text:
My Cart
Store Locator
VIP Membership Terms: Normally $24.99 per year, but your Web Exclusive deal means you enjoy 3 months of VIP Family Benefits for $1.99. Check out your savings and benefits below.
Cancellation: Rest assured you can cancel your new membership at any time during the year at a pro-rated rate, simply call your VIP Hotline at 800.543.7309 or click here to submit your cancellation request. Benefits and prices are subject to change.
Want more info? Click here and read 'til your heart's content or call your VIP Hotline at 800.543.7309
Womens Brooks Addiction 10 Running Shoe - White/Light Blue 5.5
Available in White/Light Blue
Women's BROOKS� ADDICTION 10 :: Your flat arches are no longer an issue with the women's Brooks� Addiction? 10. You'll love the support and control of these trainers, specially designed for runners with low arches. You'll notice the difference, thanks to the Addiction? 10's multiple layers of cushy comfort that ...
Womens Brooks PureConnect 4 Running Shoe - Black/Pink 5.5
Available in Black/Pink
Women's Brooks®: PureConnect 4 :: Let your feet sing with light-as-air design and fully loaded flex. The PureConnect 4 gives you just enough shoe to go where the turns are tight and spontaneous breakaways are encouraged. With and enhanced upper featuring an open vamp design, track spike tongue and wrap ...
Womens Brooks PureFlow 3 Running Shoe - Electric/Caribbean 8.5
Available in Electric/Caribbean
Women's BROOKS� PUREFLOW 3 :: Get the same unique combination of ultra-light construction and super plush cushioning that you loved on the Women's Brooks� PureFlow 2, while getting a closer, more comfortable fit from the newly updated PureFlow 3. You'll stay light on your feet while securing soft landings on ...
Womens Brooks Transcend 2 Running Shoe - Teal 13
Available in Teal
Women's BROOKS� TRANSCEND 2 :: Rise above it all, with the next evolution of the Float experience, the Women's Brooks� Transcend 2. Score your smoothest, most comfortable strides ever in the newly designed upper that works in perfect harmony with the adaptable cushioning of the Super DNA midsole. You'll feel ...
Womens Brooks PureCadence 2 Running Shoe - Anthracite/Black 6
Available in Anthracite/Black
Women's BROOKS� PURECADENCE 2 :: Your favorite natural running shoe is back better than ever! You'll love the comfort-driven updates to the women's Brooks� PureCadence 2. Enjoy all the support you need to run strong while still experiencing the light-on-your feet feel of this minimalist running shoe. The hugs just ...
Womens Brooks PR LD Track and Field Shoe - White/Gemma 11
Available in White/Gemma
Women's Brooks� PR LD :: Behold the lightest spike of them all, designed for eyebrow-raising efficiency over long distances. A super thin mono mesh upper keeps the foot securely wrapped and supremely ventilated. With a newly engineered spike plate designed to flex with the foot and deliver optimal traction, your ...
Womens Brooks PR MD 54.26 Track and Field Shoe - Fiery Coral 9.5
Available in Fiery Coral
Women's Brooks� PR MD 54.26 :: Go the mid-distance with the new and super light PR MD. We improved the medial side with a built-in stretch material to secure the arch--perfect for when your eyes are on the finish tape. New, airy mesh keeps the weight down and increases breathability ...
Womens Brooks Glycerin 10 Running Shoe - Anthracite/White 5.5
Available in Anthracite/White
Women's BROOKS� GLYCERIN� 10 :: Enjoy the plush, balanced and smooth feel of the new Brooks� Glycerin� 10 trainers. For a ride that feels truly decadent, slip on a pair of these beauties and hit the road in style. Your feet will appreciate the Glycerin� 10's range of comfort features, ...
Womens Brooks Mach 13 Racing Shoe - Flame Orange/Varsity Maize 9
Available in Flame Orange/Varsity Maize
Women's Brooks� Mach 13 :: As ounces drop, so do times. Lighter than ever but still eager to tear up any terrain, the re-tuned Mach 13 Spike is sure to make headlines. With a made-for-you fit and feel, this newly streamlined racing shoe is designed to take everything out of ...
Womens Brooks Mach 12 Spike Racing Shoe - Imperial Purple/White 8.5
Available in Imperial Purple/White
Women's Brooks� Mach 12 Spike :: As ounces drop, so do times. Lighter than ever but still eager to tear up any terrain, the re-tuned Mach XII Spike is sure to make headlines. With a made-for-you fit and feel, this newly streamlined racing shoe is designed to take everything out ...
Womens Brooks PureFlow 3 Nightlife Running Shoe - Bright Pink/Black 10
Available in Bright Pink/Black
Women's Brooks� PureFlow 3 Night Life :: Get the same unique combination of ultra-light construction and super plush cushioning in a Night Life while getting a closer, more comfortable fit from the newly updated PureFlow 3. You'll stay light on your feet while securing soft landings on every step in ...
Womens Brooks Cascadia 8 Trail Running Shoe - Black/Pink 10
Women's BROOKS� CASCADIA 8 :: Tackle the road less traveled in the rugged women's Brooks� Cascadia 8, designed with gritty, wet and uneven ground in mind. Adapt to any surface with the 4-Point Heel and Forefoot Pivot Post System that helps keep neutral footing on uneven ground. Charge ahead in ...
Womens Brooks Cascadia 9 Trail Running Shoe - Fuschia 6.5
Available in Fuschia
Women's BROOKS� CASCADIA 9 :: Run strong and sure-footed through even the roughest trails, protected in the newly updated Women's Brooks� Cascadia 9, now designed to keep you more comfortable tackling gritty, wet and uneven ground. Secure more natural balance and a connected-to-the-ground feel from the slightly lowered heel, while ...
Womens Brooks Mach 14 Spikeless Cross Training Shoe - Neon Magenta/Neon Blue 12
Available in Neon Magenta/Neon Blue
Women's Brooks� Mach 14 Spikeless :: Runs like the wind, and weighs just as much. The new Mach 14 Spikeless keeps it light with an internal skeletal structure that secures the foot using minimal materials. Screen print design details get you killer looks without the weight, though exceptional traction means ...
ENJOY YOUR BIGGEST BENEFITS EVERY DAY Get FREE Shipping every day • Your WARP SPEED PROMISE • Ultimate Test Run 90-Day Wear 'em & Love 'em Guarantee • tons more perks! JOIN NOW!! �
Get the Latest, Sign Up Now!
Become a Fan & Follow Us.


