Scheduling Algorithms for Optimizing Age of Information in Wireless Networks with Throughput Constraints

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1 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING Scheduling Algorithms for Optimizing Age of Information in Wireless Networks with Throughput Constraints Igor adota, Abhishek Sinha and Eytan odiano Abstract Age of Information (AoI) is a performance metric that captures the freshness of the information from the perspective of the destination. The AoI measures the time that elapsed since the generation of the packet that was most recently delivered to the destination. In this paper, we consider a singlehop wireless network with a number of nodes transmitting timesensitive information to a Base Station and address the problem of minimizing the Expected Weighted Sum AoI of the network while simultaneously satisfying timely-throughput constraints from the nodes. We develop four low-complexity transmission scheduling policies that attempt to minimize AoI subject to minimum throughput requirements and evaluate their performance against the optimal policy. In particular, we develop a randomized policy, a ax- Weight policy, a Drift-Plus-Penalty policy and a Whittle s Index policy, and show that they are guaranteed to be within a factor of two, four, two and eight, respectively, away from the minimum AoI possible. Simulation results show that ax-weight and Drift- Plus-Penalty outperform the other policies, both in terms of AoI and throughput, in every network configuration simulated, and achieve near optimal performance. Index Terms Age of Information, Throughput, Scheduling, Optimization, Quality of Service, Wireless Networks. I. INTRODUCTION The Age of Information (AoI) is a performance metric that measures the time that elapsed since the generation of the packet that was most recently delivered to the destination. This metric captures the freshness of the information from the perspective of the destination. Consider a cyber-physical system such as an automated industrial plant, a smart house or a modern car, where a number of sensors are transmitting timesensitive information to a monitor over unreliable wireless channels. Each sensor samples information from a physical phenomena (e.g. pressure of the tire, quantity of fuel, proximity to obstacles and engine rotational speed) and transmits this data to the monitor. Ideally, the monitor receives fresh information about every physical phenomena continuously. However, due to itations of the wireless channel, this is often impractical. In such cases, the system has to manage the use of the available channel resources in order to keep This work was supported by NSF Grants AST-54733, CNS-7375, and CNS-70964, and by Army Research Office (ARO) grant number W9NF Igor adota and Eytan odiano are with the Laboratory for Information & Decision Systems (LIDS), IT. Abhishek Sinha is currently working as a senior engineer at Qualcomm Research, San Diego. ( kadota@mit.edu; sinhaa@mit.edu; modiano@mit.edu) This paper was presented in part at IEEE INFOCO 08. the monitor updated. In this paper, we develop four lowcomplexity transmission scheduling policies and analyze their performance in terms of the freshness of the information at the monitor, namely the Age of Information. Let every packet be time-stamped with the time it was generated. Denote by τ i m the time-stamp of the mth packet delivered by sensor i to the monitor. Assume that at time t, the mth packet delivered by sensor i is the most recent. Then, the Age of Information associated with sensor i at time t is given by h i (t) = t τ i m. While the monitor does not receive new packets from sensor i, the value of h i (t) increases linearly with t, representing the information getting older. As soon as the monitor receives a new packet from sensor i, the corresponding time-stams instantaneously updated from τ i m to τ i m+, reducing the value of h i (t) by τ i m + τ i m. Notice that at the moment packet (m + ) is delivered to the monitor, the value of h i (t) matches the delay of the packet. This makes sense because, at that moment, the information at the monitor is as old as the information contained in packet (m + ). It follows naturally that a good AoI performance is achieved when packets with low delay are delivered regularly. In order to provide good AoI performance, the scheduling policy must control how the channel resources are allocated to the different sensors in the network. Depending on the channel conditions and network configuration, this can mean that some sensors get to transmit repeatedly, while other sensors less often. The frequency at which information is delivered to the monitor is of particular importance in sensor networks. Clearly, a sensor that measures the quantity of fuel requires a lower update frequency (i.e. throughput) than a sensor that is measuring the proximity to obstacles in order to avoid collisions. For capturing this attribute, we associate a minimum timely-throughput requirement with each sensor in the network. Hence, in addition to providing good AoI performance, the scheduling policies should also fulfill timelythroughput constraints from the individual sensors. A framework for modeling wireless networks with timelythroughput requirements was proposed in together with two debt-based scheduling policies that fulfill any feasible requirements. Generalizations of this model to different network configurations were proposed in 3 5. Scheduling policies that maximize throughput and also provide service regularity in wireless networks were studied in 6 8. The problem of minimizing AoI was introduced in 9. In 9 6, different queueing systems are analyzed and the optimal server utilization with respect to AoI is found. In 7 0, the

2 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING authors optimize the process of generating information updates in order to minimize AoI. The design of scheduling policies based on AoI is considered in 33. An important observation is that high throughput does not guarantee low AoI. Consider an // queue with high arrival rate and low service rate. In this system, the queue is often filled, resulting in high throughput and high packet delay. This high delay means that packets being served contain outdated information. Hence, despite the high throughput, the AoI may still be high. In this paper, we develop policies that minimize AoI subject to minimum throughput requirements, where timely-throughput is modeled as in. To the best of our knowledge, this is the first work to consider AoIbased policies that provably satisfy throughput constraints of multiple destinations simultaneously. The remainder of this paper is outlined as follows. In Sec. II, the network model and performance metrics are formally presented. Then, in Sec. III, four low-complexity scheduling policies are proposed and analyzed. In Sec. IV, those policies are simulated and compared to the state-of-the-art in the literature. The paper is concluded in Sec. V. II. SYSTE ODEL Consider a single-hop wireless network with a Base Station (BS) receiving time-sensitive information from nodes. Let the time be slotted, with slot index k,,,, and consider a wireless channel that allows at most one packet transmission per slot. In each slot k, the BS either idles or selects a node i,,, for transmission. Let u i (k) be the indicator function that is equal to when the BS selects node i during slot k, and u i (k) = 0 otherwise. When u i (k) =, node i samples fresh information, generates a new packet and sends