches may be removed, such that the sum of sizes of the remaining batches is at least ¾ times as large. Let [r, r′] be the delay interval corresponding to batch B0. There may be one batch per processor whose process starts before r and does not finish until after r. At most p batches may be kept, and in addition, the first at-most-two batches for each table that intersect with this interval may be kept. The contribution of the other batches, however many they might be, may be small. Consider the third (and higher) batches performed in this interval corresponding to a particular table. Their original tasks start no earlier than r and their release times do not exceed r′. The former is true, since otherwise, such pieces would be part of the first or second batch of this particular table; call them B1 and B2. However, suppose there exists an update J that starts before r and is not included in B1 or B2. As it is not included in B1 it preferably would have been released after the start of B1. The batch B2 cannot include any update released before r. So if it does not contain J, it should be empty, which is a contradiction. Hence, the total length of these batches (third and later) is no more than d0, as they only include jobs whose start and end times are inside the delay interval [r, r′]. Now
td0≦pd0/(4α), by definition of α. (17)
In conjunction with (16), the total (unsquared) length of the remaining at-most-3t batches is at least pd0/α−(¼)pd0/α=(¾)pd0/α. Considering that generally:
it may be inferred that the sum of squares of the at-most-3t leftover tasks is at least (¾pd0α)2/(3t), which exceeds p2d02/(6tα2).
To show that each batch receives a sufficient budget, let the “depth” of a node be the maximum number of arcs on a path from that node to a node of outdegree 0. The punctual nodes are the only nodes of outdegree 0. Induction on the depth of nodes may be used to prove, for any node Bi of depth at most Δ, that
For sinks, i.e., nodes of outdegree 0, the claim is apparent, since
Take a tardy batch B0 of depth Δ whose immediate children are B1, . . . , Bk. For any child Bi of B0, whose depth has to be less than Δ, there is the following:
Now we prove the inductive assertion as follows.
Therefore:
from (15) above, and because for jεBp, the first arc of the paths can be factored out to get
More precisely, let Pe,j, for an arc e=(u, v) and a node j, be the set of all directed paths from u to j whose second node is v. Then,
The second property of a charging scheme says that the budget available to a batch should not be overused.
For any batch Bj,
and tγ<1.
The delay intervals corresponding to batches of a single table may be disjoint, as shown: The delay interval of a batch BiεB starts at the end of Ji. Suppose this interval intersects one of Bj,j≠i, from the same table. Without loss of generality, assume that Jj starts at least as late as Ji. Thus, as Ji and Jj intersect, Jj should have been released before the delay interval of Bi ends. This is in contradiction with the definition of batching, as it implies Jj should be included in batch Bi.
To demonstrate the second property of the charging scheme, let the height of a node be the maximum number of arcs on a path from any node to that node. Induction on the height of nodes can be used to demonstrate for any node Bj of height H,
It may be noted that:
For a batch Bj at height zero (a source, i.e., a node of indegree 0), the definition of vij, which involves a sum over all i→j paths, would be 0 unless i=j, in which case vij=1/β2. Now the claim that λ≧1/β2 follows from the definition of λ and the fact that tγ<1.
As above, the last arc of the path can be factored out, except for the zero-length trivial path. Consider B0, whose immediate ancestors are B1, . . . , Bk with arcs ei=(B1, B0), . . . , ek=(Bk, B0), respectively. These incoming arcs may come from batches corresponding to different tables. However, it may be shown that the sum Σi=1kwe
Further, for any e, it may be shown that we<1. So
As the height of any ancestor Bi of B0 is strictly less than H, the inductive hypothesis ensures that the total load ΣjεBvji on Bi is no more than λ. The total load on B0 is:
by definition of vij in (15) above, noting that v00=1/β2 and the fact that for any i≠0, any path from Bi to B0 visits another batch Bi′ which is an immediate ancestor of B0,
Now
by the inductive hypothesis applied to i′ and (39)
by (37), then
as desired. The matrix (vij) is a charging scheme, as shown above.
Above the staleness measure is twice between the penalty measure and the lower bound LOW. Since the main theorem shows that these two outer values are close, staleness should also be close to the lower bound.