text:
Real-Time Model Checking is Really Simple Leslie Lamport Microsoft Research 13 June 2005 To be presented at Charme 2005, to be held 3-6 October 2005, u in Saarbr¨cken, Germany. This version is formatted differently, with different line and page breaks, from the version to appear in the Charme proceedings. Contents 1 Introduction 1 Introduction 1 2 Writing Explicit-Time Specifications 2 3 Model Checking Explicit-Time Specifications 4 3.1 Specifications and Temporal Properties . . . . . . . . . . . . 4 3.2 Symmetry . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 5 3.3 Model Checking . . . . . . . . . . . . . . . . . . . . . . . . . . 5 3.4 View Symmetry . . . . . . . . . . . . . . . . . . . . . . . . . . 7 3.5 Symmetry Under Time Translation . . . . . . . . . . . . . . . 7 3.6 Periodicity and Zeno Behaviors . . . . . . . . . . . . . . . . . 8 4 Comparison with Uppaal 10 4.1 The Leader Algorithm . . . . . . . . . . . . . . . . . . . . . . 10 4.2 Fischer's Algorithm . . . . . . . . . . . . . . . . . . . . . . . . 14 5 Conclusion 15 References 16 Real-Time Model Checking is Really Simple Leslie Lamport Microsoft Research 13 June 2005 Abstract It is easy to write and verify real-time specifications with existing languages and methods; one just represents time as an ordinary vari- able and expresses timing requirements with special timer variables. The resulting specifications can be verified with an ordinary model checker. This basic idea and some less obvious details are explained, and results are presented for two examples. 1 Introduction Numerous special languages and logics have been proposed for specifying and verifying real-time algorithms. There is an alternative that I call the explicit-time approach, in which the current time is represented as the value of a variable now and the passage of time is modeled by a Tick action that increments now . Timing constraints are expressed with timer variables. Hardly anything has been written about the explicit-time approach, per- haps because it is so simple and obvious. As a result, most people seem to believe that they must use special real-time languages and logics. It has already been shown that an explicit-time approach works fine for specifying and proving properties of real-time algorithms [1]. Here, I consider model checking explicit-time specifications. The major advantage of the explicit-time approach is that it can be used with any language and logic for describing concurrent algorithms. This is especially important for complex algorithms that can be quite difficult to represent in the lower-level, inexpressive languages typical of real-time model checkers. For example, distributed message-passing algorithms have queues or sets of messages in transit, each with a bound on its delivery time. Such algorithms are difficult or impossible to handle with most real-time model checkers. Section 2 briefly explains the explicit-time approach with a simple distributed algorithm. A complete specification of the algorithm in TLA+ [8], a high-level mathematics-based language, appears in [9]. Explicit-time descriptions can use either continuous or discrete time. Section 3 shows that when discrete time is used, these descriptions can be checked with ordinary model checkers. This simple fact has been known for quite a while and is implicit in several published results [5]. However, a direct statement of it does not seem to have appeared before in print. Moreover, there are some aspects of model checking explicit-time specifications that may not be obvious, including the use of view symmetry and a method for checking that a specification is nonZeno [1]. Section 4 describes the result of checking the algorithm described in Section 2 with TLC, a model checker for TLA+ specifications, and with Uppaal [10], the only real-time model checker I know of that can handle this example. It also compares TLC, Spin [6], and SMV [11] with Uppaal on the Fischer mutual exclusion algorithm [13]. More details appear in [9]. 2 Writing Explicit-Time Specifications In an explicit-time specification, time is represented with a variable now that is incremented by a Tick action. For a continuous-time specification, Tick might increment now by any real number; for a discrete-time specification, it increments now by 1. Timing bounds on actions are specified with one of three kinds of timer variables: a countdown timer is decremented by the Tick action, a count-up timer is incremented by Tick , and an expiration timer is left unchanged by Tick .1 A countdown or count-up timer expires when its value reaches some value; an expiration timer expires when its value minus now reaches some value. An upper-bound timing constraint on when an action A must occur is expressed by an enabling condition on the Tick action that prevents an increase in time from violating the constraint; a lower-bound constraint on when A may occur is expressed by an enabling condition on A that prevents it from being executed earlier than it should be. I illustrate how one writes explicit-time specifications using the example of a simple version of a classic distributed algorithm of Radia Perlman [12]. The original algorithm constructs a spanning tree rooted at the lowest- numbered node, called the leader. The tree is maintained by having the 1 Dutertre and Sorea [3] use a different kind of timer variable that predicts the time at which an action will occur. 2 ∆ Tick = ∃ d ∈ {r ∈ Real : r > 0} : ∧ ∀ n ∈ Node : timer [n] + TODelay ≥ d ∧ ∀ ms ∈ BagToSet(msgs) : ms.rcvTimer ≥ d ∧ now = now + d ∧ timer = [n ∈ Node → timer [n] − d ] ∆ ∧ msgs = let Updated (ms) = [ms except !.rcvTimer = ms.rcvTimer − d ] in BagOfAll (Updated , msgs) ∧ unchanged ldr , dist Figure 1: The Tick action's definition for the leader algorithm. leader periodically propagate an I'm Leader message down it that informs each node of its distance to the leader. A new tree is constructed if a failure causes some node to time out before receiving the I'm Leader message. I have simplified it by eliminating failures, so correctness means simply that every node learns the leader within some fixed length of time. A complete TLA+ specification of the algorithm appears in [9]. Here, I describe only the TLA+ specification of the Tick action. The algorithm has three timing parameters, Period , MsgDelay, and TODelay. Each node n has a countdown timer timer [n]. Setting timer [n] to τ causes a timeout to occur between τ and τ +TODelay seconds later. By letting τ be the minimum timeout interval, this models both delay in react- ing to a timeout and variation in the running rate of physical timers. When its timeout occurs, node n sends an I'm Leader message and sets timer [n] to Period . If n receives an I'm Leader message from a lower-numbered node, it resets timer [n] to a suitable value. A message is assumed to be re- ceived at most MsgDelay seconds after it is sent, a constraint enforced with a rcvTimer countdown timer field in the message. The algorithm achieves stability if, upon receiving a message from its leader, a node n sets timer [n] to a value no smaller than Period + TODelay + dist[n] ∗ MsgDelay, where dist[n] is the distance from n to the leader. Figure 1 contains the definition of the Tick action from the TLA+ spec- ification. It can't be completely understood without seeing the rest of the specification and having some knowledge of TLA+ (including the definitions of the operators BagToSet and BagOfAll from the standard Bags module). However, it will indicate how timing constraints are specified and also give an idea of the high-level nature of TLA+. This version is for a continuous-time specification, in which now is incremented by some real value d . We obtain 3 a discrete-time specification by replacing " ∃d ∈ {r ∈ Real : r > 0} : " ∆ with " let d = 1 in ". The action's first two conjuncts enforce the upper-bound constraints. The first prevents timer [n] from becoming less than −TODelay, for each node n. The second prevents the timer ms.rcvTimer from becoming nega- tive, for all messages ms in the bag (multiset) msg of messages in transit. The action's remaining conjuncts assert how the variables are changed. The third conjunct asserts that now is incremented by d . The fourth and fifth conjuncts assert that all the timers are decremented by d , the fourth for each timer [n] and the fifth for the timer component ms.rcvTimer of each message ms. The final conjunct asserts that the specification's other variables are unchanged. The complete specification asserts the additional timing constraint that a timeout action of node n cannot occur before timer [n] has counted down past 0. This constraint is expressed by the conjunct timer [n] < 0 in that action's definition. 3 Model Checking Explicit-Time Specifications Most real-time system specifications are symmetric under time translation, meaning that system actions depend only on the passage of time, not on ab- solute time values. This section explains what symmetry and model check- ing under symmetry mean and describes a simple method of model checking explicit-time specifications that are symmetric under time translation. 3.1 Specifications and Temporal Properties Let a state of a specification be an assignment of values to all the specifi- cation's variables, and let its state space be the set of all such states. A state predicate is a predicate (Boolean function) on states, and an action is A a predicate on pairs of states. The formula s −→ t asserts that action A is true on the pair s, t of states. A behavior is a sequence of states. A temporal property is a predicate on behaviors. Temporal properties are represented syntactically as temporal formulas. Assume a specification S that consists of an initial predicate Init, a next- state action Next, and a liveness assumption L that is a temporal property, possibly equal to true. The initial predicate and next-state action form the safety part S of specification S. A behavior s 1 , s 2 , . . . satisfies S iff s 1 Next satisfies Init and s i −→ s i+1 for all i ; it satisfies S iff it satisfies both S and L. 4 3.2 Symmetry A symmetry is an equivalence relation on states. A state predicate P is symmetric with respect to a symmetry ∼ iff, for any states s and t with s ∼ t, predicate P is true in state s iff it is true in state t. An action A is symmetric with respect to ∼ iff, for any states s 1 , s 2 , and t 1 , A A s 1 −→ t 1 s 1 −→ t 1 implies there exists t 2 such that A s2 s 2 −→ t 2 In other words, for any states s 1 and s 2 with s 1 ∼ s 2 and any state t 1 , if A A s 1 −→ t 1 then there exists a state t 2 with t 1 ∼ t 2 such that s 2 −→ t 2 . A symmetry ∼ is extended to an equivalence relation on behaviors in the obvious way by letting two behaviors be equivalent iff they have the same length and their corresponding states are equivalent. A temporal property is symmetric (with respect to ∼) iff, for every pair of behaviors σ and τ with σ ∼ τ , the property is true of σ iff it is true of τ . A temporal formula is constructed from state predicates and actions by applying temporal operators, logical connectives, and ordinary (non- temporal) quantification. The formula is symmetric if each of its component state predicates and actions is symmetric. 3.3 Model Checking An explicit-state model checker works by computing the directed graph G of a specification S's reachable states. The nodes of G are states, and G is the smallest graph satisfying the following two conditions: (i) G contains Next all states satisfying Init, and (ii) if state s is a node of G and s −→ t, then G contains the node t and an edge from s to t. Paths through G (which may traverse the same node many times) starting from an initial state correspond to behaviors satisfying S. Those behaviors that also satisfy its liveness assumption are the ones that satisfy S. The model checker constructs G by the following algorithm, using a set U of unexamined reachable states. Initially, G and U are both empty. The checker first sequentially enumerates the states satisfying Init, adding each state not already in G to both G and U. It does the following, while U is nonempty. It chooses some state s in U and enumerates all states t satisfying Next s −→ t. For each such t: (i) if t is not in G then it adds t to G and to U; (ii) if there is no edge from s to t in G, then it adds one. 5 Model checking under a constraint P is performed by constructing a subgraph of G containing only states that satisfy the state predicate P . To compute the subgraph, this procedure is modified to add a state to G and U only if the state satisfies P . Model checking under a symmetry ∼ consists of constructing a smaller graph E by adding a state to E and U only if E does not already contain an equivalent state. The graph E constructed in this way satisfies the following properties: (i) s ∼ t for every distinct pair of nodes s, t of E; (ii) for every state s satisfying Init, there is a node t in E such that t satisfies Init and Next s ∼ t; (iii) for every node s of E and every state t such that s −→ t, the graph E contains a node t with t ∼ t and an edge from s to t . The specification is then checked as if E were the reachable-state graph. Here, I ignore practical concerns and assume a theoretical model checker that can perform this algorithm even if the state graph is infinite. All the results apply a fortiori if the state graph is finite. For model checking with symmetry to be equivalent to ordinary model checking, the following condition must hold: SS. A behavior satisfies S iff it is equivalent (under ∼) to a behavior de- scribed by a path through E starting from an initial state. This condition does not imply that the behaviors described by paths through E satisfy S, just that they are equivalent to ones that satisfy S. Condition SS is true if the specification satisfies the following two properties: S1. (a) Init is symmetric, or (b) No two states satisfying Init are equivalent. S2. Next is symmetric. The specification is defined to be safety symmetric iff it satisfies S1 and S2. An explicit-state model checker checks that a correctness property F holds by checking that L ⇒ F holds for every behavior described by a path through the reachable-state graph starting from an initial state, where L is the specification's liveness assumption. A symmetric property is true of a behavior iff it is true of any equivalent behavior. Condition SS therefore implies that model checking with symmetry is equivalent to ordinary model checking for verifying that a safety symmetric specification with a symmetric liveness assumption satisfies a symmetric property. The simplest kind of temporal property is a state predicate P , which as a temporal formula asserts that P is true initially. If the specification satisfies S1(b), then model checking with symmetry is equivalent to ordinary model checking for verifying that P is satisfied, even if P is not symmetric. 6 3.4 View Symmetry A view symmetry is defined by an arbitrary function on states called a view. Two states are equivalent under a view V iff the value of V is the same in the two states. Many explicit-state model checkers test if a state s is in the state graph G constructed so far by keeping the set of fingerprints of nodes in G and testing if G contains a node with the same fingerprint as s. Such a checker is easily modified to implement checking under view symmetry by keeping fingerprints of the views of states rather than of the states themselves. TLC supports view symmetry as well as symmetry under permutations of a constant set. View symmetry is equivalent to abstraction [2, 4] for a symmetric speci- fication S. Abstraction consists of checking S by model checking a different specification A called an abstraction of S. The view corresponds to the abstraction mapping from states of S to states of A. 3.5 Symmetry Under Time Translation Time-translation symmetry is a special kind of symmetry in which two states are equivalent iff they are the same except for absolute time. I now define what this means, using the notation that s.v is the value of variable v in state s. A time translation is a family of mappings T d on the state space of the specification S that satisfies the following properties, for all states s and all real numbers d and e: (i) T d (s).now = s.now + d , (ii) T 0 (s) = s, and (iii) T d+e (s) = T d (T e (s)). Specification S is defined to be invariant under this time translation iff it satisfies the following two conditions, for all real numbers d . T1. (a) A state s satisfies Init iff T d (s) does, or (b) s.now = t.now for any states s and t satisfying Init. Next Next T2. s −→ t iff T d (s) −→ T d (t), for any states s and t. Given a time translation, we define the time-translation symmetry ∼ by s ∼ t iff s = T d (t) for some d . T1 and T2 imply S1 and S2 for this symmetry. Hence, a specification that is invariant under a time translation is symmetric under the corresponding time-translation symmetry. Invariance under time translation is stronger than time-translation symmetry because, in addition to implying SS, it implies the following property. 7 TT. Let s 1 , . . . , s k and t 1 , t 2 , . . . be two behaviors satisfying S (the second behavior may be finite or infinite). If s k = T d (t j ), then the behavior s 1 , . . . , s k , T d (t j +1 ), T d (t j +2 ), . . . also satisfies S. To define a time translation, we must define T d (s).v for every real number d , state s, and variable v . Explicit-time specifications have three kinds of variables: now , timer variables, and "ordinary" variables that are left unchanged by the Tick action. We know that T d (s).now equals s.now + d . Time translation should not change the value of an ordinary variable v , so we should have T d (s).v = s.v for such a variable. For a timer variable t, we should define T d (s).t so that the number of seconds in which t will time out is the same in s and T d (s). The value of a countdown or count-up timer directly indicates the number of seconds until it times out, so T d (s).ct should equal s.ct for such a timer ct. Whether or not an expiration timer et has timed out depends on the value of et − now . The time translation T d preserves the number of seconds until et times out iff T d (s).et − T d (s).now equals s.et − s.now , which is true iff T d (s).et = s.et + d . With this definition of the T d , any explicit-time specification is invariant under time translation, and hence safety symmetric under time-translation symmetry, if it expresses real-time requirements only through timer vari- ables. Let v 1 , . . . , v m be the specification's ordinary variables and count- down and count-up timer variables, and let et 1 , . . . , et n be its expiration timer variables. Then symmetry under time translation is the same as view symmetry with the view v 1 , . . . , v m , et 1 − now , . . . , et n − now . 