this packet over the wireless channel. The packet from node i is successfully received by the BS with probability (0, and a transmission error occurs with probability. The probability does not change with time, but may differ between nodes. The transmission scheduling policy controls the decision of the BS in each slot k, which is represented by the set of values u i (k). The interference constraint associated with the wireless channel imposes that u i(k), k,,, () meaning that at any given slot k, the scheduling policy can select at most one node for transmission. Let d i (k) be the random variable that indicates when a packet from node i is delivered to the BS. If node i transmits a packet during slot k, i.e. u i (k) =, then d i (k) = with probability and d i (k) = 0 with probability. On the other hand, if node i does not transmit, i.e. u i (k) = 0, then d i (k) = 0 with probability one. It follows that E d i (k) u i (k) = u i (k) and, applying the law of iterated expectations E d i (k) = E u i (k). () In this paper, we consider non-anticipative scheduling policies, i.e. policies that do not use future knowledge in making decisions. Denote by Π the class of non-anticipative policies and let π Π be an arbitrary admissible policy. Our goal is to design low-complexity scheduling policies that belong to Π, provide close to optimal AoI performance and, at the same time, guarantee a minimum throughput level for each individual destination. Next, we formally introduce both performance metrics, throughput and AoI, and define a measure for closeness to optimality. A. inimum Throughput Requirement Let q i be a strictly positive real value that represents the minimum throughput requirement of node i. Using the random variable d π i (k), we define the long-term throughput of node i when policy π is employed as ˆq i π := Ed π i (k). (3) k= Then, we express the minimum throughput constraint of each individual node as ˆq π i q i, i,,. (4) In this paper, we assume that q i is a feasible set of minimum throughput requirements, i.e. there exists a policy π Π that satisfies all interference constraints in () and all throughput constraints in (4) simultaneously. As shown in, Lemma 5, the inequality q i, (5) is a necessary and sufficient condition for the feasibility of q i. Throughout this paper, we assume that (5) is satisfied with strict inequality. Next, we present the AoI metric. B. Age of Information The Age of Information depicts how old the information is from the perspective of the BS. Let h i (k) be the positive integer that represents the AoI associated with node i at the beginning of slot k. If the BS does not receive a packet from node i during slot k, then h i (k + ) = h i (k) +, since the information at the BS is one slot older. In contrast, if the BS receives a packet from node i during slot k, then h i (k+) =, because the received packet was generated at the beginning of slot k. The evolution of h i (k) follows h i (k + ) =, if d i (k) = ; h i (k) +, otherwise. The average AoI of node i during the first slots is captured by E k= h i(k) /, where the expectation is with respect to the randomness in the channel and the scheduling policy. For measuring the freshness of the information of the entire network when policy π is employed, we use the Expected Weighted Sum AoI E J π = E h i (k) h(), (7) k= where h() = h (),, h () T is the vector of initial AoI in (6) and > 0 is the weight of node i. For simplicity, we assume that h i () =, i, and omit h() henceforth. (6)

3 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING 3 C. Optimization Problem With the definitions of AoI and throughput, we present the optimization problem that is central to this paper. AoI Optimization OPT = min π Π E k= h i (k) (8a) s.t. ˆq i π q i, i ; (8b) u i(k), k. (8c) The minimum throughput constraints are depicted in (8b) and the interference constraints are in (8c). The scheduling policy that results from (8a)-(8c) is referred to as AoI-optimal. For a given network setup (,, q i, ), let OPT be the Expected Weighted Sum AoI achieved by the AoI-optimal policy π. Similarly, let OPT η be the AoI achieved by some policy η Π. The optimality ratio of η is given by ψ η = OPT η OPT, (9) and we say that policy η is ψ η -optimal. Naturally, the closer ψ η is to, the better is the AoI performance of policy η. The optimality ratio is used in the upcoming sections to compare the performance of different scheduling policies. III. SCHEDULING POLICIES In this section, we propose four low-complexity scheduling policies with strong AoI performances. The first three provably satisfy the throughput constraints for every feasible set q i and the fourth accounts for the throughput constraints, but provides no guarantee. To evaluate the AoI performance of each policy, we find their corresponding optimality ratio ψ η. oreover, in Sec. IV, we simulate and compare these policies to the state-of-the-art in the literature. Prior to introducing the policies, we obtain a lower bound to the AoI optimization (8a)-(8c) which is used in the derivation of the optimality ratios ψ η. Then, we present four scheduling policies: ) Optimal Stationary Randomized policy; ) ax- Weight policy; 3) Drift-Plus-Penalty policy; and 4) Whittle s Index policy. The first is obtained by solving the AoI optimization (8a)-(8c) over the class of Stationary Randomized Policies. The second and third policies are derived using Lyapunov Optimization 34. The fourth policy is obtained by using the Restless ulti-armed Bandit framework 35. A. Lower Bound In this section, we use a sample path argument to derive a lower bound to the AoI optimization (8a)-(8c). Theorem. The optimization problem in (0a)-(0c) provides a lower bound L B to the AoI optimization (8a)-(8c), namely L B OPT for every network setup (,, q i, ). Lower Bound L B = min π Π + ) (0a) ( ˆq i π s.t. ˆq i π q i, i ; (0b) u i(k), k. (0c) Proof. Consider a scheduling policy π Π that satisfies all throughput and interference constraints running on a network for the time-horizon of slots. Let Ω be the sample space associated with this network and let ω Ω be a sample path. For a given sample path ω, the total number of packets delivered by node i during the slots is denoted D i () = k= d i(k) and the inter-delivery time associated with each of those deliveries is denoted I i m. In particular, let I i m be the number of slots between the (m )th and mth packet deliveries from node i, m,, D i (). After the last packet delivery from node i, the number of remaining slots is R i. Hence, the time-horizon can be written as = D i() m= I i m + R i, i,,,. () According to the evolution of h i (k) in (6), the slot that follows the (m )th packet delivery from node i has an AoI of h i (k) =. Since the mth packet is delivered only after I i m slots, we know that h i (k) evolves as,,, I i m. This pattern is repeated throughout the entire time-horizon, including the last R i slots. As a result, the time-average Age of Information of node i can be expressed as h i (k) = k= = D i() D i() m= D i() D i () (I i m + )I i m m= + (R i + )R i Ii m + R i +, () where the last equality uses () to replace the two linear terms by. Define the operator x that computes the sample mean of any set x. In particular, let the sample mean of I i m and Ii m be I i = I i = D i() I i m ; (3) D i () m= D i() Ii m. (4) D i () m= Substituting I i into () and then applying Jensen s inequality, yields h i (k) ( Di () k= ( Ii ) R ) + i +, (5) Naturally, I i is the number of slots between the first packet delivery from node i and the first slot k =.