It can be argued that LOW is also a lower bound on staleness. Staleness is an integration on how out-of-date each table is. Tables can be considered separately. For each particular table, one can look at portions corresponding to different updates. If an update starts at time r and ends (i.e., is released) at r′, then at point r≦x≦r′, the staleness is no less than x. Thus, the total integration is at least ½ΣiεAai2. Staleness in most or all cases cannot be larger than half the penalty paid. For each specific table, the time frame is partitioned into intervals, marked by the times when a batch's performance is finished. The integration diagram for each of these updates consists of a trapezoid. It can be denoted by y the staleness value at time r. The staleness at r′ is then y+r′−r. Total staleness for this update is
as y≧0 and AB≦(A+B/2)2 for A,B≧0,=ρ, (44)
where ρ is the penalty for this batch according to our objective.
There is no known online algorithm which is competitive with respect to stretch. With regard to this, suppose there is one processor and two tables. A large update of size S1 arrives on the first table. At some point, it needs to be applied. As soon as this is done, a very small update of size S2 appears on the second table. Since preemption is not allowed, the small update needs to wait for the large update to finish. The stretch would be at least αS1/S2. But if there was advanced knowledge of this, the larger job could be delayed until completion of the smaller update.
Even if there is an offline periodic input, the situation might be less than optimal. In some cases, stretch can be large. Again, if there are two tables and one processor, one table may have a big periodic update of size S1. The other table may have small periodic updates of size S2. At some point, an update from table one should be performed. Some updates from table two may arrive during this time and need to wait for the large update to finish. So their stretch is at least a(S1−S2)/S2.
The above examples all work with one processor, but similar constructions show that with p processors, the stretch can be as large as desired because it is not bounded. To do this, p+1 tables and p processors are needed. The ith table has a period which is much larger than the (i+1)th one. The argument above was a special case for p=1.
The identity of a condition that allows stretch to be bounded is sought. Suppose each table has updates of about the same length (i.e., it is semi-periodic). In other words, the updates from each table have size in [A, cA] for some constant c. Any constant c would work (yet give a different bound finally), but c=2 is picked for ease of exposition. Further assume that tables can be divided into a few (g, to be precise) groups, such that periods of updates in each group is about the same thing (the same condition for the update lengths being in [A, 2A]). Then at least as many processors are needed as compared to the number of groups. Otherwise, there can be examples to produce arbitrarily large stretch values. Additionally, a reasonable assumption can be made that p is much larger than g. Each group is assigned some processors, in an amount proportional to their load. That is, the number of processors given to each group is proportional to the number of tables in the group. The algorithm is given in
At that point, each group forms an independent instance. Let us assume that for the t′ tables in a specific group with p′ processors, the upper bound is α≦p′/8t′ on each αi.
It can be shown that if all the updates of a group have sizes between A and 2A, and α≦p′/8t′, stretch is bounded by 3. Taking any one task, it can be shown it cannot wait for long, and thus its stretch is small. Note that stretch also considers the effect of batching this task with some other tasks of the same table. To this end, the execution of tasks is divided into several sections. Each section is either tight or loose. A tight section is a maximal time interval in which all the processors are busy. Loose is defined in the example as not tight.
Jobs can be ordered according to their release times, and ties may be arbitrarily decided. Let ωi denote the wait time (from release time to start of processing) of the ith job (say, Ji). Let θk be the length of the kth tight section (call it Sk). Recursive bounds can be established on values ωi and θk, and then induction may be used to prove they cannot be too large. There is inter-relationship between them, but the dependence is not circular. Generally, θk depends on ωi for jobs which are released before Sk starts. On the other hand, ωi depends on ωi′ for i′<i and θk for Sk in which Ji is released.