3.6 Periodicity and Zeno Behaviors Let NZ be the temporal property asserting that time increases without bound. A specification S is nonZeno iff every finite behavior satisfying S can be extended to an infinite one satisfying S and NZ [1]. Property NZ is not symmetric under time translation; by replacing states of a behavior with ones translated back to the behavior's starting time, we can construct an equivalent behavior in which now never changes. Thus, model checking with time-translation symmetry cannot be used to check that a specification is nonZeno. However, we can take advantage of time-translation invariance as follows to use ordinary model checking to show that a specification is nonZeno. Let S be a specification that is invariant under time translation. For sim- plicity, we assume that the initial condition of S asserts that now equals 0, so s.now ≥ 0 for all reachable states s. For any reachable state s, let 8 LeastTime(s) be the greatest lower bound of the values t.now for all states t equivalent to s (under time-translation symmetry). The period of S is de- fined to be the least upper bound of the values LeastTime(s) for all reachable states s of S. Intuitively, if a system's specification has a finite period λ, then all its possible behaviors are revealed within λ seconds. More pre- cisely, any λ-second segment of a system behavior is the time translation of a segment from the first λ seconds of some (possibly different) behavior. Define the condition NZλ as follows, where λ is a positive real number. NZλ . Every finite behavior satisfying S that ends in a state s with s.now ≤ λ can be extended to a behavior satisfying S that ends in a state t with t.now ≥ λ + 1. It can be shown that if a specification S is time-translation invariant, has a period less than or equal to the real number λ, and satisfies NZλ , then it is nonZeno. Therefore, we can check that S is nonZeno by verifying that S has a period of at most λ and that it satisfies NZλ . Here is how we can use model checking under time-translation symmetry to find an upper bound on the period of S. Let E be the state graph con- structed by model checking under this symmetry. Because every reachable state is equivalent to a node in E, the period of S is less than or equal to the least upper bound of the values s.now for all nodes s of E. (Since all initial states have now = 0, the period of most specifications will equal this least upper bound for a model checker that, like TLC, uses a breadth-first construction of the state graph.) Debugging features allow the TLC user to insert in the specification expressions that always equal true, but whose evaluation causes TLC to perform certain operations. Using these features, it is easy to have TLC examine each state s that it finds and print the value of s.now iff s.now > t.now for every state t it has already found.2 This makes computing an upper bound on the period of S easy. An explicit-state model checker that lacks the ability to compute the upper bound can ver- ify that λ is an upper bound on the period by verifying the invariance of now ≤ λ, using time-translation symmetry. To check that S satisfies NZλ , we must show that from every reachable state with now ≤ λ, it is possible to reach a state with now ≥ λ + 1. We can do this by model checking with the constraint now ≤ λ + 1, in which the model checker ignores any state it finds with now > λ+1. It is easy to verify NZλ under this constraint with a model checker that can check possibility properties. With one like TLC that checks only linear-time temporal proper- ties, we must show that S together with fairness assumptions on subactions 2 One of the features needed was added to TLC after publication of [8]. 9 of its next-state action imply that the value of now must eventually reach λ + 1 [1, 7]. That is, we add fairness assumptions on certain actions and check that eventually now ≥ λ + 1 holds, using the constraint now ≤ λ + 1. All of this, including the definition of period, has been under the assump- tion that now = 0 for all initial states. Extending the definition of period to the general case is not hard, but there is no need to do it. Invariance under time translation requires that either (a) the set of initial states is invariant under time translation, or (b) the value of now is the same in all initial states. In case (b), that value will probably either be 0 or else a parameter of the specification that we can set equal to 0. In case (a), we conjoin the requirement now = 0 to the initial predicate. Invariance under time trans- lation implies that, in either case, modifying the specification in this way does not affect whether or not it is nonZeno. 4 Comparison with Uppaal 4.1 The Leader Algorithm I have checked the TLA+ specification of the leader algorithm with the TLC model checker. Although the specification is time-translation invariant, the correctness property is not. It asserts (now > c(n)) ⇒ P (n) for each node n, where c(n) is a constant expression and P (n) does not contain now . We could add a timer variable and restate the property in terms of it. (This is what is done in the Uppaal model.) However, I instead had TLC check the property under a symmetry ∼ defined as follows. Let Σ be the maximum of c(n) for all nodes n. Then s ∼ t iff s.now and t.now are both equal or both greater than Σ. Both the specification and the correctness property are symmetric under ∼. This symmetry is view symmetry under the view consisting of the tuple v 1 , . . . , v k , if now > Σ then Σ + 1 else now , where the v i are all the variables except now . Real-time model checkers use much lower-level modeling languages than TLA+. Uppaal [10] is the only one I know of whose language is expressive enough to model this algorithm. Arne Skou, with the assistance of Gerd Behrmann and Kim Larsen, translated the TLA+ specification to an Uppaal model. Since Uppaal's modeling language is not as expressive as TLA+, this required some encoding. In particular, Uppaal cannot represent the potentially unbounded multiset of messages in the TLA+ specification, so the Uppaal model uses a fixed-length array instead. To ensure that the model faithfully represents the algorithm, Uppaal checks that this array does not overflow. 10 TLC and Uppaal were run on different but roughly comparable machines. As indicated, some Uppaal executions were run on a 30-machine network. More detailed results are presented in [9]. The parameters of the specification are the number N of nodes, a con- stant operator that describes the graph, and the timing constants Period , TODelay, and MsgDelay. The latter two are upper-bound constraints, which implies that the number of reachable states is an increasing func- tion of their values. Figure 2 shows the results of checking the correctness property on two different graphs, with 3 and 4 nodes, for some haphazardly chosen values of the timing bounds. Uppaal timings are given for a single machine and for the 30-machine network; fail means that Uppaal ran out of memory. We expect that increasing a timing bound will increase the number of reachable states, and hence TLC's execution time, since it increases the number of possible values of the timer variables. The time required by Uppaal's algorithm depends only on the ratios of the timing bounds, not on their absolute value. The results show that Uppaal's execution time is strongly dependent on the ratio MsgDelay/Period . For ratios significantly less than .6, Uppaal's execution time depends almost entirely on the graph and not on the other parameters. TLC's execution time depends on the magnitude of the parameters as well as on this ratio. Hence, if Uppaal succeeds, it is usually faster than TLC for small values of the parameters and much faster for larger values. Using 30 processors extends the range of parameters for which Uppaal succeeds. TLC can be run on multiple com- puters using Java's RMI mechanism. Tests have shown that execution speed typically increases by a factor of about .7 times the number of computers. This suggests that, run on a network of processors, TLC's execution speed is comparable to Uppaal's for the range of instances tested. However, since increasing the timing-constraint parameters increases the number of reach- able states, TLC will be slower than Uppaal for large enough values of these parameters. The overall result is that Uppaal can check models with larger timing- constraint parameters, and hence with a finer-grained choice of ratios be- tween the parameters. However, TLC can check a wider range of ratios among the parameters. For finding bugs, the ability to check parameter ra- tios of both 1:2 and 2:1 is likely to be more useful than the ability to check ratios of both 1:2 and 11:20.3 3 The Uppaal model was subsequently rewritten to improve its performance. Because the TLA+ specification was written to be as simple as possible, with no consideration of model-checking efficiency, the fairest comparison seems to be with the first Uppaal model. 11 N =3 N =4 ¨2 1r 1 2 3 4 3 MsgDelay 30-proc N Period MsgDelay TODelay Period TLC Uppaal Uppaal 3 10 3 5 .3 255 9.4 2.9 3 1 1 .33 4 9.4 13.4 5 2 5 .5 70 11.2 2.9 5 3 1 .6 13 30.8 3.0 5 3 5 .6 265 fail 20.9 3 2 1 .67 7 10.2 3.0 3 2 2 .67 20 fail 16.6 5 4 1 .8 27 32.5 9.2 5 4 5 .8 980 fail fail 2 2 1 1 11 fail fail 1 2 1 2 270 fail fail 1 2 2 2 1280 fail fail 4 10 3 5 .3 1385 42.2 2.5 3 1 1 .33 6 43.9 2.7 5 2 2 .4 42 48.3 4.2 5 2 5 .4 390 93.0 4.3 2 1 1 .5 6 48.2 3.7 5 3 1 .6 28 72.8 3.8 5 3 5 .6 1770 fail 84.6 3 2 1 .67 12 73.1 9.8 3 2 2 .67 44 fail 73.1 5 4 5 .8 6760 fail fail 2 2 1 1 13 fail fail 1 2 1 2 390 fail fail 1 2 2 2 1650 fail fail Figure 2: Comparison of Uppaal and TLC execution times in seconds for the indicated graphs with 3 and 4 nodes. 12 MsgDelay reachable msgs in transit Period MsgDelay TODelay Period states max mean 2 2 1 1 6579 6 3.46 1 2 1 2 240931 12 6.57 3 2 2 .67 20572 6 3.69 10 3 5 .33 247580 6 3.85 Figure 3: The number of messages in transit. The dependence on the MsgDelay/Period ratio can be explained as fol- lows. Since Period is a lower bound on the time between the sending of messages and MsgDelay is an upper bound on how long it takes to deliver the message, the maximum number of messages that can be in transit at any time should be roughly proportional to this ratio. The table of Figure 3 gives some idea of what's going on, where the results are for the 3-node graph. The first two rows show the dramatic effect of changing Period and leaving the other parameters the same. The second two rows show that the MsgDelay/Period ratio is just one of the factors determining the number of messages in transit and the number of reachable states. It is possible that these results reflect some special property of this ex- ample. However, the sensitivity to the MsgDelay/Period ratio suggests that it is the messages in transit that pose a problem for Uppaal. Each message carries a timer, and the performance of real-time model checkers tends to depend on the number of concurrently running timers. Perhaps the most common use of real time in systems is for constraints on message transit time—constraints that are modeled by attaching timers to messages. This suggests that Uppaal might have difficulty checking such systems if there can be many messages in transit. However, more examples must be tried before we can draw any such conclusion. TLC was also used to check that some of the instances in Figure 2 were nonZeno. For N = 3, this took about twice as long as checking the correctness property; for N = 4 the two times were about the same. Uppaal can check the new model on a single computer an average of 4.5 times faster for the N = 3 instances of Figure 2 and 50 times faster for the N = 4 instances, but it still fails when MsgDelay/Period is greater than about 1. The new model therefore does not alter the basic result that Uppaal is faster than TLC for the range of parameter ratios it can handle, but it cannot handle as wide a range. 13 Safety Liveness K states TLCs TLC Spin SMV TLC Spin SMV 2 155976 9 29 .7 1.3 128 3.7 2.5 3 450407 10 78 2.4 3.8 385 13 6.3 4 1101072 16 194 6.9 6.5 1040 49 10 5 2388291 26 399 19 10 3456 171 16 6 4731824 47 784 51 14 5566 468 22 7 8730831 78 1468 142 25 13654 1317 40 8 15208872 132 2546 378 35 3593 54 9 25263947 244 4404 977 46 5237 73 10 40323576 446 7258 2145 62 95 Uppaal 135 Figure 4: Execution times in seconds for a simple version of Fischer's al- gorithm with 6 threads, where TLCs is TLC with symmetry under thread permutations. 4.2 Fischer's Algorithm I also compared the explicit-state approach to the use of Uppaal on a ver- sion of Fischer's mutual exclusion algorithm [13] that is distributed with Uppaal. Because TLA+ is a very high-level language, TLC must "execute" a specification interpretively. It is therefore significantly slower than con- ventional model checkers for verifying simple systems. I also obtained data for two other popular model checkers whose models are written in lower- level languages: the explicit-state model checker Spin [6] and the symbolic checker SMV [11] that uses binary decision diagrams. The Spin model was written and checked by Gerard Holzmann, and the SMV model was written and checked by Ken McMillan. Checked were the safety properties of mu- tual exclusion and deadlock freedom (except for SMV) and a simple liveness property. This version of Fischer's algorithm uses a parameter K that is both an upper- and lower-bound timing constraint. All the models were tested for 6 threads, which is the smallest number for which Uppaal takes a significant amount of time. The results for different values of K are shown in Figure 4. Uppaal's execution time is independent of K . For checking safety, TLC was run both with and without symmetry under permutations of threads. (The liveness property is not symmetric.) The speedups obtained by the 6-fold symmetry should not be taken very seriously; in real examples one at best obtains only 2- or 3-fold symmetry. 14 Since Uppaal's execution time is independent of K , we know that for large enough values of K it will be faster than a model checker whose run- ning time depends on K . All of the model checkers could check the spec- ification for large enough values of K to provide reasonable confidence of its correctness, though the numbers do not bode well for the ability of TLC and Spin to check liveness for more complicated examples. We do not ex- pect TLC's performance on liveness checking to be good enough for large applications. But because Fischer's algorithm is so simple, it is dangerous to infer from these numbers that the performance of Uppaal and SMV would be good enough. 5 Conclusion Experts in the field will not be surprised that one can write and check explicit-time specifications using ordinary model checkers. But this is ap- parently not widely appreciated because it has not been stated clearly in the literature. Moreover, the use of view symmetry and the method described here for checking that a specification is nonZeno may be new even to experts. I know of no previous comparisons of the explicit-state approach with the use of a real-time model checker. The results reported here do not tell us how the two methods will compare on other examples. But they do indicate that verifying explicit-time specifications with an ordinary model checker is not very much worse than using a real-time model checker. Indeed, the results for the leader algorithm suggest that the explicit-time approach is competitive with Uppaal for distributed algorithms. The results of using TLC to check two more complicated versions of Fischer's algorithm are reported in [9]. They too suggest that TLC can be used in practice to check explicit-time specifications. The main advantage of an explicit-time approach is the ability to use lan- guages and tools not specially designed for real-time model checking. There are practical reasons for using a higher-level language like TLA+ instead of one designed expressly for model checking. As one industrial user remarked, "The prototyping and debug phase through TLA+/TLC is so much more efficient than in a lower-level language." 15 References ın [1] Mart´ Abadi and Leslie Lamport. An old-fashioned recipe for real time. ACM Transactions on Programming Languages and Systems, 16(5):1543–1571, September 1994. [2] Edmund M. Clarke, Orna Grumberg, and David E. Long. Model check- ing and abstraction. ACM Transactions on Programming Languages and Systems, 16(5):1512–1542, September 1994. [3] Bruno Dutertre and Maria Sorea. Modeling and verification of a fault- tolerant real-time startup protocol using calendar automata. In Formal Techniques, Modelling and Analysis of Timed and Fault-Tolerant Sys- tems, Joint International Conferences on Formal Modelling and Analy- sis of Timed Systems, FORMATS 2004 and Formal Techniques in Real- Time and Fault-Tolerant Systems, FTRTFT 2004, Grenoble, France, September 22-24, 2004, Proceedings, volume 3253 of Lecture Notes in Computer Science, pages 199–214. Springer, 2004. [4] Susanne Graf and Claire Loiseaux. Property preserving abstractions under parallel composition. In Marie-Claude Gaudel and Jean-Pierre Jouannaud, editors, TAPSOFT'93: Theory and Practice of Software Development, volume 668 of Lecture Notes in Computer Science, pages 644–657. Springer, 1993. [5] Thomas A. Henzinger and Orna Kupferman. From quantity to quality. In Oded Maler, editor, Proceedings of the International Workshop on Hybrid and Real-Time Systems (HART '97), volume 1997 of Lecture Notes in Computer Science, pages 48–62. Springer-Verlag, 1997. [6] Gerard J. Holzmann. The Spin Model Checker. Addison-Wesley, Boston, 2004. [7] Leslie Lamport. Proving possibility properties. Theoretical Computer Science, 206(1–2):341–352, October 1998. [8] Leslie Lamport. Specifying Systems. Addison-Wesley, Boston, 2003. A link to an electronic copy can be found at http://lamport.org. [9] Leslie Lamport. Real time is really simple. Technical Report MSR-TR- 2005-30, Microsoft Research, March 2005. 16 [10] Kim Guldstrand Larsen, Paul Pettersson, and Wang Yi. UPPAAL in a nutshell. International Journal of Software Tools for Technology Transfer, 1(1/2):134–152, December 1997. [11] K. L. McMillan. Symbolic Model Checking. Kluwer, 1993. [12] Radia Perlman. An algorithm for distributed computation of a span- ningtree in an extended LAN. In Proceedings of the Ninth Symposium on Data Communications, pages 44–53. SIGCOMM, ACM Press, 1985. [13] Fred B. Schneider, Bard Bloom, and Keith Marzullo. Putting time into proof outlines. In J. W. de Bakker, C. Huizing, W.-P. de Roever, and G. Rozenberg, editors, Real-Time: Theory in Practice, volume 600 of Lecture Notes in Computer Science, pages 618–639, Berlin, Heidelberg, New York, 1992. Springer-Verlag. 17
Pages to are hidden for
"charme2005"Please download to view full document

subset: commoncrawl


text:
Find commute times to your work, school, gym and more.
Emerald Place
Contact this property
Floor Plans
Prices, specials, features and availability subject to change.
Controlled Access, Swimming Pool
Property Details
Large, spacious floorplans...Convenient Location...Close to Sac State, ARC Colleges...
Emerald Place Apartments are professionally managed, offering an open serene atmosphere for relaxed living. Comfortable surroundings and a peaceful residential atmosphere are yours to enjoy. Conveniently located close to schools, restaurants and shopping. The apartment you want at the price you can afford!
Leasing Terms
Lease Terms: 6, 9, and 12 month terms offered. Security Deposits reduced to just $300.
Pet Policy: 25lbs max pet weight, 2 max pets per unit, refundable deposit varies upon move-out. Current prices OAC.
Pet Policy
Cats Allowed. Dogs Allowed. Call for service animal policy.
Community Information
204 units
Building Type
Contact Property
Check Availability or call
Similar Properties
Similar Properties
Neighborhood Info
My Commute
Edit Address
Save your commutes
Scores are compared to the national average. The higher the bar, the higher the crime risk. The lower the bar, the lower the crime risk.
near Sacramento, 95825Transit
Ratings & Reviews
5 stars
4 stars
3 stars
2 stars
1 stars
Emerald Place Resident
I have a comfortable apartment.
November 6, 2014
Emerald Place Resident
The place that we were shown wasn't what we were moving into, it was the model. The front of our apartment doo Read Morer looked as if it had been previously broken in to. Our neighbors smoked downstairs which came into our vents, when we requested a unit non smoking, my mother has mild asthma, she basically had to rely on her inhaler A LOT more. When we left when our lease was up they accused us of damages that we never created, they were already there.
July 15, 2014
Tamara R.
The facility was poorly managed and the property was horribly maintained. What repairs that I had seen done we Read Morere done by a courteous and nice staff, but I got the impression that they were not allowed to spend money on repairs and had to use remnants and left overs from other jobs. Absolutely atrocious!
July 10, 2014
Emerald Place Resident
At first I was happy to find, what seemed, a decent place at reasonable rent. The office staff was friendly an Read Mored accommodating and maintenance was also friendly and helpful. Over the 2 years I lived at Emerald Place the property slowly went downhill. The grass wasn't green anymore, maintenance was turning lazy in repair work requested, the sidewalks, stairs, and parking lot becoming stained and unsightly, and the "quality" of tenants moving in were lacking and had no control over their kids.
January 14, 2014
Contact Property
Contact this property
Office Hours
Monday-Friday: 9:00am-6:00pm
Saturday-Sunday: 10:00am-5:00pm
Managed by
JLE Properties