4 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING 4 combining () into (3) and then substituting the result in (5), gives h i (k) ( ( R i ) ) + R i D i () +. (6) k= By minimizing the LHS of (6) analytically with respect to the variable R i, we have h i (k) ( ) D i () + +. (7) k= Taking the expectation of (7) and applying Jensen s inequality, yields E h i (k) Di () k= E + +. (8) Applying the it to (8) and using the definition of throughput in (3), gives E h i (k) ( ) +. (9) k= Combining (9) and the objective function in (7), yields E J π = E h i (k) ˆq π i k= ) +. (0) ( ˆq i π Finally, substituting (0) into the AoI optimization (8a)-(8c) gives the Lower Bound (0a)-(0c). To obtain the expression in (0), we applied Jensen s inequality twice and minimized (6) analytically with respect to R i. Each of those steps could have led to a loose lower bound L B. However, in the next section, we use this lower bound to obtain a tight optimality ratio, ψ R <, for a Stationary Randomized policy. oreover, we evaluate the tightness of L B using numerical results in Sec. IV. B. Optimal Stationary Randomized policy Denote by Π R the class of Stationary Randomized Policies and let R Π R be a scheduling policy that, in each slot k, selects node i with probability µ i (0, and idles with probability µ idle. Each policy in Π R is fully characterized by the set of scheduling probabilities µ i, where µ i = Eu i (k), i, k and µ idle = µ i. Next, we find the Optimal Stationary Randomized policy R that solves the AoI optimization (8a)-(8c) over the class Π R Π and derive the associated optimality ratio ψ R. Proposition. Consider a policy R Π R with scheduling probabilities µ i. The long-term throughput and the expected time-average AoI of node i can be expressed as ˆq i R = ; () Eh i (k) =. () k= Proof. In any given slot k, the BS receives a packet from node i if this node is scheduled and the corresponding packet transmission is successful. The probability of this event is. oreover, the inter-delivery times I i m of node i are i.i.d. with PI i m = n = ( ) n, n,,. Clearly, under policy R, the sequence of packet deliveries is a renewal process. Thus, we can use renewal theory to derive () and (). In particular, by the definition of longterm throughput (3) and the expression for the expected timeaverage AoI of node i, we have ˆq R i = k= k= Ed i (k) (a) = EI i m = µ i ; (3) Eh i (k) (b) = EI i m EI i m + =. (4) where (a) follows from the elementary renewal theorem and (b) from its generalization for renewal-reward processes 36, Sec Substituting both expressions from Proposition into the AoI optimization (8a)-(8c) gives the equivalent optimization problem over the class Π R presented below. Optimization over Randomized policies OPT R = min R Π R s.t. q i, i ; µ i. (5a) (5b) (5c) Notice that under the class Π R, conditions (5c) and (8c) are equivalent. The Optimal Stationary Randomized policy R is characterized by the set µ i that solves (5a)-(5c). Theorem 3 (Optimality Ratio for R ). The optimality ratio of R is such that ψ R <, namely the Optimal Stationary Randomized policy is -optimal for every network setup. Proof. Let ˆq i L be the throughput associated with the policy that solves the Lower Bound (0a)-(0c). Consider the policy R Π R with long-term throughput ˆq i R = = ˆq i L for each node i. Since ˆq i R = ˆq i L, it follows that R satisfies all throughput constraints. Comparing L B in (0a) with the objective function associated with R, namely OPT R, yields OPT R < L B ψ R = OPT R OPT OPT R L B <, (6) where OPT comes from (8a) and OPT R from (5a). Recall that L B OPT OPT R OPT R. Corollary 4. The Optimal Stationary Randomized policy R is also the solution for the Lower Bound problem (0a)-(0c). Proof. Using the same argument as in the proof of Theorem 3, in particular ˆq i R = = ˆq i L, it follows that the scheduling policy that solves the Optimization over Randomized policies (5a)-(5c) also solves the Lower Bound (0a)-(0c). Theorem 5 (Optimal Stationary Randomized policy). The scheduling probabilities µ i that result from Algorithm

5 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING 5 are the unique solution to (5a)-(5c) and, thus, characterize the Optimal Stationary Randomized policy R. Algorithm Unique solution to T Conditions : γ i /qi, i,,, : γ max i γ i 3: µ i (q i / ) max ; γ i /γ, i 4: S µ + µ + + µ 5: while S < do 6: decrease γ slightly 7: repeat steps 3 and 4 to update µ i and S 8: end while 9: µ i = µ i, i, and γ = γ 0: return (µ, µ,, µ, γ ) Proof. To find the set of scheduling probabilities µ i that solve the optimization problem (5a)-(5c), we analyze the T Conditions. Let λ i be the T multipliers associated with the relaxation of (5b) and γ be the multiplier associated with the relaxation of (5c). Then, for λ i 0, i, γ 0 and µ i q i /,, i, we define L(µ i,λ i, γ) = + p i µ i ( ) + λ i (q i ) + γ µ i, (7) and, otherwise, we define L(µ i, λ i, γ) = +. Then, the T Conditions are (i) Stationarity: µi L(µ i, λ i, γ) = 0; (ii) Complementary Slackness: γ( µ i ) = 0; (iii) Complementary Slackness: λ i (q i ) = 0, i; (iv) Primal Feasibility: q i, i, and µ i ; (v) Dual Feasibility: λ i 0, i, and γ 0. Recall that q i is a strictly positive real value. Since L(µ i, λ i, γ) is a convex function, if there exists a vector (µ i, λ i, γ ) that satisfies all T Conditions, then this vector is unique. Hence, the scheduling policy R Π R that optimizes (5a)-(5c) is also unique and is characterized by µ i. Next, we find the vector (µ i, λ i, γ ). To assess stationarity, µi L(µ i, λ i, γ) = 0, we calculate the partial derivative of L(µ i, λ i, γ) with respect to µ i. It follows from the derivative that µ i + λ i = γ, i. (8) From complementary slackness, γ( µ i ) = 0, we know that either γ = 0 or µ i =. Equation (8) shows that the value of γ can only be zero if λ i = 0 and µ i, which violates µ i (0,. Hence, we obtain γ > 0 and µ i =. (9) Notice that µ i = implies in µ idle = 0. Based on dual feasibility, λ i 0, we can separate nodes i,, into two categories: nodes with λ i > 0 and nodes with λ i = 0. Category ) node i with λ i > 0. It follows from complementary slackness, λ i (q i ) = 0, that µ i = q i. (30) Plugging this value of µ i into (8) gives the inequality λ i = γ γ i > 0, where we define the constant γ i := q i. (3) Category ) node i with λ i = 0. It follows from (8) that ( ) qi γ = γ i µ i = q i γi γ. (3) In summary, for any fixed value of γ > 0, the scheduling probability of node i is µ i = q i γi max ;. (33) γ Notice that for a decreasing value of γ, the probability µ i remains fixed or increases. Our goal is to find the value of γ that gives µ i satisfying the condition µ i =. Proposed algorithm to find γ : start with γ = maxγ i. Then, according to (33), all nodes have µ i = q i / and, by the feasibility condition in (5), it follows that q i µ i =. (34) Now, by gradually decreasing γ and adjusting µ i according to (33), we can find the unique γ that fulfills µ i =. The solution γ exists since γ 0 implies in µ i. The uniqueness of γ follows from the monotonicity of µ i with respect to γ. This process is described in Algorithm and illustrated in Fig.. Algorithm outputs the set of scheduling probabilities µ i and the parameter γ. The set λ i is obtained using (8). Hence, the unique vector (µ i, λ i, γ ) that solves the T Conditions is found. In order to fulfill the throughput constraints (5b), every scheduling policy in Π R must allocate at least µ i q i / to each node i. What differentiates policies in Π R is how they Fig.. Illustration of Algorithm in a network with 3 nodes. On the left, the initial configuration with γ = maxγ i. On the right, the outcome γ implies that under policy R node will operate with minimum required scheduling probability µ = q /p, while the other two nodes will operate with a scheduling probability that is larger than the minimum.

6 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING 6 distribute the remaining resources, q i/, between nodes. According to Algorithm, the Optimal Stationary Randomized policy R supplies additional resources, µ i > q i/, to nodes with high value of γ i, namely nodes with a high priority or a low value of q i /. Notice that if a node with low q i / was given the minimum required amount of resources, it would rarely transmit and its AoI would be high. In contrast, policy R allocates the minimum required, µ i = q i/, to nodes with low priority or high q i /. The policies R Π R discussed in this section are as simple as possible. They select nodes randomly, according to fixed scheduling probabilities µ i calculated offline by Algorithm. Despite their simplicity, it was shown that R is -optimal regardless of the network setup (,, q i, ). In the following sections, we develop scheduling policies that take advantage of additional information, such as the current AoI of each node, for selecting nodes in an adaptive manner. C. ax-weight policy Using techniques from Lyapunov Optimization, we derive the ax-weight policy associated with the AoI optimization (8a)-(8c). ax-weight is a scheduling policy designed to reduce the expected increase in the Lyapunov Function. The Lyapunov Function outputs a positive scalar that is large when the network is in undesirable states, namely when nodes have high AoI or less throughput than the minimum required q i. Intuitively, the ax-weight policy keeps the network in desirable states by controlling the growth of the Lyapunov Function. Prior to presenting the ax-weight policy, we introduce the notions of throughput debt, network state, Lyapunov Function and Lyapunov Drift. Let x i (k) be the throughput debt associated with node i at the beginning of slot k. The throughput debt evolves as x i (k + ) = kq i k d i (t). (35) The value of kq i can be interpreted as the minimum number of packets that node i should have delivered by slot k + and k t= d i(t) is the total number of packets actually delivered in the same interval. Define the operator (.) + = max(.), 0 that computes the positive part of a scalar. Then, the positive part of the throughput debt is given by x + i (k) = maxx i(k); 0. A large debt x + i (k) indicates to the scheduling policy π Π that node i is lagging behind in terms of throughput. In fact, strong stability of the process x + i (k), namely t= Ex + i (k) <, (36) k= is sufficient to establish that the minimum throughput constraint, ˆq i π q i, is satisfied 34, Theorem.8. Denote by S k = (h i (k), x + i (k)) the network state at the beginning of slot k and define the Lyapunov Function by L(S k ) := ( h i (k) + V x + i (k) ), (37) where V is a (strictly) positive real value that depicts the importance of the throughput constraints. Observe that L(S k ) is large when nodes have high AoI or high throughput debt. To measure the expected change in the Lyapunov Function from one slot to the next, we define the Lyapunov Drift (S k ) := E L (S k+ ) L (S k ) S k. (38) The ax-weight policy is designed to keep L(S k ) small by reducing (S k ) in every slot k. Next, we present an upper bound on (S k ) that can be readily used to design the ax- Weight policy. The derivation of this upper bound is centered around the evolution of h i (k) in (6) and the evolution of x + i (k) in (35). The complete derivation can be found in Appendix A and the upper bound follows (S k ) E u i (k) S k W i (k) + B(k), (39) where W i (k) and B(k) are given by W i (k) = h i(k)h i (k) + + V x + i (k) ; (40) B(k) = h i (k) + + V x + i (k)q i +. (4) Both W i (k) and B(k) are fully characterized by the network state S k and network setup (,, q i, ). Hence, both can be used by admissible policies for making scheduling decisions. However, notice that the term B(k) in (39) is not affected by the choice of u i (k). Thus, for minimizing the upper bound in (39), the ax-weight policy selects, in each slot k, the node with highest value of W i (k), with ties being broken arbitrarily. Denote the ax-weight policy as W. Theorem 6. The ax-weight policy satisfies any feasible set of minimum throughput requirements q i. Theorem 7 (Optimality Ratio for W ). For any given network setup (,, q i, ), the optimality ratio of W is such that ψ W 4 + V. (4) L B In particular, for every network with V /, the ax-weight policy is 4-optimal. The proofs of Theorems 6 and 7 are provided in Appendices B and C, respectively. Both theorems follow from the analysis of the expression in (39). Recall that the Optimal Stationary Randomized policy R selects nodes randomly, according to fixed scheduling probabilities µ i. In contrast, the ax-weight policy W uses feedback from the network, namely h i (k) and x + i (k), to guide scheduling decisions. Despite the added complexity, we expect the feedback loop to improve the performance of W. In fact, numerical results in Sec. IV demonstrate that W outperforms R in every network configuration simulated. However, by comparing Theorems 3 and 7, it might seem that R yields a better performance than W. This is because the analysis associated with W is more challenging, leading to