Let ik be the index of the last job released before Sk starts. The topological order of recursive dependence is then as follows: the ωi's are sorted according to i, and θk is placed between ωi
To derive a bound on θk, one can look more closely at the batches processed inside Sk. Let r and r′ be the start and end time of the section Sk. These batches correspond to updates which are released before r′. Let the so-called “load” at time r be the total amount of updates (released or not) until time r that has not yet been processed. Part of a batch that is released might have been processed (although its effect would not have appeared in the system yet). After half of processing time is passed, it can be considered that half of the batch has been processed. Updates which have not been released, but have a start time before r, may be considered to be part of the load (not all of it, but only the portion before r). The contribution to load by any single table is at most 2A+maxi≦i
where the first term corresponds to a (possibly not yet released) batch which is being formed while the other batch is running, the second term bounds the length of the running batch and the last term takes out the amount of load that is processed during the period from z to r. Noting that α≦1, equation (45) gives X≦2A+maxi≦i
Hence, the total load at time r is at most t(2A+maxi≦i
Rearranging yields:
Inequalities for ωi may be considered. Without loss of generality, consideration is made of the waiting time for the first update of a batch. It may be noted that they have the largest wait time among all the updates from the same batch. If a task has to wait, it should have one of two reasons. Either all the processors are busy; or another batch corresponding to this table is currently running Consider two cases:
The first case is when Ji is released in the loose section before Sk. If ωi>0, it is waiting for another batch from the same table. The length of the batch is at most maxi′<iωi′+2A. If as soon as this batch is processed at time τ, Ji processing may begin. Any job with higher priority than Ji should have been released before it is. When Ji is released, all these other higher-priority jobs are either running, or waiting for one batch of their own table. So, their count cannot be more than p−2 (there is one processor working on the table corresponding to Ji and one for each of these higher-priority jobs, and at least one processor is idle). In other words, there are at most p−2 tables which might have higher priority than Ji's table at time τ. Thus, Ji cannot be further blocked at time τ. Hence,
The other case is when Ji is released inside the tight section Sk. If Ji is processed after ω>θk time units pass since the start of Sk, a batch from the same table has to be under processing at the moment Sk finishes; otherwise, Ji would start at that point. However, processing of this batch must have started before Ji was released; or else, Ji has to be part of it. Moreover, similarly to the argument for the first case, it can be shown that as soon as the processing of the blocking batch from the same table is done, the batch corresponding to job Ji will start to be processed. More precisely, there can be at most p−1 other batches with higher priorities than Ji. So they cannot block Ji at time τ when the lock on its table is released. Hence,
since, it either waits for the tight section to end, or for a batch of its own table whose length cannot be more than maxi′<iωi′+2A.
One can pick of ω*=θ*=A/3 and use induction to show that ∀i:ωi≦ω* and ∀k:θk≦θ*, in part because α≦⅛ and p/α≧8t. Note that the right-hand side of Equations (47), (49) and (48) would be no more than θ*=ω* if one replaces these values for the wi and θk values in the formula. It only remains to observe that the dependence is indeed acyclic, which is clear by the ordering and by the fact that each formula uses the values to its left in the ordering.
The length of a batch is at most A′=A/3≦7/3A, where the first term comes from length of the first job in the batch (A≦A′≦2A), and the second term comes from the bound on its wait time. The resulting maximum stretch for any update piece would be bounded by 7/3(1+α)<3.
The algorithm preferably keeps the stretch low. The algorithm in
If the sum of (p−g)|Ti|/t for different groups is p−g, then Σi[(p−g)|Ti|/t≦p. Thus, one would use, at most, as many processors as were available. Then, in each group p′≧(p−g)t′/t. So the following applies: α≦p′/8t′. The arguments above, regarding if all of the updates of a group have sizes between A and 2A, apply to show that stretch is bounded by 3.
So-called “weighted staleness” can also be considered. That is, each table has a weight wi that should be multiplied by the overall staleness of that table. This takes into account the priority of different tables. In an online setting, no known algorithm can be competitive with respect to weighted staleness. As soon as a job from the low priority table is scheduled, a job appears from the high priority table which will then cost too much.
In the semi-periodic instance, weighted staleness of the algorithm in
To the maximum extent allowed by law, the scope of the present disclosure is to be determined by the broadest permissible interpretation of the following claims and their equivalents, and shall not be restricted or limited to the specific embodiments described in the foregoing detailed description.
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