text:
A spinal osteosynthesis system, in particular for anterior spinal fixation, comprising at least an anchoring member, an elongated linking element and a connector capable of being fixed to the anchoring member and to the linking element by selecting an angular position of the linking element relative to the connector. The system comprises a second elongated linking element, the connector being capable of being fixed simultaneously to both linking elements.
1. A clamp for clamping a pair of rods to the vertebra of the spinal column comprising:
first and second branches each having respective first ends, second ends, and lengths between said first and second ends, said first and second branches being integral at their respective said first ends and free at their respective said second ends, said first and second branches having facing surfaces with rod receiving portions formed thereon for receiving each of said pair of rods therebetween;
at least one bone screw to facilitate attaching said first branch to the vertebra; and
a clamping element engaging at least said second branch for clamping said pair of rods within said rod receiving portions of said first and second branches,
wherein at least one of said pair of rods is substantially transverse to at least one of said respective lengths of said first and second branches.
2. The clamp of
3. The clamp of
4. The clamp of
5. The clamp of
6. The clamp of
7. The clamp of
8. The clamp of
9. The clamp of
10. The clamp of
11. The clamp of
12. The clamp of
13. The clamp of
14. The clamp of
15. The clamp of
16. The clamp of
17. The clamp of
18. The clamp of
19. The clamp of
20. The clamp of
21. The clamp of
This application is a National Stage of International Application No. PCT/FR99/02402, filed Oct. 7, 1999, the disclosure of which is hereby incorporated by reference herein.
The invention relates to spinal osteosynthesis systems, in particular for anterior fixation.
Spinal osteosynthesis systems for anterior fixation are known in which the connection element is formed by a plate, and others are known in which the connection element is formed by a rod. Because of their bulk, the plate systems are difficult to use, if they can be used at all, via the endoscopic route. Moreover, their limited size (length) means that they can only be used for simple vertebrectomies involving a single vertebra, or perhaps two. It is impossible to treat scoliosis with this type of system. Finally, a plate is difficult to adapt to the morphology of the vertebrae in which it is anchored. Moreover, the rod systems generally comprise fairly voluminous connectors which cannot always be used via the endoscopic route.
A spinal osteosynthesis device is also known from U.S. Pat. No. 5,938,663, which device comprises a clamp or connector with two branches which can clamp a connection rod between them, the branches being able to be engaged on a vertebral pedicle screw. The rod is received between the branches with interposition of a slotted ring which has a spherical face and is engaged on the rod in order to control the angular orientation of the connector relative to the rod before clamping. However, this device, while well suited for posterior fixation of the spine on the vertebral pedicles, does not always provide sufficient rigidity with a view to anterior fixation of the spine.
An object of the invention is to make available a different spinal osteosynthesis system, adapted for anterior fixation, easy to fit, ensuring good rigidity of the system on the spine, and compatible with being fitted via the endoscopic route.
To achieve this object, the invention provides a spinal osteosynthesis system, in particular for anterior fixation, comprising at least one anchoring member, an elongate connection element and a connector which is able to be fixed to the anchoring member and to the connection element by choosing an angular position of the connection element relative to the connector, the system comprising a second elongate connection element, the connector being able to be fixed simultaneously to the two connection elements.
Thus, the presence of the two connection elements gives the system very great rigidity, without complicating its assembly, without increasing the volume of its various components (which renders it compatible with fitting via the endoscopic route), and while maintaining the possibility of controlling the angular position of the connector relative to the first connection element. The system according to the invention does not require identical bending on the two connection elements. Moreover, the number of connectors can remain small.
Advantageously, the system is designed in such a way that the second connection element can be fixed to the connector only in a single angular position relative to the connector.
Thus, the shape of the second connection element dictates the relative angular position of the connectors which are fixed to it. This angular position can therefore be chosen in advance depending on the prior curvature given to this connection element, either at the time of manufacture or, better still, during the surgical intervention.
Advantageously, the second connection element has less resistance to bending than the first connection element.
Thus, the first connection element has mainly a function ensuring support of the connectors, and the second connection element has mainly a function ensuring angular positioning of the connectors.
Advantageously, the connector comprises two branches which can clamp at least one of the two connection elements.
Advantageously, the branches have first ends via which they are connected to each other and second free ends.
This simplifies production of the connector. Moreover, by appropriately selecting the flexibility of the connector and/or the diameter of the second connection element, this element can be introduced laterally between the free ends of the branches, which makes fitting easier.
Advantageously, the connector is able to receive the second connection element by lateral insertion between the free ends.
Advantageously, the branches can simultaneously clamp the two connection elements.
Advantageously, the anchoring member comprises a vertebral screw, at least a first of the two branches being able to be engaged on the screw.
Advantageously, the vertebral screw has a threaded orifice, the system comprising a clamping screw which can constitute a screw-nut connection with this orifice and is able to bear on one of the branches in order to clamp the branches.
Advantageously, the system comprises a second vertebral screw, the first branch being able to be engaged simultaneously on the two vertebral screws.
Advantageously, the system is designed in such a way that the second connection element, when fixed to the connector, extends in a trajectory of the second vertebral screw for its disengagement from the connector.
Thus, the second screw is prevented from starting to come out at an inopportune moment.
Advantageously, the system comprises a ring which can be engaged on the first connection element and received between the branches in order to choose the angular position of the first connection element before the branches are clamped.
Other characteristics and advantages of the invention will become more apparent from the following description of a preferred embodiment given as a nonlimiting example. In the attached drawings:
FIG. 1 is a perspective view of the system according to the invention;
FIG. 2 is a partial and exploded perspective view of the system in FIG. 1;
FIGS. 3 and 4 are two perspective views, from above and below, respectively, showing one of the connectors of the system in FIG. 1;
FIG. 5 is a view, half in elevation and half in axial section, of a ring of the system in FIG. 1;
FIG. 6 is a view partly from above and partly in section of the connector in FIG. 3 receiving the first rod;
FIG. 7 is a partial perspective view showing the head of the main screw; and
FIG. 8 shows the system from FIG. 1 fixed on vertebrae.
Referring to FIGS. 1 and 2, the system according to the invention comprises, in the preferred embodiment, a first elongate connection rod 2, or main rod, of circular cross section, a second elongate connection rod 3 or secondary rod of circular cross section and several connector sub-assemblies 4 which can be fixed simultaneously to the latter. Each of these sub-assemblies, of which only two can be seen in FIG. 1 and of which only one can be seen in FIG. 2, comprises a connector 6, a first vertebral screw or main screw 8, a clamping screw 10, a second vertebral screw or secondary screw 12, and a ring 13.
In the present example, the connector 6 is made in one piece. The different parts of the system are made of biocompatible metal.
Referring to FIGS. 3 and 4, the preferred connector 6 includes two branches 16 extending opposite to and at a distance from each other, giving the connector a general U-shaped profile. The connector 6 includes a plane of symmetry S visible in FIGS. 3 and 6, perpendicular to the width of the branches 16 and parallel to their length. Thus, each branch has a first end and a second end. The branches are connected to each other via their first ends. The second ends are free ends Referring to FIG. 6, at the point of origin of the branches 16, that is to say of their first end, the connector has two cylindrical and coaxial inner faces 18, 20 with axis 22 perpendicular to the plane S and with different radii. The face 20 of greater radius is in two distinct parts and extends on either side of the face 18 of lesser radius, which is traversed by the plane S. At their junctions, the two faces 18, 20 form two circular edges 24 with axis 22.
Referring to FIGS. 5 and 6, the preferred ring 13 has a cylindrical inner face 26 and a spherical outer face 28 which are coaxial. The cylindrical inner face 26 has a radius about equal to that of the main rod 2 in such a way that the ring 13, slotted on one side along its axis, can be received as a sliding fit on the main rod 2. Moreover, the ring 13 can be lodged between the branches 16 opposite the cylindrical faces 18, 20. The spherical outer face 28 of the ring has a radius which is adapted such that in this position the edges 24 of the connector 6 are in linear contact with the spherical outer face 28 of the ring 13 and serve as bearings for it. In this position, before clamping of the branches 16, the angular position of the main rod 2 engaged in the ring 13 can be controlled in two mutually perpendicular planes over an amplitude of, for example, 15° on either side of a mean position of the rod in which the rod is perpendicular to the plane S.
Referring to FIGS. 3 and 4, the branches 16 have two respective smooth cylindrical openings which, in this case, are through-orifices 30 extending coaxially opposite each other. The preferred main screw 8 is a bicortical vertebral screw and has a threaded body for this purposes, in a manner known per se. Referring to FIG. 7, it has a head 32 having a smooth cylindrical outer face 34. At the junction between the head and the body, the screw includes an annular flange 36 having a plane lower face perpendicular to a longitudinal axis of the screw, and a frustoconical upper face 38 with the narrowest cross section of the frustum situated towards the head 32 of the screw. The head 32 has a threaded orifice 39 coaxial to the body of the screw and, formed in the threaded face of the orifice 39, a noncircular shape such as a hexagon socket. The clamping screw 10 includes a threaded body 42 which is able to form a screw-nut connection with this orifice 39, and a screw head 44 in which a hexagon socket is formed. The head 44 has a spherical and convex lower outer face 46 whose narrowest cross section is situated towards the point of the screw.
In the preferred embodiment, one of the branches 16 of the connector 6, which for the sake of clarity we will here call the lower branch, has an extension 50 extending in the direction away from the cylindrical faces 18, 20 of the connector. This is the branch intended to be adjacent to the vertebra. This extension has the second, free, end of the branch. The two branches 16 are able to be engaged simultaneously on the head 32 of the main screw 8 introduced in the orifices 30 starting from the lower branch against which the upper face 38 of the flange 36 comes into abutment. The clamping screw 10 is then introduced into the head 32 of the main screw 8 starting from the upper branch 16. The tightening of the clamping screw 10 in the head 32 of the main screw 8 causes the branches 16 to move towards each other and causes, inter alia, frictional blocking of the first rod 2 in the chosen position relative to the connector 6.
The orifice 30 of the lower branch 16 has a lower edge, remote from the upper branch and intended to be towards the vertebra, having a concave spherical recess 40 intended to come into contact with the upper face 38 of the flange 36 in order to effect, by friction, rotational blocking of the connector 6 relative to the axis of the main screw 8. The orifice 30 of the upper branch 16 has an upper edge, remote from the lower branch and intended to be remote from the vertebra, having a concave spherical recess 40 intended to come into contact with the convex and spherical lower face 46 of the head 44 of the clamping screw 10 and making it possible to fix the latter and the main screw 8 by controlling the angular orientation of the main screw 8 relative to the connector.
In the preferred embodiment, the extension 50 has an opening in the form of a through-orifice 52. The lower branch 16 is curved in the area of the extension 50 in a direction away from the upper branch 16 in such a way that the axes of its orifices 30 and 52 are not quite parallel. The secondary screw 12 is a vertebral screw, here a monocortical screw, having a threaded body and a head 56 with a spherical and convex lower face 58 whose narrowest cross section is situated towards the body. Its head has a hexagon socket. The orifice 52 of the extension has an upper edge 60, oriented towards the other branch 16 and intended to be remote from the vertebra, having a spherical and concave or preferably frustoconical form 60, intended to come into contact with the spherical and convex lower face 58 of the head 56 of the secondary screw 12, making it possible to control the angular orientation of this screw relative to the connector 6.
Certain characteristics of the connector 6 which have not been expanded on in detail here will be found in U.S. Pat. No. 5,938,663.
The lower branch 16 can be bent in order to accentuate or reduce its curvature for better adaptation to the shape of the anterior part of the vertebra for which it is intended. Once bent, this branch 16 is no longer stressed in flexion since it is fixed to the vertebra by two screws 8, 12 along its length. The two screws, namely the main screw 8 and the secondary screw 12, are self-tapping and include bone threads.
In an alternative embodiment, the main screw 8 does not have a hexagon socket in this threaded orifice 39, and instead the flange 36 has a hexagonal shape or has two parallel and diametrically opposite flats which can cooperate with a tightening wrench for rotating this screw 8 relative to the connector 6.
The two connection rods 2, 3 each have a profiled rectilinear shape, the profile here being circular. The secondary rod 3 has a cross section, transverse to its longitudinal axis, having a diameter smaller than that of the main rod 2. The main rod 2 will, for example, have a diameter of 6 mm. The diameter of the secondary rod 3 will, for example, be between 30% and 80% of the diameter of the main rod 2. This small diameter allows the surgeon to choose the curvature of the secondary rod 3 corresponding to that of the level of the spine which is being operated on. By contrast, since the rings 13 allow relative angular positioning of the two connectors 6, the main rod 2 does not have to be bent. It can thus have a substantial diameter in order to be very robust.
The branches 16 of the connector have respective cylindrical recesses, grooves or jaws 74 formed in the faces of the branches opposite each other. The recesses 74 extend opposite each other and have axes parallel to each other and perpendicular to the plane of symmetry S.
On the upper branch 16, the recess 74 extends at a free end of the branch such that the orifice 30 is interposed between the faces 18, 20, on the one hand, and the recess 74 on the other. On the lower branch 16, the recess 74 extends between the two orifices 30 and 52, at the origin of the extension 50. It is contiguous with the orifice 52 so that it engages on its edge 60, i.e. recess 74 intercepts the upper opening of orifice 52 as shown in FIG. 3.
The secondary rod 3 is intended to be received in the recess 74 of the lower branch 16 in a unique angular position relative to the connector, perpendicular to the plane of symmetry S. The connector and the rod are dimensioned in such a way that the secondary rod can be introduced laterally between the free ends of the branches (that is to say without engaging the rod via one end in the connector). Before clamping, the space between the branches is wider than the diameter of the rod which can thus be freely inserted. When the two branches 16 are then clamped in the direction of each other, the recess 74 of the upper branch comes into contact with the secondary rod 3 which is thus in surface contact with each of the two recesses, which effect frictional blocking of the secondary rod 3 relative to the connector 6, which are thereby rigidly fixed to each other.
The secondary rod 3 is placed in the recess 74 of the lower branch after the secondary screw 12 has been introduced into the orifice 52. The position of the recess 74 of the lower branch is such that the secondary rod 3 then extends in the trajectory of the head of the secondary screw 12 for its disengagement from the connector and its exit from the orifice 52. Consequently, once the secondary rod 3 has been fixed to the connector, the secondary screw 12 can no longer be separated from the connector.
The upper branch 16 of the connector has at its free end a notch 76 which engages recess 74 with which it is contiguous, and facilitates maneuvering of the clamping screw 12 by means of a tool despite the space occupied by the upper branch.
Such a device is fitted in the following manner, with reference to FIG. 8. After exposing the affected vertebra 70 and the two adjacent vertebrae 72, a vertebrectomy is performed while preserving, if possible, the respective plates of these vertebrae. For each subassembly 4, a pilot hole is made on the lateral side of the associated vertebra 72 at an equal distance from the upper and lower plates, and at the limit of the most posterior quarter of the vertebral body. The main screw 8 is then inserted into this pilot hole as far as the limit flange 36. The connector 6 is then positioned on the main screw 8, blocked in translation by the conical face 38 of the main screw 8 matching the recess 40 of the connector 6. The fit of the lower branch 16 of the connector 6 on the vertebra is then checked and can be adjusted by withdrawing the connector in order to bend, to a greater or lesser extent, the lower branch 16, which is its most anterior part.
The secondary screw 12 is then screwed relative to the main screw 8 into the second orifice 52 of the lower branch 16 until the spherical seat 60 of the extension, provided for this purpose, comes into contact with the spherical part 58 of the secondary screw 12. It is desirable to position the connector 6 as parallel as possible to the plates.
After the two adjacent vertebrae 72 have been thus equipped, the main rod 2 is positioned by engaging it via one end in the rings 13 of the connectors 6 and the angular position of each sub-assembly. 4 relative to this rod 2 is controlled. The secondary rod 3 is then introduced laterally, and generally in the medial-lateral direction, into the recesses 74 of the connectors 6 after it has first been bent manually to obtain the curvature required for the corresponding level of the spine. In the event of an error, this rod 3 can be removed laterally in order to correct its curvature and then put back in place. FIG. 1 shows the system before the clamping of the branches. Final clamping is effected by virtue of the clamping screw 10 which is inserted into the main screw 8 and thereby compresses the connector 6 in order to clamp its two branches 16 towards each other. During this clamping, the clamping force is directed first on the main rod 2 via the ring 13, until the recess 74 of the upper branch comes into contact, with the secondary rod 3. Thereafter, the clamping force is distributed on the two rods 2, 3. Thus, the reaction at the level of the pairing of main screw 8 and clamping screw 10 is substantially coaxial to these.
The characteristics relating to the association of the first screw 8 with the clamping screw 10 will be able to be implemented independently of the presence of the extension 50 and of the secondary screw 12.
Although less advantageous, the extended branch can be the one which is intended to be farthest from the vertebra.
Although the invention herein has been described with reference to particular embodiments, it is to be understood that these embodiments are merely illustrative of the principles and applications of the present invention. It is therefore to be understood that numerous modifications may be made to the illustrative embodiments and that other arrangements may be devised without departing from the spirit and scope of the present invention as defined by the appended claims.
Patent Citations
Referenced by
Legal Events
Original Image


