Patent Application: US-21833405-A

Abstract:
a method of transmitting data packets , where randomness is added to the schedule . universal broadcast schedules using encoding and randomization techniques are also discussed , together with optimal randomized schedules and an approximation algorithm for finding near - optimal schedules .

Description:
in this paragraph , it will be shown that any deterministic schedule has poor worst - case performance . in particular , the worst - case waiting time of such schedule is equal to the length of the longest packet . indeed , suppose that i is the longest packet in the schedule and denote by x i the length of packet i and by [ t i , t i + 1 ] the transmission interval of this packet . then , if the information request arrives immediately after time t i + 1 − x i , the waiting time will be close to x i time units . in order to improve the worst - case performance of wireless clients , the notion of random schedules is introduced . in a random schedule each packet is transmitted , in an encoded form , over a time interval whose length is a random variable , distributed according to a certain probability distribution . a random schedule can be viewed as a probability distribution over a set of deterministic schedules with unequal transmission intervals . with random schedules , a waiting time at any time t is a random variable . thus , the design goal is to minimize the expected worst - case waiting time of the schedule , where the expectation is taken over the probability distribution of the server . in order to implement random schedules a random bit generator can be used . the random bit generator outputs a stream of random bits . this stream is then converted to a sequence of random numbers , which are distributed according to a given probability distribution function . these numbers are then used to determine the length of the intervals during which packets are being transmitted . in order to evaluate the worst - case performance of a schedule , the notion of an adversarial client or an adversary is employed . the goal of such a client is to maximize waiting time by generating requests at times which are the least desirable for the schedule . the adversary knows the probability distribution of the schedule and can place its request based on the observed history . the adversary captures the situations in which clients base their requests on information previously broadcasted over the channel . for example , consider the schedule depicted in fig2 . in this schedule the length of each interval is exactly one time unit . suppose that a client has to receive a stock market update every 10 time units . thus , after listening to an update at time t , it expects the next update at time t + 10 . however , due to clock skew , the client begins to listen to a channel at time t + 10 + ε , where ε is a small value . in this case , the expected waiting time of the client is close to one time unit . for clarity , the rest of the application focuses on settings in which the broadcast channel is dedicated to a single information source . in such settings , each packet carries the same information , e . g ., stock quotes . the content of each packet , however , is different , because each packet captures the most recent state of the information source . it is also assumed that all packets have an identical size and that the transmission of a packet ( without encoding ) requires one time unit . in accordance with the present disclosure , there are two ways to transmit data : ( 1 ) with encoding : each interval contains an encoding of a packet , such that the length of the encoded packet is equal to the length of the interval . ( 2 ) no encoding : each interval contains a portion during which no data is being transmitted followed by a portion during which the original data of unit length ( without additional encoding ) is transmitted . ( 1 ) is discussed in the section “ encoded transmission ” of the present application , while ( 2 ) is discussed in the section “ broadband schedules with no encoding ” of the present application . a broadcast schedule specifies the times at which the packets are generated and transmitted . each packet is allocated a time interval whose length is at least one time unit . each packet is periodically broadcasted ( in correct bit order ) over a corresponding time interval . such encoding allows the client to restore the original packet from any portion of the interval whose length is at least one time unit . definition 1 [ schedule s ] a schedule is a sequence { x 1 , x 2 , . . . }, x i ≧ 0 , such that x i + 1 specifies the length of the time interval allocated for packet i . a schedule s ={ x 1 , x 