7 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING 7 an optimality ratio ψ W that is less tight than ψ R. Next, we develop a policy called Drift-Plus-Penalty policy and show that it is -optimal. D. Drift-Plus-Penalty policy The Drift-Plus-Penalty policy is derived using a similar technique as the ax-weight policy. The main difference between these two policies is that the Drift-Plus-Penalty is designed to reduce the sum of the Lyapunov Drift and a Penalty Function, while the ax-weight policy reduces only the Lyapunov Drift. As we will see, this change will improve the optimality ratio of the Drift-Plus-Penalty policy significantly. Based on the AoI minimization (8a), we define the Penalty Function as follows P (S k ) := β i Eh i (k + ) S k, (43) where S k = (h i (k), x + i (k)) is the network state at the beginning of slot k and β i is a positive real value associated with node i. Similarly to (38), we define the Lyapunov Drift as (S k ) := E L (S k+ ) L (S k ) S k, (44) with associated Lyapunov Function L (S k ) := V x + i (k), (45) where V is a (strictly) positive real value that represents the importance of the throughput constraints. Notice that, as opposed to the Lyapunov Function in (37), the expression in (45) does not contain the term with h i (k). This is because the AoI term is already present in the Penalty Function. The Drift-Plus-Penalty policy is designed to minimize an upper bound on (S k ) + P (S k ) at every slot k. The upper bound is derived in Appendix D by manipulating (43)-(45). The expression for the upper bound follows (S k )+P (S k ) E u i (k) S k W i (k) + B (k), (46) where W i (k) and B (k) are given by W i (k) = β i h i(k) + V x + i (k) ; (47) B βi (k) = h i(k) + + V x + i (k)q i + V. (48) The values of W i (k) and B (k) can be easily calculated by any admissible policy and thus can be used for making scheduling decisions. For minimizing the upper bound in (46), the Drift-Plus-Penalty policy selects, in each slot k, the node with highest value of W i (k), with ties being broken arbitrarily. Denote the Drift-Plus-Penalty policy as DP P. Theorem 8. The DPP policy satisfies any feasible set of minimum throughput requirements q i. Theorem 9 (Optimality Ratio for DP P ). For any given network setup (,, q i, ), by choosing the constant β i = /µ i, the optimality ratio of DP P is such that ψ DP P + L B V. (49) In particular, for every network with V /, the Drift-Plus-Penalty policy is -optimal. The proofs of Theorems 8 and 9 are provided in Appendices E and F, respectively. The Lyapunov Function (45) with a quadratic term in x + i (k) has a central role in showing that the DP P policy satisfies any feasible requirements q i. Notice that this is also true for W. The Penalty Function (43) with a linear term in h i (k) is central to show that the DP P policy is -optimal. Recall that the W policy was designed based on a quadratic h i (k) and was shown to be 4-optimal. Comparing Theorems 7 and 9, we can clearly see this improvement in the optimality radio. However, this improvement is ited to the mathematical analysis. Numerical results in Sec. IV suggest that DP P and W have similar performances. Next, we develop an index policy based on Whittle s Index 35 that is surprisingly similar to W. E. Whittle s Index policy To find Whittle s Index, we transform the AoI optimization (8a)-(8c) into a relaxed Restless ulti-armed Bandit (RAB) problem. This is possible because every node in the network evolves as a restless bandit. To obtain the relaxed RAB problem, we first substitute the interference constraints in (8c) by the single time-averaged constraint k= Eu i (k). (50) Next, we relax this time-averaged constraint, by placing (50) into the objective function (8a) together with the associated Lagrange ultiplier C 0. The resulting optimization problem is called relaxed RAB and its solution lays the foundation for the design of Whittle s Index. A detailed description of this method can be found in 35, 37, 38. One of the challenges associated with this method is that Whittle s Index is only defined for problems that are indexable. Unfortunately, it can be shown that due to the throughput constraints, ˆq π i q i, the relaxed RAB resulting from the transformation of the AoI optimization is not indexable. To overcome this, we relax the throughput constraints (8b), placing them into the objective function of (8a)-(8c) as follows Relaxed AoI Optimization ÕPT = min π Π k= + θ i ( qi Eu i (k) s.t. θ i 0, i ; u i(k), k. E h i (k) + ) (5a) (5b) (5c)

8 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING 8 Each Lagrange ultiplier θ i is associated with a relaxation of ˆq i π q i. These multipliers are called throughput incentives for they represent the penalty incurred by scheduling policies that deviate from the corresponding throughput constraint. Applying the transformation described at the beginning of this section to the relaxed AoI optimization (5a)-(5c) yields Doubly relaxed RAB ÕPT D = min π Π k= + (C θ i )E u i (k) C + θ iq i s.t. θ i 0, i ; C 0. E h i (k) + (5a) (5b) (5c) Next, we solve the doubly relaxed RAB, establish that the relaxed AoI optimization is indexable and obtain a closed-form expression for the Whittle s Index. The doubly relaxed RAB is separable and thus can be solved for each individual node. Observe that a scheduling policy running on a network with a single node i can only choose between selecting node i for transmission during slot k or idling. The scheduling policy that optimizes (5a)-(5c) for a given node i is characterized next. Proposition 0 (Threshold policy). Consider the doubly relaxed RAB problem (5a)-(5c) associated with a single node i. The optimal scheduling policy is a Threshold policy that, in each slot k, selects node i when h i (k) H i and idles when h i (k) < H i. For positive fixed values of C and θ i, if C > θ i, the expression for the threshold is H i = 3 + ( Otherwise, if C θ i, the threshold is H i =. ) + (C θ i). (53) Proposition 0 follows from 3, Propostion 4. Next, we define the condition for indexability and establish that the relaxed AoI optimization is indexable. For a given value of C, let I i (C) = h i (k) N h i (k) < H i be the set of states h i (k) in which the Threshold policy idles. The doubly relaxed RAB associated with node i is indexable if the set I i (C) increases monotonically from to N, as the value of C increases from 0 to +. Furthermore, the relaxed AoI optimization is indexable if this