text:
Yvonne Castino joined the Serian UK trustees in November 2007. She is now retired from a long teaching and teacher training career, her last position being Senior Lecturer at the University of Portsmouth where she was the Programme Manager of the Language Centre.
Her areas of specialisation are English as a Foreign Language and Cultural Awareness. She remains an external examiner for the Cambridge University suite of examinations in English as a Foreign Language.
She first met Gemma and her parents when she was Gemma's English teacher and followed her career with great interest.
She lives in Netley Abbey, Southampton, and she and her late husband Franco, who died in December 2013, shared a great interest in Serian UK. Together they visited Noonkodin School in February 2007 and met their sponsored student for the first time.
They made a second visit in 2009 and were very pleased to see how the hard work was continuing and how well the students were progressing. Yvonne hopes to fulfil the plans she and Franco had to re-visit Noonkodin before too long.
Yvonne raises funds for Serian UK through the sale of greetings cards which she designs and makes throughout the year.
An amateur in the art of web design, Yvonne set up the original Serian UK website.


text:
New found from yesterday (3.9.2008), beautiful, abandoned house on countryside.
beautiful shot,,,love the tones and composition,,,, well done
A perfectly beautiful photo….I agree with midzing…comp and tones are perfect.
This is beautiful!! Gorgeous tones!
Thank you. With this place it was love at first sight – I must go there with more time as soon as I gonna get my car fixed, once again. There are lot of hidden photos still, I think… ;)
Love the composition and lighting. Well done.
You've captured some mystery here – beautiful!
beautiful- lovely sepia tones!
Melancholy and beautiful!!


text:
Mountain Views and Peaceful Setting in Carefree
From $250 / night
Hide Property Details
Description from the owner
- 2 bedrooms
- 3 bathrooms
- Sleeps 6
- Villa
Bedrooms: 2
Bathrooms: 3
Sleeps: 6
Type: VillaMountain Views and Peaceful Setting in Carefree
Description from the owner
Two bedroom, first floor townhome located in gated, secure community. Enjoy the year-round heated pool with whirlpool, putt on the nearby putting green, or enjoy hikes to nearby Carefree and Cave Creek. Located near the foothills,the area is peaceful yet allows you the flexibilty of visiting nearby Sedona, Flagstaff, and the Grand Canyon. The unit inlcudes two separate bedrooms, one with a king bed and the other with two queen beds and sleeps 6 comfortably. Numerous golf courses are nearby, and tennis courts are available at the nearby Carefree Inn Resort.
About the owner
About the owner
Jeff Schuh
Jeff Schuh
Response rate 80%
Response time Within a day
Calendar updated 3 days ago
Overall rating5.0 / 5.0 based on 5 reviews
Mountain Views and Peaceful Setting in Carefree
Carefree, Arizona
Additional Location Information
Located in quiet area of Carefree yet close to shopping and the offerings of the greater Phoenix area. Black Mountain is a short distance away and is awesome for a moderate hike. Cave Creek is also just down the road and offers many restaurant and entertainment venues.
More About This Location
Mountain nearby
Also Consider
FlipKey has rentals in over 11,000 destinations around the world:
Palm Springs vacation rentals
574 rentals from $125/night
Flagstaff cabin rentals
69 rentals from $89/night
Fajardo vacation rentals
44 rentals from $116/night
Barbados vacation rentals
2,027 rentals from $80/night
Ocracoke vacation rentals
5 rentals from $131/night