2 , . . . } can also be defined by its transmission sequence { t i , t 2 , . . . }, where t n represents the beginning of the n th interval , that is , t 1 = 0 and t n = σ i = 1 n = 1 x i + n − 1 for all n & gt ; 1 . fig2 depicts schedule s 1 ={ 0 , 0 , 0 , . . . }. in this schedule packet i is transmitted over the interval [ i , i + 1 ], for i = 0 , 1 , . . . . fig3 depicts a schedule s 2 ={ 2 , 2 , 2 , . . . }. while in the first schedule each packet is sent without encoding , in the second schedule each packet is periodically broadcasted over an interval of length 3 . efficient data broadcast systems require schedules that minimize waiting time , i . e ., the amount of time spent by the client waiting for data . let s be a schedule , and suppose that a client request is placed at time t . also , let n be the current interval , i . e ., the interval for which t n ≦ t & lt ; t n + 1 . the waiting time depends on the time left in the current interval , i . e ., t n + 1 − t . specifically , if t n + 1 − t ≧ 1 then the client request can be satisfied within the current interval , hence the waiting time is zero . otherwise , the client must wait until the beginning of the next interval , hence its waiting time is t n + 1 − t . definition 2 . [ waiting time , wt ( s , t )] the waiting time wt ( s , t ) for a request at time t using a schedule s ={ x 1 , x 2 , . . . } is defined as follows : wt ⁢ ⁢ ( s , t ) = { t n + 1 - t if ⁢ ⁢ t n + 1 - t & lt ; 1 0 otherwise where n is the current interval , i . e ., the interval for which it holds that t n ≦ t & lt ; t n + 1 . for example , in schedule s 1 , the waiting time of the request that arrives at time t is ┌ f ┐− t . in contrast , in schedule s 2 , the waiting time is zero for many requests . indeed , suppose that the client arrives at time t , 3 ≦ t ≦ 6 . if the remainder of the current interval is more than one unit , i . e ., 6 − t ≧ 1 , then the waiting time of the client is zero . otherwise , the client must wait 6 − t time units for the beginning of the next interval . the staleness captures the age of the information delivered to the client . the staleness depends on both the amount of time that has passed from the beginning of the current interval n , i . e ., t − t n , and the amount of time left in the current interval , i . e ., t n + 1 − t . specifically , if t n + 1 − t ≧ 1 then the client request can be satisfied within the current interval . in this case the client receives the data t − t n time units after it was obtained from the database , hence the staleness is t − t n . if t n + 1 − t & lt ; 1 , then the client must wait to the beginning of the next interval , and the information it receives will be up - to - date , i . e ., the staleness will be zero . definition 3 [ staleness , st ( s , t )] the staleness st ( s , t ) for a request at time t using a schedule s ={ x 1 , x 2 , . . . } is defined as follows . st ⁢ ⁢ ( s , t ) = { 0 if ⁢ ⁢ t n + 1 - t & lt ; 1 t - t n otherwise where n is the current interval , i . e ., the interval for which it holds that t n ≦ t & lt ; t n + 1 . note that in the schedule s 1 the staleness is zero for any request . in contrast , in schedule s 2 , the staleness is positive for some requests . for example , suppose that the client arrives at time t , 3 ≦ t ≦ 6 . if the remainder of the current interval is more than one unit , i . e ., 6 − t ≧ 1 , then the client request is satisfied within the current interval and the staleness is t − 3 . otherwise , the client must wait for the beginning of the next interval , hence the staleness is zero . staleness and waiting time have certain duality properties . namely , for any specific value of t it is the case that exactly one of the two is not zero . in general , as shown below , a lower waiting time can be achieved at the cost of higher staleness . the goal of universal schedules is to minimize waiting time for any client , regardless of its behavior . for that end , it is assumed that the requests are generated by an adversary , whose purpose is to generate requests resulting in high waiting time . both schedules s 1 and s 2 have a poor performance in the presence of an adversary . for example , suppose that an adversary puts its request at time t = t i − 1 + δ , for some small value of δ & gt ; 0 , where t i is the beginning of some interval i . in this case , the waiting time is 1 − δ , which can be arbitrary close to one time unit . now , it is shown that the worst - case waiting time of any deterministic schedule is close to one time unit ( the only exception is schedule in which one packet is broadcasted over an infinite interval . this schedule , however , has unbounded staleness and hence cannot be used for