condition holds for all nodes. The condition on I i (C) follows directly from Proposition 0 and is true for all nodes i. Thus, we establish that the relaxed AoI optimization (5a)-(5c) is indexable. Given indexability, we define Whittle s Index. Let C i (h i (k)) be the Whittle s Index associated with node i in state h i (k). By definition, C i (h i (k)) is the infimum value of C that makes both scheduling decisions (transmit or idle) equally desirable to the Threshold policy while in state h i (k). The scheduling decisions are equally desirable when the multiplier C is such that H i = h i (k) +. Using (53) to solve this equation for the value of C gives the following expression for the Index C i (h i (k)) = α i h i(k) h i (k) + + θ i. (54) After establishing indexability and obtaining the expression for C i (h i (k)), we define Whittle s Index policy. The Whittle s Index policy selects, in each slot k, the node with highest value of C i (h i (k)), with ties being broken arbitrarily. Denote the Whittle s Index policy as W I. Theorem (Optimality Ratio for W I). For any given network setup (,, q i, ), the optimality ratio of W I is such that ψ W I 8 + L B θ i 7. (55) In particular, for every network with θ i 7 /, the Whittle s Index policy is 8-optimal. The proof of Theorem is provided in Appendix G. The arguments used for deriving ψ W I are analogous to the ones for deriving ψ W in Theorem 7. Those similarities come from the fact that policies W and W I are almost identical. Comparing the expressions for W i (k) and C i (h i (k)), in (40) and (54), respectively, we can see that both have the term h i (k)/ and both have an isolated throughput term: W i (k) has V x + i (k) and C i(h i (k)) has θ i. Naturally, we expect the performance of both policies to be similar in terms of AoI. The key difference between W and W I lies in the throughput term. While the term V x + i (k) guarantees that W satisfies the throughput constraint, ˆq i π q i, the positive scalar θ i represents an incentive for W I to comply with the constraint, but provides no guarantee. The benefit of using a fixed θ i is that there is no need to keep track of x + i (k) for each node and at every slot k. The results in this section hold for any given set of positive throughput incentives θ i. Next, we propose an algorithm that finds the values of θ i which maximize a lower bound on the Lagrange Dual problem associated with the relaxed AoI optimization (5a)-(5c). Observe that ÕPT D in (5a) is the Lagrange Dual function associated with (5a)-(5c). Thus, we can define the Lagrange Dual problem as max ÕPT C,θi D subject to C 0 and θ i 0, i. Since this dual problem is challenging to address, we consider a lower bound: max L(C, χ i ) max ÕPT D OPT. (56) C,χ i C,θ i subject to χ i = C θ i, C 0 and θ i 0 for all nodes i, where L(C,χ i ) = C q i + (57) p i p i + χ i + χ iq i. The throughput incentives θ i that result from the maximization of L(C, χ i ) are given by Algorithm. They are used in the

9 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING 9 next section to simulate the Whittle s Index policy. Simulation results show that the values of θi from Algorithm reduce the throughput debt when compared to θ i = 0. Algorithm Throughput Incentives : χ i (/q i ) (/ /) /, i : C max i χ i 3: φ i minc; χi /( ) + (/ /), i 4: S φ + φ + + φ 5: while S < do 6: decrease C slightly 7: repeat steps 3 and 4 to update φ i and S 8: end while 9: C = C and χ i = minc ; χ i and θi = C χ i, i 0: return (θ, θ,, θ ) the throughput debt. Figs. and 4 show that the performance of W, DP P, W I and W P are comparable to the lower bound. Since the lower bound is only associated with policies that fulfill the throughput requirements, we conclude that the performances of W and DP P are close to optimal. V. CONCLUDING REARS In this paper, we considered a single-hop wireless network with a number of nodes transmitting time-sensitive information to a Base Station over unreliable channels. We addressed the problem of minimizing the Expected Weighted Sum AoI of the network while satisfying minimum throughput requirements IV. SIULATION RESULTS In this section, we simulate six transmission scheduling policies: ) Optimal Randomized, R ; ) ax-weight, W ; 3) Drift-Plus-Penalty 3, DP P ; 4) Whittle s Index, W I; 5) Largest Weighted-Debt First, LD; and 6) Whittle s Index without throughput constraints, W P. The first four policies are developed in Sec. III and the last two are proposed in and 3, respectively. Policy LD selects, in each slot k, the node with highest value of x i (k)/, where x i (k) is the throughput debt (35). It was shown in that LD satisfies any set of feasible throughput requirements q i. Notice that LD does not account for AoI. Policy W P was proposed in 3 for minimizing the AoI in broadcast wireless networks. It is analogous to W I but with θ i = 0, i and it does not account for minimum throughput requirements. We simulate a network with nodes, each having different parameters. Node i has weight = ( + i)/, channel reliability = i/ and minimum throughput requirement q i = ε /, where ε 0, ) represents the hardness of satisfying the throughput constraints ˆq i π q i. The larger the value of ε, the more challenging are the constraints. Notice that ε < is necessary for the feasibility of q i. Each simulation runs for a total of = 0 6 slots. Figs. and 3 show simulation results of networks with different sizes, namely 5, 0,, 5, 30, while Fig. 4 shows networks with varying throughput constraints, in particular ε 0.7, 0.75,, 0.9, 0.95, Two performance metrics are used to evaluate scheduling policies. Figs. and 4 measure the Expected Weighted Sum AoI, EJ π, defined in (7) and compare it with the lower bound L B in (0a). Fig. 3 measures the maximum normalized throughput debt, defined as max i x + i ( + )/q i. Each data point in Figs., 3 and 4 is an average over the results of 0 simulations. Our results clearly demonstrate the superior performance of the ax-weight and Drift-Plus-Penalty policies. Fig. 3 shows that, as expected, only W I and W P violate the throughput requirements. Nevertheless, by comparing W I and W P, it is evident that the incentives θi from Algorithm reduced The ax-weight policy is simulated with V =. 3 The Drift-Plus Penalty policy is simulated with V = /. Fig.. Simulation of a network with fixed ε = 0.9 and varying size. Fig. 3. Simulation of a network with fixed ε = 0.9 and varying size. Fig. 4. Simulation of a network with size = 30 and varying hardness ε.