text:
Keeping Up with Python: the 2.2 Release
Generators extend from the idea of iterators. The main motivation for generators, however, comes from a different angle: they allow saving state across function calls. Static variables, such as in C functions, have the ability to maintain their value across multiple calls of that function. This partially solves the state problem, but what would be really nice would be to yield a value just like an iterator but be able to freeze execution only to resume exactly where you left off when it is called again. This is exactly what generators do. They represent the idea of merging iteration with state along with functions that are resumable. When they do, they pick up right where they left off, keeping intact all the state information they need to deliver the next item. Note that we use the term yield here for two reasons: to hint that it's not a true return (along with frame object stack pop) and to introduce the new keyword yield.
For backward compatibility (in case there is code out there that uses yield as an identifier), you must include the "from __future__ import generators" directive to use generators. Generators will become standard soon (2.3?) so importing will not be a necessity. Generators behave in another manner similar to iterators: when a real return or end-of-function is reached and there are no more values to yield, a StopIteration exception is raised. Here's a simple example:
def simpleGen(): yield 1 yield '2 --> punch!'
Now that we have our function, let's call it and get a generator object:
>>> myG = simpleGen() >>> myG.next() 1 >>> myG.next() '2 --> punch!' >>> myG.next() Traceback (most recent call last): File "", line 1, in ? myG.next() StopIterationor more aptly: for eachItem in simpleGen(): print eachItem. Of course that was a silly example. I mean, why not use a real iterator for that? More motivation comes from being able to iterate through a sequence that requires the power of a function, rather than static objects already sitting in some sequence.
In the following example, we are going to create a random iterator that takes a sequence and returns a random item from that sequence:
from random import randint def randIter(seq): while len(seq) > 0: yield seq.pop(randint(0, len(seq)-1))
The difference is that each item returned is also consumed from that sequence, sort of like a combination of list.pop() and random.choice():
>>> for eachItem in randIter([123, 'xyz', 45.678, 9j]): ... print eachItem ... 'xyz' 9j 45.678 123Table 1 is a summary of the differences between iterators and generators. You can find more details on both iterators and generators in their respective PEPs (234 and 255).
This is perhaps the most controversial update to Python so far. There are many pros and cons, but finally those who believe in true division have won out. To highlight this change let's define (or redefine) some division terms and their functionality with integer and floating-point operands.
When presented with integer operands, classic division truncates the decimal place, returning an integer (see the "Floor Division" section below). When given a pair of floating-point operands, it returns the actual floating-point quotient (see the "True Division" section). Here is an example of what Python's division has been and still is today (actually a mix of true and floor division):
>>> 1 / 2 # perform integer result (floor) 0 >>> 1.0 / 2.0 # returns real quotient 0.5
This is the case where the result should always be the actual quotient, regardless of the type of the operands. This is the big change that is to come our way when Python 3.0 nears reality. For now, to take advantage of true division, one must give the from __future__ import division directive. Once that happens, the division operator ( / ) performs only true division:
>>> from __future__ import division >>> >>> 1 / 2 # returns real quotient 0.5 >>> 1.0 / 2.0 # returns real quotient 0.5
- What continuous engineering is
- How to continuously improve complex product designs
- How to anticipate and respond to markets and clients
- How to get the most out of your engineering resources
Sponsored by IBM
Learn to:
- Define service virtualization
- Select the most beneficial services to virtualize
- Improve your traditional approach to testing
- Deliver higher-quality software faster
Sponsored by IBM
Trending Topics
Free DevOps eBooks, Videos, and more!
Regardless of where you are in your DevOps process, Linux Journal can help!
We offer here the DEFINITIVE DevOps for Dummies, a mobile Application Development Primer, and advice & help from the expert sources like:
- Forrester
- Gartner
- Linux Journal
|Users, Permissions and Multitenant Sites
|Mar 26, 2015
|New Products
|Mar 24, 2015
|Flexible Access Control with Squid Proxy
|Mar 23, 2015
|Solving ODEs on Linux
|Mar 19, 2015
|DevOps: Everything You Need to Know
|Mar 19, 2015
|Tighten Up SSH
|Mar 18, 2015
- Be a Mechanic...with Android and Linux!
- New Products
- Users, Permissions and Multitenant Sites
- Flexible Access Control with Squid Proxy
- Security in Three Ds: Detect, Decide and Deny
- High-Availability Storage with HA-LVM
- Solving ODEs on Linux
- DevOps: Everything You Need to Know
- Tighten Up SSH
- Non-Linux FOSS: MenuMeters


text:
Trask is a populated place located in Howell County at latitude 36.978 and longitude -91.802.
The elevation is 1,207 feet. Trask appears on the Trask U.S. Geological Survey Map. Howell County is in the Central Time Zone (UTC -6 hours).
Howell County: Cities & Towns | ZIP Codes | Schools | Features (airports, churches, hospitals, etc.)
BIG Map |
Driving Directions |
Data Source - USGS GNIS FID: 727762
BOUNDARY MAPS: Township of Hutton Valley | Howell County
Trask, MO
(< 0.1 mile)
Hutton Valley, MO (4.7 miles WSW)
Turnerville, MO (5.5 miles NE)
Mountain View, MO (5.5 miles ENE)
Hattie, MO (7.2 miles NNE)
Peace Valley, MO
(8.1 miles SSE)
Harlow Ford, MO (8.4 miles NE)
Ratcliff Ford, MO (8.4 miles NE)
White Church, MO (8.9 miles S)
Cardinal Acres, MO (9.1 miles NE)
Guild, MO
(6.4 miles N)
Chapel, MO (7.5 miles ESE)
Varvol, MO (12.1 miles NW)
Nagle, MO (13.1 miles N)
Not, MO (13.5 miles ESE)


text:
Be part of the debate
Policy Research Shop
Policy Research Shop
Students at the Policy Research Shop provide nonpartisan research on critical issues to state legislators in Vermont and New Hampshire.
Tuck School of Business Professor Paul Argenti weighs in on the dynamics of failed marketing campaigns.
TDI professor and physician H. Gilbert Welch writes about why it is important for doctors to avoid "knee-jerk medicine."
Geisel Professor Jay Dunlap and colleagues take a closer look at glowing fungi.
Springfield, Mass., will be the site of a museum honoring Theodor Geisel, Class of 1925.
Happy first day of spring term! Have a great first day of classes!
.@NashuaTelegraph features Sumant T'16 who went on a Global Insight Expedition in Israel w/ Tuck prof Adam Kleinbaum. http://t.co/LDBCH6wVpb
Considering an #MBA? Join #TuckSchool AdCom Thursday, April 2, 12-1pm ET for open Q&A; http://t.co/0zE8EAd8OA. #bschool
Tuck sponsors sled hockey game w/ Ice Vets, a nonprofit offering sport opps for disabled vets. http://t.co/j2PxCuNWo2 http://t.co/HxSP8Zqh95
.@Dartmouth Intro to #Engineering Project: RunRite Smartsole Shoe Insert: https://t.co/k7wEu8oelV #running
Today @ 4: "Health Systems Design and Implementation: A User-Centered Perspective" http://t.co/7Tj9GvcxUd
Science Day is tomorrow! Grad students will be hosting activities for students in grades 6-9: http://t.co/wF22q8kIug http://t.co/NGduoB02pz
Dr. Melody Brown Burkins was appointed Associate Director for Programs at the Dickey Center: http://t.co/IfnWqHEbyl http://t.co/wVEXpi2Ro1
Nicholas Warren, PEMM, and Erin Shoemaker-Kiess, MCB, met with Senator Jeanne Shaheen: http://t.co/lEQ9qfQvmY http://t.co/IsbHWHJf8s
RSVP for Melina Gehring's Poster Session Workshop for International Grad Students by Wednesday, April 1st: http://t.co/NAvWhU2uZr
Will an apple a day really keep the doctor away? Researchers test an old proverb http://t.co/oKD2cOML9d
Find out more about President Phil Hanlon's Moving Dartmouth Forward Plan, which will create higher expectations of students while strengthening Dartmouth's longstanding commitment to leadership in teaching and learning.


text:
Brandy is gearing up to drop her new album Two Eleven on October 16, and today, she released her second single "Wildest Dreams." The song is produced by Tha Bizness and written by Sean Garrett, and Brandy definitely proves on the track that she's still got it going on. Click the link in her tweet to listen.
Here it is!!"Wildest Dreams"!!! fb.me/1q9y7csN1— Brandy Norwood (@4everBrandy) August 21, 2012
[via Rap Radar]
Brandy Tells All: The Stories Behind Her Classic Records


text:
Is there an STO Affiliate Program?
View Single Post
Archived Post
Lt. Commander
Join Date: Dec 2007
Posts: 120
# 3
I'm talking about selling subscriptions to the game, so say somebody goes to my website and has a look around and then decides to go and check out STO through an advertisment on my page, if they then sign up to a paid account can I make real world money from that, or C-Store points maybe? Though I'd prefer real money of course. I can buy things like food and shelter with it.
A promo code that I could use. That would work. I seem to remember having the option to enter a promo code when I signed up.