practical purposes ). since the adversary knows the schedule , it can generate a request 1 − δ units of time before the transmission of the next item . in order to deal with such a powerful adversary and to improve the worst - case behavior of the system the applicants propose to add randomness to the schedule . in the proposed schedules according to the present embodiment , each packet is transmitted , preferably in an encoded form , over an interval of random length . as mentioned above ( paragraph [ 0025 ]), there are two ways the randomness can be incorporated into the schedule . in the first approach , detailed in this section , a certain level of redundancy is added to the schedule . in this approach every packet is transmitted over a larger time interval in an encoded form . this time interval is larger than the time required to transmit the original packet . encoding ensures that client is able to satisfy its request by listening to a sufficiently large portion of the interval . one possible form of encoding is to transmit the packet periodically over the interval in the correct bit order , but other forms can also be considered . those forms are assumed to be known to the person skilled in the art and will not be discussed here in detail . in the second approach , detailed in section 2 , an empty interval of random length is inserted into the schedule before the transmission of each packet . in the proposed random schedules the lengths x i of all intervals are random variables . this implies , in turn , that the values of waiting time wt ( s , t ) and staleness st ( s , t ) for any request time t are also random variables . in randomized settings , there are several types of adversaries that can be considered ( see reference [ 5 ], chapter 7 . 1 ). one type is an oblivious adversary , i . e ., an adversary that decides about its requests in advance , before the broadcast begins . this adversary is relatively weak and can be dealt with by transmitting an empty interval of random length followed by a deterministic schedule . in what follows , the adversary is assumed to be adaptive , i . e ., a request generated at time t is based on the history of the schedule from the beginning of the transmission up to time t . such an adversary models the worst possible access pattern , including possible correlations between requests and prior transmissions . in order to define the expected staleness and waiting time for adaptive adversaries the probability distribution of a given random schedule s is conditioned on the history of s up to time t . intuitively , the history of a schedule can be described by the lengths of the intervals transmitted up to time t . definition 4 a history h =( t , x 1 , x 2 , . . . , x l ) of a random schedule s ={ x 1 , x 2 , . . . } at time t is the event in which ( a ) for all i , 1 ≦ i ≦ l , it holds that x i = x i ; and ( b ) σ i = 1 l x i + 1 ≦ t & lt ; σ i = 1 l = 1 x i + l + 1 . in other words , h =( t , x 1 , x 2 , . . . , x l ) is the event in which ( a ) for the first l random variables in s it holds that x i = x i , and ( b ) the number of intervals that are completely broadcasted up to time t is l . formally , let h be a history event . h is said to be admissible if it occurs with a non - zero probability . for admissible histories h , let s | h be the schedule obtained by conditioning s on the event h . notice that s | h is also a random schedule . the worst - case expected waiting time of a schedule is defined as follows : ewt ⁢ ⁢ ( s ) = sup h , t ⁢ e ⁢ [ wt ⁢ ⁢ ( s ⁢ | ⁢ h , t ) ] est ⁢ ⁢ ( s ) = sup h , t ⁢ e ⁢ [ st ⁢ ⁢ ( s ⁢ | ⁢ h , t ) ] ewt ( s ) and est ( s ) capture the worst - case expected waiting time and staleness of a schedule up to time t . the expectation is taking over the schedule distribution s | h , and the maximization is over admissible history events h . consider the schedule in which the length of each interval is uniformly distributed on [ 1 , 2 ]. it can be easily verified that the worst - case expected waiting time of this schedule is 0 . 5 , which is a significant ( 50 %) improvement over deterministic schedules . a simple calculation shows that the worst - case expected staleness of this schedule is just 0 . 