P2PSP (https://p2psp.github.io) is an application-layer protocol that provides real-time broadcasting of media streams on the Internet. Peers collaborate to diseminate the stream that is generated by a single source, generating a controlled latency and protocol overhead. P2PSP overlays (teams) are push-based (topology-driven) dynamic meshes. The chunks of data are forwarded without explicit requests, using a set of dynamic routes, that except in the case of losing chunks, show a maximum bounded latency, deﬁned by the overlay users.
Cursive is used the ﬁrst time a P2PSP-related term/concept is introduced, and for key concepts or ideas.
P2PSP has a modular design organized in sets of rules, where each module is especialized in implementing diﬀerent functionalities.
P2PSP supposes that there is a collection of channels that are broadcasted in parallel.1 The channels are available at one or more2 streaming servers, and each channel has a diﬀerent URL (Universal Resource Locator), usually expressed as a Web address with the structure:
Notice that a server can be serving several channels.
P2PSP does not perform data-ﬂow control over the stream. The transmission bit-rate between P2PSP entities is controlled by the servers (Icecast servers, for example), which provides the stream to the P2PSP teams. Fig. 1 shows an example of a streaming overlay where several servers relay a set of channels generated by a set of source-clients, directly or through other servers. As can be seen, a listener (which usually plays the stream) can be replaced by a splitter, a P2PSP entity that sends the received stream (a single channel) to a set of P2PSP peers.
In a pure CDN system, users request the channels directly to the servers. Unfortunately, this simple procedure has a drawback: normally, users do not know the load nor the distance to the servers. This problem can be solved by using a load balancer. The listeners, which know the URL of the required channel, connects ﬁrst to a load balancer which redirects them (with an HTTP 302 code) to a suitable server.
This idea can be extended to minimize the response time of hybrid CDN/P2PSP structures. When a user (who knows an URL of the channel) runs a local peer, it provides to his peer the URL of the channel (the URL pointing to a server and a mount point). Then, the peer (as any other listener does) contacts a load balancer which in this case sends a list of splitters which are broadcasting the channel.3 Then, the peer tries to connect with all the splitters in parallel, and the ﬁrst establised connection determines the selected splitter (the rest of connections are closed). If only those splitters with space in their teams answer to the peer, this procedure should select the “nearest” splitter for the peer in terms of response time.
For the case of the incorporation of new splitters to the network, the procedure is similar. A new splitter (which is instantiated knowing an URL of a channel) contacts the load balancer which returns a list of servers and peers, which are serving the channel. Then, the splitter tries to connect with all of them in parallel, and the ﬁrst successfull connection is ﬁnally selected.4
Using the idea of the extended load balancer, when a player (listener) connects to it, if there is a local peer running in the same host or the same private network that the player, the balancer will redirect the player to the local peer.
Finally, it is compulsory that all the splitters associated to the same channel to generate exactly the same chunks (content and header). See Section 9 for more information.
|Maximum number of peers in a team|
|Buﬀer size in chunks in the peers|
|Size of the set of last peers served by the splitter|
|Maximum allowed number of lost chunks|
|Number of monitors|
|Number of rounds to compute .|
|Number of peers in the team|
|Physical network latency|
|Latency experimented by the end-user|
DBS provides ALM  of a media stream in unicast environments . First, the media is sent by a streaming server, and received by a splitter (see Sec. 1). Then, the splitter divides the stream into a sequence of chunks of data, and relay them to its team using a round-robing schema. A team is composed by peers and each peer gathers the chunks from the splitter and the rest of peers of the team, and sends them to at least one player5 .
A team is a set of one or more peers (referenced by their end-points) that share the same stream. By deﬁnition, in a team of size one (the corresponding splitter is considered out of the team if feeds), the only peer is known as a monitor peer, and in a team with more than one peer, at least one of them must be a monitor, which are instantiated by the overlay administrator to monitorize diﬀerent aspects of the broadcasting, such as, the expected quality of the rendered video at the peers or the expected average end-user latency.
The number of peers (normal peers and monitors) in a team has a maximum (see Tab. 1). This parameter has a impact on the latency of the protocol (see Sec. 2.4) and usually is deﬁned by the administrator of the overlay.
The splitter divides the stream into chunks of constant length (chunk size), and sends exclusively each chunk to a diﬀerent origin6 peer, using a round-robin schema. Chunks are enumerated to distinguish them, and this information is transmitted as a part of a chunk header.
We deﬁne a round as the process of transmitting diﬀerent chunks from the splitter to a team of peers. Therefore, for a team of size , the round time . Notice that is generally variable, and depends on the current number of peers in the team (), and the chunk time () (which depends on the chunk size () and the average bit-rate of the media stream).
In DBS, all the peers of the team are origin of a diﬀerent chunk, in each round.