text:
Asset allocation? Think asset location
Annuities can be very valuable
Join the NASDAQ Community today and get free, instant access to portfolios, stock ratings, real-time alerts, and more!
"Buy one, get one free" deals aren't confined to bags of chips
at the supermarket or sneakers at the shoe store: Many
travel-rewards credit cards offer free or heavily discounted
companion tickets on plane flights.
Those perks can save travelers hundreds of dollars, but there
are typically requirements and restrictions that can make the
offers less appealing.
"It sounds good, but a lot of times it's a marketing angle,"
says Geoff Whitmore, founder of
, a site that focuses on travel rewards. "You think, 'Oh, a
companion ticket.' Then you find out it's not as flexible as you
thought when you applied for the card."
The best advice, Whitmore says, is to read the terms and
conditions about the companion tickets before you apply for a card,
especially if the ability to travel cheaply with a friend or
relative is a major reason for applying.
As far as credit card travel perks go, companion tickets have
traditionally been a minor selling point, an extra tacked on
alongside other features. But with airlines making frequent flier
more difficult to earn and spend
for many travelers, the appeal of a companion ticket could become
Not all airlines offer companion tickets through credit card
rewards programs, and those that do, run the programs in different
ways (See chart, "
airline cards' companion ticket policies
"). Some send yearly coupons. Others are targeted at big spenders
or are redeemable only in conjunction with tickets that cost a
certain amount. Many can be redeemed online, but with some, you
have to book travel the old-fashioned way: by calling the airline
or, in one case, a travel agent (remember them?).
Southwest Airlines' program -- one of the more-popular ones --
issues a companion pass good for an unlimited number of flights for
up to two years. To get it, you have to earn 110,000 frequent-flier
points in a calendar year, but that includes spending and sign-up
bonuses on Chase's Southwest Airlines Rapid Rewards credit
Child-sleep consultant Lori Strong, 38, of Austin, Texas,
qualified for the companion pass last year, through a combination
of flights and spending on a Rapid Rewards business card, including
a 50,000-point sign-up bonus. She designated her 4-year-old
daughter as her companion, and the two have taken trips to
California, New York and Maryland, usually to visit relatives with
her husband and 8-year-old son.
"It's really been a great thing to have for our family," Strong
says. She estimates it has saved more than $1,000. When she books a
ticket for herself, she clicks a button on the Southwest website
that lets her add a ticket for her daughter. She has to pay an
airport security fee of up to $10, but says she has had no problems
with availability, even over Christmas holidays.
Some programs 'convoluted'
Matt Hague, founder of finance and travel blog site
, says the Southwest companion deal is one of the best, especially
when Chase offers 50,000-point sign-up bonuses, as it does
periodically, including earlier this year. If you were to apply for
both the personal and business versions of that card, you would
very nearly reach the 110,000 points needed for the pass.
Hague says that approach can be a great deal if you happen to
live near an airport served by Southwest. Other companion ticket
programs can be "more convoluted," he says.
Some airlines offer not passes, but free or deeply discounted
companion tickets, issued usually once a year. For instance,
holders of cards affiliated with US Airways and Alaska Airlines
receive certificates enabling them to buy a second ticket for $99.
Holders of certain Delta Airlines-affiliated cards receive annual
certificates for a companion to fly for free in coach, in certain
fare classes. A card affiliated with Virgin America offers an
annual $150 off a second fare.
Of course, there are restrictions. On Barclaycard's US Airways
Premier World MasterCard, for example, the fare on the original
ticket must be $250 or more, it's available only for coach travel
in the U.S. and Canada, and for 2014, the certificate cannot be
used on 28 specified dates throughout the year. Plus there are
restrictions on flying to certain cities around the Super Bowl,
Mardi Gras and the Kentucky Derby.
Other cards make it even more difficult to use a free companion
ticket. For instance, with the Citi ThankYou Premier card -- one of
the few nonairline cards to offer the benefit -- cardholders
receive one free companion ticket a year for a round-trip coach
flight in the U.S., but it and the primary ticket must be booked by
calling a particular travel company to figure out fares and
availability. The travel company's prices can be higher than those
on airlines' websites, and the terms and conditions warn that the
travel company "may not be able to honor requests for specific
airports, airlines, flight times or routings."
International tickets, too
While some cards offer breaks on domestic coach travel, others
target big-spending international travelers and are probably out of
reach for most consumers. Take Chase's British Airways Visa
Signature card. If you make $30,000 in purchases on that card in a
year, you'll receive a "Travel Together Ticket," which enables a
second person to accompany you when you redeem BA Avios miles for a
free trip. It's good on BA flights in coach, business or first
class -- but you still have to pay large fuel surcharges estimated
between $682 and $1,250 per coach ticket.
Other international carriers such as Lufthansa offer a yearly
free companion ticket in coach to holders of the Miles & More
World MasterCard by Barclaycard. But downsides include having to
pay expensive fuel surcharges and being unable to use the ticket
unless you buy one of the more expensive coach fares.
Not every companion-ticket program is for everybody, but most
every program will appeal to somebody.
Comedian Dan Nainan, 32, of New York City, says he likes using
his free companion ticket on Delta Air Lines every year -- even if
it doesn't count toward building up frequent flier miles.
In January, he bought a $253 round-trip coach ticket from New
York to Florida to "get away from the freezing cold" for a long
weekend visiting his cousin. His girlfriend traveled on the free
companion ticket. One drawback to using the pass, Nainan says, is
that it makes him ineligible for first-class upgrades he might
receive as one of Delta's elite-level frequent fliers. Or, he
muses, maybe it's a blessing that he won't catch grief for enjoying
first-class perks while his girlfriend suffers in coach.
"Overall, I like it, and it's a great benefit," he says. "It
kind of tips the scale toward getting the card and keeping it."
Enter up to 25 symbols separated by commas or spaces in the text box below. These symbols will be available during your session for use on applicable pages.
You have selected to change your default setting for the Quote Search. This will now be your default target page;
unless you change your configuration again, or you delete your
cookies. Are you sure you want to change your settings?


text:
Was £28.00
Now £14.00
Now £8.40
70% Off
Save a total of £19.60
Black 'Gracie' long sleeve fuller bust pyjama top
Item No.1510106328
Gracie Secret Support Chemis. Gracie soft fabrics feature lace trims plus a cute button and bow detail. While the great support and new sizes mean the fit is as fine as the look. Comfort with a classic feel. Freya loungewear is available from an XS to XL and you also have the choice between a medium cup (B-DD) and a fuller cup (E-G) so you can buy loungewear that fits your bust as well as your body.
This item is available for
UK Home Delivery Free (UK orders over £30) or £3.99
Delivered within 4 working days
Next or Named Day Delivery
7 days a week between 8am - 6pm
Monday to Friday - order by 7pm - £4.99
Saturday & Sunday - order by 6pm - £4.99
Next or Named Evening Delivery
7 days a week between 6pm - 10pm
Monday to Sunday - order by 10pm - £6.99
Collect from store the next day - FREE
Order before 7pm for next day collection after 12pm midday
Show participating stores
Full Delivery details
- Secret support for added comfort
- 92% viscose, 8% elastane
- Machine washable


text:
'Smart Guns' Setting Off Debate: How Smart Will They Really Be?
Published: October 22, 1998
(Page 2 of 2)
However far it is from production or widespread acceptance, the new ''safe gun'' or ''personalized gun,'' as it is also known, is still the biggest news in years to come out of the firearms industry, which has been rooted here in the Connecticut River Valley since Samuel Colt began making revolvers more than 160 years ago.
In fact, the smart gun is the first technological countermeasure the gun industry has mounted to the wave of legal and marketplace challenges that have confronted the industry in recent decades -- from laws against assault-style rifles to recent efforts by anti-gun forces to use product liability lessons of the tobacco wars against gun makers.
But while Colt's officials are enthusiastic about the gun's prospects, critics inside and outside the industry see it as a waste of time.
Smith & Wesson, Colt's big rival a few miles to the north in Springfield, Mass., suspended its smart gun program some time ago to wait for the technology to advance. L. E. Shultz, president and chief executive at Smith & Wesson, echoes the main argument against smart guns, that technology will do nothing to stop the misuse of guns by their proper owners.
''If you want to shoot your wife, you will go out and get your gun and shoot her,'' Mr. Shultz said. ''It's the aura of the smart gun as the solution to that situation that I have a problem with, because I don't think we will be able to develop in our lifetime or even in our children's lifetime a gun that can determine its user's attitude.''
Sturm, Ruger, based in Southport, Conn., has also rejected the current leading technology for the smart gun: the radio transmitter, called a transponder, which, when brought close to the gun, engages its firing system.
''The major failure that we see with any of these so-called smart guns is the battery,'' said Stephen L. Sanetti, vice president and general counsel at Ruger, which is the only publicly held American firearms company. ''Whether you are on a target range or on a hunting trip you saved for for years, or, in the worst-case scenario, defending your family, if the battery has failed, your options are that either the gun doesn't work or the safety doesn't work, and that makes the technology incompatible with firearms.
''Because one of the major components of the safety of a firearm is its reliability under all conditions.''
A safe gun would also continue to work if its transmitter bracelet were stolen along with the weapon, as well as if, in a struggle with an assailant, the gun were fired while still in or touching the owner's hand. The Colt's smart gun's police transponder will be programmable so that officers can fire one another's guns.
Testing to Satisfy Police Officers
Colt's officials say that their smart-gun prototype still has to undergo real-life testing -- dropping it in the mud and hard surfaces, using it in the rain -- -- before it will be ready to go on the market. The testing will be intended to satisfy police officers, who are expected to be the first users of the guns; 13 percent of officers killed in the line of duty are shot by their own guns.
The transponder used by Colt's has a range of only a few inches, Colt's officials said, making a gun useless even if held within feet of an officer by an assailant. Use by the police would go a long way to reassuring civilian customers that the smart gun will fire when it should, Colt's officials believe.
But it is the unproven nature of the smart gun that causes the lobbying arm of the National Rifle Association, the country's largest gun owners' group, to resist any laws mandating their adoption, said James Jay Baker, the N.R.A.'s executive director.
''I don't think there is any harm in proceeding with the development of a safer firearm,'' he said, ''but I think there is a problem with adopting a technology that doesn't even exist yet. Some of the anti-gun proponents in Congress are already talking about mandating it for everybody, and it's not even on the market yet, and the police haven't even accepted the technology yet.''
Researchers Concerned About Battery Life
Sandia National Laboratories in Albuquerque, N.M., a Federal research center, issued a 172-page review of smart gun technology for use in police work in 1996 with a highly iffy conclusion. After reviewing more than a dozen possible technologies, from Colt's radio frequency tag system to biometric devices that cause the gun to recognize, say, the length of a particular police officer's finger, it found that no approach scored higher than a B, in part because of concerns about battery life.
''It may take a generation of smart gun systems to come and go before a smart gun is not only common, but favored over a non-smart gun,'' the report, by D. R. Weiss, concluded.
In a telephone interview, Mr. Weiss predicted that smart guns would eventually be perfected, but he cautioned that the technology should be settled before laws mandating them are enacted.
''I think we will see them eventually,'' he said, ''but we will have to make sure that the job is done properly, because it would be very detrimental if something were put out and widely distributed and it didn't work. The mistake could cause smart guns their reputation.''
Outside the gun industry, the smart gun has provoked conflicting reactions. Mr. Brady, the former White House press secretary, went to Trenton in September to endorse New Jersey's proposed smart-gun legislation, which is strongly opposed by the N.R.A. and by New Jersey's influential gun-owners lobby, the Coalition of New Jersey Sportsmen. (Although the the New Jersey legislation is supported by some members of both parties, it faces steep obstacles in getting to the floors of the House and Senate.)
But the more stringently anti-gun Violence Policy Center, which favors phasing out private handgun ownership entirely, opposes them. The center maintains that smart guns are unlikely to reduce suicides, the largest single cause of death from guns, and that the second-highest category, homicides, are most often committed among acquaintances.
Moreover, the center maintains, most gun owners have more than one handgun, leaving plenty of ''dumb'' guns still at large.
Finally, the Center objects to any technology that is likely to lead to increased handgun sales -- and additional sales are just what Mr. Hendrickson at Colt's expects to get from smart guns.
''Let's just say this thing hits next year and it works, and it works well,'' Mr. Hendrickson said. ''You're going to see so much legislation that says that only the smart gun can be sold that it's going to put the other gun factories out of business.''