25 . as detailed below , a lower waiting time can be achieved under the same staleness constraint . in what follows schedules s ={ x 1 , x 2 , . . . } in which all random variables x i are independent and identically distributed ( i . i . d .) are considered . such schedules are referred to as i . i . d . schedules . it can be shown that for any schedule s there exists an i . i . d . schedule s ′ which is at least as good as s , i . e ., ewt ( s ′)≦ ewt ( s ) and est ( s ′)≦ est ( s ). the objective in the design of data broadcast schedules can be summarized as follows : given a staleness constraint s , find a schedule s whose worst - case expected waiting time ewt ( s ) is minimal subject to the staleness constraint est ( s )≦ s . in what follows opt ( s ) denotes the minimum worst - case expected waiting time of a schedule that satisfies staleness constraint s . in the next theorem optimal universal schedules for small values of the staleness constraint s , i . e ., s ≦ 0 . 13 are presented . the proof of the following theorem is omitted , for the sake of clarity . theorem 1 let s ≦ 0 . 13 be a staleness constraint . then , the optimal schedule that satisfies s has distribution function f ( x )= min { 1 , c ( s ) e x } and yields worst - case expected waiting time c ( s ), where c ⁢ ⁢ ( s ) = ( 1 - 2 ⁢ s s + s ⁢ ⁢ ( 4 + s ) ) ⁢ ⁢ ⅇ - 1 2 ⁢ ⁢ ( s + s ⁢ ⁢ ( 4 + s ) ) in the next paragraphs , an approximation algorithm that computes optimal broadcast schedules will be detailed . the algorithm receives as input a staleness constraint s and any ( arbitrarily small ) approximation parameter ε , and returns a schedule s whose worst - case expected staleness is at most s and whose worst - case expected waiting time is at most opt ( s )+ ε . the computational complexity of the algorithm is polynomial in s /( ε ). the approximation algorithm has two steps . first , it is shown that for any ε 1 & gt ; 0 there exists a schedule s 1 such that est ( s 1 )≦ s , ewt ( s 1 )≦ opt ( s )+ ε 1 and the support of s 1 ( the support of a cumulative distribution function f ( x ) is a set of values of x at which the function is strictly less than 1 , i . e ., { x | f ( x )& lt ; 1 }) is bounded by s /( ε 1 ). in other words , the optimal distribution can be approximated by a distribution with bounded support . second , it is shown that for any ε 2 & gt ; 0 the schedule s 1 can be approximated by a schedule s 2 whose distribution is a piecewise - constant function that includes at most s /( ε 1 ε 2 ) segments . this schedule satisfies the staleness constraint , i . e ., est ( s 2 )≦ s , and its maximum waiting time is more than that of s 1 by at most ε 2 , i . e ., ewt ( s 2 )≦ ewt ( s 1 )+ ε 1 ≦ opt ( s )+ ε 1 + ε 2 . moreover , a linear program that computes s 2 can be formulated . the running time of this program is polynomial in s /( ε 2 ). as a result , for any ε & gt ; 0 a schedule that satisfies the staleness constraint s and whose maximal waiting time is at most opt ( s )+ ε can be computed . indeed , the above requirement can be satisfied by setting ε 1 = ε 2 =( ε )/ 2 . the optimal and approximation algorithm presented in the previous sections has been used in order to compute the attainable values of worst - case waiting time for a broad range of staleness constraints . the obtained results establish a trade - off between the staleness and waiting time of universal broadcast schedules . the trade - off is depicted on fig4 ( series a ). this trade - off has a surprising behavior referred to as the “ knee ” phenomenon : for small values of staleness ( typically below 0 . 3 ) the minimum waiting time decreases drastically with only a minor increase in the staleness constraint ; however , for large values of the staleness constraint ( above 0 . 3 ), any increase in the staleness constraint results in only a minor decrease of waiting time . a direct result of the knee phenomenon is the existence of a schedule that has small maximum expected waiting time ( 0 . 31 ) and whose worst - case expected staleness is also small ( at most 0 . 3 ). this point represents a reasonable trade - off between waiting time and staleness . the corresponding schedule reduces the worst - case waiting time by 70 % compared to a deterministic schedule while ensuring that the distributed information is up - to - date . the study of analytical ( closed form ) approximate solutions to the problem at hand gave rise to the following empirical observation . for arbitrary values of s , the distribution function g s ⁡ ( x ) = 1 - γ ⁢ ⁢ ( s + a ⁢ ⁢ ( s ) ) ⁢ ( s x + 1 - s - a ⁢ ⁢ ( s ) γ ⁢ ⁢ ( x + 1 ) ) yields worst - case waiting time which is very close to optimum . here a ( s ) is a constant between 0 and 1 , and γ ( x ) represents the standard gamma function . the staleness / waiting time trade - off of the schedules defined by g s ( x ) are depicted in fig4 ( series b ). now , broadcast schedules that use random time intervals and do not employ encoding are detailed . again , for clarity , it is assumed that the length of each packet is exactly one time unit . in the proposed schedule , each packet is transmitted over a time interval whose length is at least one time unit . in