10 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING 0 from the individual nodes. Four low-complexity scheduling policies were developed: Optimal Stationary Randomized policy, ax-weight policy, Drift-Plus-Penalty policy, and Whittle s Index policy. The performance of each policy was evaluated both analytically and through simulation. The ax- Weight policy and the Drift-Plus-Penalty policy demonstrated the best performance in terms of both AoI and throughput. Interesting extensions include consideration of unknown channel probabilities and periodic generation of packets. APPENDIX A UPPER BOUD ON THE LYAPUNOV DRIFT OF W In this appendix, we obtain the expressions in (39)-(4), which represent an upper bound on the Lyapunov Drift. Consider the network state S k = (h i (k), x + i (k)), the Lyapunov Function L(S k ) in (37) and the Lyapunov Drift (S k ) in (38). Substituting (37) into (38), we get (S k ) = E h i (k + ) h i (k) S k + (58) + V E x + i (k + ) x + i (k) S k. Next, we find expressions for x + i (k + ) x + i (k) and h i (k + ) h i (k) which are then substituted into (58). To obtain the expression associated with the throughput debt, we use the following recursion x i (k + ) = x i (k) d i (k) + q i, k, (59) with x i () = 0. Notice that (59) is equivalent to (35). Squaring x + i (k + ), yields x + i (k + ) = maxxi (k) d i (k) + q i ; 0 anipulating (60), gives maxx + i (k) d i(k) + q i ; 0 x + i (k) d i(k) + q i. (60) x + i (k + ) x + i (k) x + i (k)d i(k) q i +. (6) Finally, by taking the conditional expectation of (6) and applying (), we get the upper bound E x + i (k + ) x + i (k) Sk (6) x + i (k) (Eu i (k) S k q i ) +. To obtain the expression associated with the AoI, we calculate Eh i (k + ) S k using the evolution of h i (k) in (6). It follows that E h i (k + ) Sk = pi E u i (k) S k + anipulating (63), we get + (h i (k) + ) ( E u i (k) S k ). (63) E h i (k + ) h i (k) Sk = (64) = E u i (k) S k h i (k) h i (k) + + h i (k) +. Substituting (6) and (64) into the Lyapunov Drift in (58), yields the expressions in (39)-(4). APPENDIX B PROOF OF THEORE 6 Theorem 6. The ax-weight policy satisfies any feasible set of minimum throughput requirements q i. Proof. The expression for the Lyapunov Drift (39) is central to the analysis in this appendix and is rewritten below for convenience. (S k ) E u i (k) S k W i (k) + B(k), where W i (k) and B(k) are given by W i (k) = h i(k)h i (k) + + V x + i (k) ; ( B(k) = h i (k) + ) ( + V x + i (k)q i + ). Recall that the ax-weight policy minimizes the RHS of (39) by selecting i = arg maxw i (k) in every slot k. Hence, any other policy π Π yields a lower (or equal) RHS. Consider a Stationary Randomized Policy R Π R that, in each slot k, selects node i with probability µ i (0,. Then, it follows that E u i (k) S k W i (k) µ i W i (k). (65) Substituting (65) into the equation of the Lyapunov Drift gives (S k ) µ i W i (k) + B(k) + V V h i (k) (µ i q i )x + i (k). (66) Consider the Cauchy-Schwarz inequality h i (k) + p i µ i h i (k) +. (67) Applying this inequality to (66) yields (S k ) V (µ i q i )x + i (k)+ (68) p i µ i + V h i (k) +

11 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING and rearranging the terms V (µ i q i )x + i p i µ (k) + i + h i (k) + p i µ i (S k )+ p i µ i + + V. (69) For simplicity of exposition, we divide inequality (69) into four terms LHS + LHS RHS + RHS. Taking their expectation with respect to S k, summing them over k,,, and then dividing them by, gives LHS = (70) V = (µ i q i )E x + i (k) k= LHS = (7) = E h i (k) + k= RHS = RHS = E (S k ) (7) k= + V. (73) From the definition of Lyapunov Drift (38) and the fact that the Lyapunov Function (37) is non-negative, the expression of RHS can be simplified as follows RHS L(S ), (74) recall that h i () = and x i () = 0. Hence, the Lyapunov Function L(S ) is a positive finite constant. Since LHS is non-negative, it follows that the inequality can be reduced to LHS RHS + RHS. Using (74) and applying the it yields (µ i q i ) V E x + i (k) k= + V (75) Hence, by rearranging the terms, we can show that for any given node i, we have strong stability E x + i (k) <, (76) k= what establishes condition (36). APPENDIX C PROOF OF THEORE 7 Theorem 7 (Optimality Ratio for W ). For any given network setup (,, q i, ), the optimality ratio of W is such that ψ W 4 + V. (77) L B In particular, for every network with V /, the ax-weight policy is 4-optimal. Proof. Consider the analysis in Appendix B. In particular, the inequality LHS + LHS RHS + RHS presented in (70)-(73). Applying Jensen s inequality twice to LHS, yields E k= E J W ( h i (k) + ) LHS ( ) LHS. (78) Since LHS is non-negative, it follows that the inequality can be reduced to LHS RHS + RHS. Using equations (74) and (78), and then applying the it yields ) ( p i µ i E J W E J W OPT W + ( ) ( + V + V + V (79) Analogously to the proof of Theorem 3, let ˆq i L be the longterm throughput associated with the policy that solves the Lower Bound optimization (0a)-(0c). Then, evaluating L B from (0a) gives L B = +. (80) ˆq L i Now, for each node i, we impose the following scheduling probability µ i = ˆq i L/. Then, evaluating (79) gives OPT W + V. (8) ˆq L i Comparing (80) and (8), yields L B OPT W 4L B + V ; (8) ψ W 4 + V, (83) L B )