the first portion of the interval no data is being transmitted , and the packet is transmitted in the second portion of the interval . the random interval is defined by specifying , for each packet i , the amount of time that passes between the end of the transmission of packet i − 1 and the beginning of the transmission of packet i ( for simplicity , it is assumed that the transmission of packet 0 ends at time 0 ). this time is referred to as the interleaving time . for a real random variable x , f x ( t )= pr [ x & lt ; t ] denotes the distribution function of x and μ x =∫ 0 ∞ ( 1 − f x ( x )) dx denotes the expected value of x . this section uses a slightly different definitions of schedule , waiting time and histrory of the schedule the than those presented in section 1 ( paragraphs [ 0027 ] and [ 0031 ]). definition 5 [ schedule s ] a schedule is a sequence of random variables { x 1 , x 2 , . . . } such that x i is the interleaving time for packet i . a schedule s ={ x 1 , x 2 , . . . } can also be defined by its transmission sequence { t 1 , t 2 , . . . }, where t n represents the time in which packet n was transmitted . namely , t n = σ i = 1 n x i + n − 1 for all n ≧ 1 . let s be a schedule , and suppose that a client request is placed at time t . the client &# 39 ; s waiting time is defined as the time between t and beginning of the next packet . definition 6 [ waiting time , wt ( s , t )] the waiting time for a request at time t using a schedule s is defined to be wt ( s , t )= t n − t , where n is the first packet for which it holds that t n ≧ t . waiting time wt ( s , t ) is a random variable . the expectation of wt ( s , t ) is denoted by ewt ( s , t )= e [ wt ( s , t )]. a random schedule in which the variables x i are i . i . d . is closely related to so called renewal processes ( e . g ., [ 12 ], chapter xiii ). moreover , in this context , t n − t is well studied . nevertheless , the questions of interest regarding the expected worst - case value of wt ( s , t ) for general schedules have not been addressed in the literature . in order to capture the worst - case performance of the schedule , the notion of an adversary is used . this notion has being introduced in paragraph [ 0023 ]. in what follows , terms client and adversary are used interchangeably . in this paragraph a few definitions are introduced . a realization r of a random schedule s ={ x 1 , x 2 , . . . } is a deterministic schedule { x 1 , x 2 , . . . } that is in the domain of s . now , the notion of a history of s observed at time t is defined . let v t ( x 1 , . . . , x l ) be the event in which : ( a ) ∀ n ≦ l x n = x n , ( b ) σ i = 1 l x i + l − 1 ≦ t and ( c ) σ i = 1 l + 1 x i + l & gt ; t . that is , the event in which ( a ) for n ≦ 1 , the random variables x n are equal to x n , ( b ) the number of ( partial ) packets broadcasted until time t is at least l , and ( c ) the l + 1 &# 39 ; th package has not been transmitted up to time t . such an event is referred to as a history of s at time t . namely , any realization rεv t ( x 1 , . . . , x l ) is completely described up to time t by the interleaving times { x 1 , . . . , x l }. let v ( s , t ) be the set of possible histories of s at time t . finally , for any vεv ( s , t ) let s | v be the schedule distribution obtained by conditioning s on the event v . the goal is to design schedules that perform well with any behavior of incoming requests independently of what the viewed history of the channel was before the requests . in particular , the case in which the adversary is adaptive , i . e ., its behavior on time t depends on the history of the schedule up to time t is considered . the adversaries might have different degrees of adaptivity . the degree of adaptivity measures the ability of the adversary to generate request based on the history of the schedule . a lower degree of adaptivity corresponds to a more powerful adversary . definition 7 [ degree of adaptivity , ω ] an adversary is said to be ω - adaptive if its actions at time t are based on a history vεv ( s , t − ω ). the worst case expected waiting time of the schedule s on ω - adaptive adversaries , w ( s , ω ), is now defined as : max ⁢ ⁢ ( max 0 ≤ t ≤ ω ⁢ ewt ⁢ ⁢ ( s , t ) , max t ≥ ω ⁢ max v ∈ ϑ ⁢ ⁢ ( s , t - ω ) ⁢ e ⁢ [ wt ⁢ ⁢ ( s ⁢ | ⁢ v , t ) ] ) namely , w ( s , ω ) bounds the waiting time of a client no matter at what time t its request is placed or what the history of the schedule was at time t − ω . the first expression above addresses the case in which the client placed a request at time t & lt ; ω . this implies that the adversary has not based his request on prior knowledge of the schedule . in the following section , the design of scheduling strategies in the case in which our adversaries are ω - adaptive for ω = 1 is considered . a schedule s for which w ( s , 1 ) is strictly less than 1 is presented . namely , the schedule has w ⁢ ⁢ ( s , 1 ) = 1 2 . the presented schedule is of a simple nature as the random variables x 1 , x 2 , . . . that define it are independent and identically distributed ( i . e ., i . i . d .). this schedule is optimal . that is every other schedule s ′={ x ′ 1 , x ′ 2 , . . . } has a corresponding waiting time w ( s ′, 1 ) of value at least finally , it is shown that that the transmission rate of the optimal schedule is for larger values of r , a scheduling strategy is presented that has rate r , and worst case expected waiting time which is bounded by time units . the schedule is defined by i . i . d . random variables , and is the best possible under such a construction . in this section a schedule s that has an expected waiting time ( i . e ., w ( s , 1 )) which is bounded by is presented . the schedule is defined by a single random variable x . that is , s is defined to be { x 1 , x 2 , . . . }, where each random variable x i is independent and equals x . the random variable x has a simple structure , and is defined as follows . let μ & gt ; 0 be a parameter that will be fixed in a later stage of our discussion . let z be a “ random ” variable which obtains the value 0 with probability 1 . let u [ 0 , s ] be the uniform distribution on the interval [ 0 , s ]. finally let s = 2 ⁢ μ ⁢ ⁢ ( μ + 1 ) . in the schedule s ={ x 1 , x 2 , . . . } each random variable x i is independent and identically distributed . namely , x i = x for all i where x = pz +( 1 −) u [ 0 , s ]. it is not hard to verify that e [ x ]= μ , and that the support of x is [ 0 , s ]. the worst case expected waiting time of s is w ( s , 1 )= 1 + 2μ −√{ square root over ( 2μ ( μ + 1 ))}. specifically , setting w ⁢ ⁢ ( s , 1 ) = 1 2 . theorem 3 let s ={ x 1 , x 2 , . . . } be a schedule in which each random variable x n may be arbitrarily distributed and may depend on x i for i & lt ; n . then w ⁢ ⁢ ( s , 1 ) ≥ 1 2 . in some practical settings the transmission rate , i . e ., the average number of packets sent over a period of time , is important . indeed , along with clients that listen to the channel from time to time , there might be clients that monitor the information all the time . such clients prefer schedules with high transmission rates , which allow them to receive as many updates as possible . for universal schedules there exists a tradeoff between the transmission rate and minimum worst case waiting time . the transmission rate of a schedule s ={ x 1 , x 2 , . . . } is defined to be the expected amount of time in which the channel is in use . definition 8 [ transmission rate ] let r t be the expected number of packets sent in s ={ x 1 , x 2 , . . . } up to time t . the transmission rate r of s is defined to be in this section , schedules s which are defined by a series of i . i . d . random variables x are considered . for such schedules it can be shown ( e . g ., [ 12 ], chapter xiii ) that the transmission rate of s is where μ is the expectation of x . schedules of given rate r that have optimized waiting time ( with respect to the universal objective ) are presented . now , a schedule of rate r with expected waiting time of no more than time units is presented . the schedule is defined by i . i . d . random variables , and is the best possible under such a construction . the tradeoff between the transmission rate of our schedules and the worst case waiting time is depicted in fig5 . there exists a schedule s ={ x 1 , x 2 , . . . } with rate r and worst case expected waiting time of w ⁢ ⁢ ( s , 1 ) = 2 - r - 2 - 2 ⁢ r r . moreover , any schedule s ′={ x ′ 1 , x ′ 2 , . . . } in which x ′ n are i . i . d with rate r satisfies w ( s ′, 1 )≧ w ( s , 1 ). while several illustrative embodiments of the invention have been shown and described in the above description , numerous variations and alternative embodiments will occur to those skilled in the art . such variations and alternative embodiments are contemplated , and can be made without departing from the scope of the invention as defined in the appended claims . s . acharya , r . alonso , m . franklin , and s . zdonik . broadcast disks : data management for asymmetric communication environments . in proceedings of acm sigmod , pages 199 - 210 , san jose , calif ., may 1995 . m . h . ammar and j . w . wong . the design of teletext broadcast cycles . performance evaluation , 5 ( 4 ): 235 - 242 , 1985 . m . h . ammar . response time in a teletext system : an individual user &# 39 ; 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