12 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING what establishes the expression in (4). Substituting (87) into the RHS of (46) gives APPENDIX D UPPER BOUD ON THE LYAPUNOV DRIFT OF DPP In this appendix, we derive the expressions in (46)-(48), which represent an upper bound on (S k ) + P (S k ). This derivation is similar to the one in Appendix A. First, we analyze the Lyapunov Drift (44). Substituting the Lyapunov Function (45) into the Drift gives (S k ) = V E x + i (k + ) x + i (k) S k. (84) Then, from the bound in (6), results (S k ) V x + i (k) (Eu i (k) S k q i ) + V /. (85) Now, we analyze the Penalty Function (43) by utilizing the evolution of h i (k) in (6) and the expression in (). It follows that P (S k ) := β i Eh i (k + ) S k = = β i h i (k) + h i (k)ed i (k) S k β i h i (k) + h i (k)eu i (k) S k. (86) Substituting (85) and (86) into (S k ) + P (S k ) yields the expressions in (46)-(48). APPENDIX E PROOF OF THEORE 8 Theorem 8. The DPP policy satisfies any feasible set of minimum throughput requirements q i. Proof. The expression of the upper bound in (46) is central to the analysis in this appendix and is rewritten below for convenience. (S k )+P (S k ) E u i (k) S k W i (k) + B (k), where W i (k) and B (k) are given by W i (k) = β i h i(k) + V x + i (k) ; B βi (k) = h i(k) + + V x + i (k)q i + V. Recall that the Drift-Plus-Penalty policy is designed to minimize the RHS of (46). Hence, a Stationary Randomized Policy R Π R that, in each slot k, selects node i with probability µ i (0, yields a lower (or equal) RHS, i.e. E u i (k) S k W i (k) µ i W i (k). (87) (S k ) + P (S k ) µ i W i (k) + B (k) V x + i (k)µ i q i + V + β i + + β i h i (k) µ i, (88) and by substituting the expression of P (S k ) and rearranging the terms, we get V x + i (k)µ i q i + β i h i (k)µ i V + β i (S k )+ β i Eh i (k + ) h i (k) S k, (89) For simplicity of exposition, we divide inequality (89) into five terms LHS +LHS RHS +RHS +RHS 3. Taking their expectation with respect to S k, summing them over k,,, and then dividing them by, gives LHS = (µ i q i ) V Ex + i (k) (90) LHS = RHS = V (β i µ i ) RHS = k= Eh i (k) (9) k= V + β i (9) E x + i () x + i ( + ) (93) RHS 3 = β i Eh i () h i ( + ). (94) Notice that the expression of the Lyapunov Drift (84) was utilized in RHS. Since h i ( + ) and x + i ( + ) are nonnegative, the expression of RHS + RHS 3 can be simplified as follows RHS +RHS 3 E V x + i () + β i h i () (95) Recall that h i () = and x i () = 0. Hence, in the it, we have RHS + RHS 3 0.

13 SUBITTED TO IEEE/AC TRANSACTIONS ON NETWORING 3 Since LHS is non-negative, it follows that the inequality can be reduced to LHS RHS + RHS + RHS 3. Applying the it and using (95) yields V (µ i q i ) Ex + i (k) V + β i k= (96) By rearranging the terms, it is easy to see that strong stability holds for any given node i, i.e. E x + i (k) <, (97) k= what establishes condition (36). APPENDIX F PROOF OF THEORE 9 Theorem 9 (Optimality Ratio for DP P ). For any given network setup (,, q i, ), by choosing the constant β i = /µ i, the optimality ratio of DP P is such that ψ DP P + L B V. (98) In particular, for every network with V /, the Drift-Plus-Penalty policy is -optimal. Proof. Consider the analysis in Appendix E. In particular, the inequality LHS + LHS RHS + RHS + RHS 3 presented in (90)-(94). Given that LHS is non-negative, it follows that the inequality can be reduced to LHS RHS + RHS + RHS 3. Applying the it and using (95) yields (β i µ i ) Eh i (k) V + β i (99) k= Analogously to the proof of Theorem 3, let ˆq i L be the longterm throughput associated with the policy that solves the Lower Bound optimization (0a)-(0c). Then, evaluating L B from (0a) gives L B = +. (00) ˆq L i Now, for each node i, we impose the following scheduling probability µ i = ˆq i L/ and constant β i = /ˆq i L. Then, evaluating (99) gives OPT DP P + V. (0) ˆq L i Comparing (00) and (0), yields L B OPT DP P L B + V ; (0) ψ DP P + V. (03) L B Recall from Corollary 4 that ˆq i L = µ i. Hence, we know that β i = /µ i, i. The proof is complete. APPENDIX G PROOF OF THEORE Theorem (Optimality Ratio for W I). For any given network setup (,, q i, ), the optimality ratio of W I is such that ψ W I 8 + L B θ i 7. In particular, for every network with θ i 7 /, the Whittle s Index policy is 8-optimal. Proof. Whittle s Index policy selects, in each slot k, the node with highest value of C i (h i (k)). It is easy to see that this choice maximizes E u i (k) S k C i (h i (k)), in every slot k. From this perspective, the difference between W I and W is only the tern multiplying E u i (k) S k. Thus, if we find an upper bound to the Lyapunov Drift (S k ) that has the Whittle s Index policy as its minimizer, then similar arguments as the ones utilized in Appendix C can be used to derive an optimality ratio for W I. The upper bound associated with the ax-weight policy (39) is rewritten below for V = 0 (S k ) E u i (k) S k W i (k) + B(k), where W i (k) and B(k) are given by W i (k) = h i(k)h i (k) + ; B(k) = h i (k) +. We can manipulate this upper bound as follows (S k ) where + E u i (k) S k C i (h i (k)) + B(k)+ E u i (k) S k C i (h i (k)) W i (k), (04) C i (h i (k)) W i (k) = α i h i(k) + θ i h i(k) h i (k) + θ i. (05)