After connecting with a splitter, incoming peers request (using a reliable communication) to the splitter the current set of peers in the team. To minimize the joining time, the peer sends a message to each other peer of the team, in parallel with the reception of the set. When a peer of the team receives a , it adds the sender of the message to a table7 of peers called (see in peer.py). If a peer has an entry , then each chunk received by and originated at will be forwarded to . When an incoming peer has received the set of peers, its forwarding table has been initialized to . Notice that, as long as the forwarding table contains this information, all chunks received from the splitter will be forwarded to the rest of the team. So, in absence of communication constraints, the team will be organized as a full-connected overlay (see Fig. 2a).
The splitter, in an inﬁnite loop: (1) listens to the incoming peers, (2) sends to them the set of peers of the team, and (3) includes the incoming peer to the set. Notice that only those peers that are in the set of peers of the splitter are considered to be in the team served by such splitter.
Note: See in peer_dbs.py.
In order to hide the jitter generated by the physical network and the protocol itself, peers need to store the received chunks in a buﬀer during a period of time, before sending them to a player. A chunk with number is inserted in the position of the buﬀer, where is the maximum number of chunks that the buﬀer can store. In a peer’s life, is a constant especiﬁed by the user, but it is not compulsory that all peers of a team use the same buﬀer size. The larger the buﬀer size, the higher the buﬀering delay, but also the lower the probability of lossing chunks.
The buﬀer is implemented as a circular queue of chunks, which is ﬁlled up to only chunks during the buﬀering time (which is the main part of the start-up time that the users experiment). Chunks with a higher number (newer chunks) are inserted in the head of the buﬀer. The (received) chunk pointed by the tail of the buﬀer is sent to the player. This action is carried out each time a new chunk is received8 . Empty cells in the buﬀer (caused by the chunks that have not been received on time) are skipped until to ﬁnd the next cell with content.
determines how long the peers must wait for start playing the chunks. In general, should be as small as possible, and to achieve this we can reduce and . Unfortunately, these reductions generate another drawbacks. On the one hand, the overhead of the header of the transport protocol is inversely proportional to , and therefore, should be selected enough large to keep under control this overhead. On the other hand, if is too small (for example, if ) the peer will not have enought space to buﬀer all the chunks of a round, and due to the probability of receiving all the chunks in order is very small, some chunks will overwrite others before they can be played. This problem can also happen even if , because the maximum jitter for a given peer (generated by DBS) that a chunk can experiment is the sum of the maximum jitter produced by the splitter for this peer, that can be , and the maximum jitter produced by the team, , in the case of a full-connected mesh such as the shown in the Fig. 2a. Therefore, users should select .
Given a value, DBS peers may buﬀer a diﬀerent number of chunks that depends on the order in which chunks are received. If is the (number of the) ﬁrst received chunk (the ﬁrst chunk to be played), the buﬀering time ﬁnishes when a chunk with number equal or greater than is received.9 Lets analyze some interesting cases.
Lets suppose that the ﬁrst received chunk is and that the rest of chunks of the buﬀer of size are received, being the chunk the last one (this is the ideal scenario). In this case, the stream can be played without artifacts. Because the playing of the chunks starts after the buﬀering process, the end-latency experimented by users in the ideal case would be , being the latency generated by the physical transmission media.
Lets imagine now one of the worst possible scenarios, in which after receiving the chunk is received. In this case, the chunks have been lost (or delayed too much) by the physical transmission media or the transmission protocol, but again (and considering constant), the buﬀering time is because the chunk was generated chunk times after . Therefore, in this case the end-latency is also .
After considering these two extreme situations, we can conclude that the end-latency does not depend on the loss chunk ratio during the buﬀering time (always that this ratio is smaller than one), but only on , and .
DBS implements a push-based protocol. When a peer receives a chunk, it can be retransmitted to a large number of neighbors (depending on the number of diﬀerent destinations in its forwarding table). Therefore, even by controlling the chunk rate at the servers10 , some kind of ﬂow control must be performed in order to reduce network congestion while peers perform the ﬂooding.
The congestion (in particular, the one caused by how DBS nodes use the physical links) may be avoided by means of a basic idea: only if I have received a chunk, I send a chunk (not necessary to the sender neither the same chunk). It is easy to see that, in a fully connected overlay (Fig. 2a), this allows to control the data ﬂow. However, in more realistic scenarios (such as those in which ﬁrewalls and symmetric NATS are used), where the physical media imposes interconnexion constraints, peers can not be “connected” with the rest the team, and therefore, if the splitter follows a pure round-robin strategy, some peers can send more chunks that they receive (Fig. 2b). In these scenarios, the simple rule of sending a chunk for each received one does not work.
The previous idea can be adapted to handle a variable connectivity degree (also called neighborhood degree) if each peer uses a table of sets, , indexed by the neighbor’s end-points, where each set indicates the positions in the buﬀer of those chunks that must be transmited to the corresponding neighbor, the next time such neighbor be selected in the ﬂooding process. For example, if , chunks found at positions and of the buﬀer have to be sent to peer .
Notice that using this procedure, more than one chunk can be sent to a neighbor in a transmission burst, which could congest the switching devices. However, except in very unbalanced overlays (Fig. 2b), the bursts are very short on average (only one chunk in most of cases). As an advantage, if a burst is produced, all the chunks of the burst travel between the two same hosts, which usually increases the performance of the physical routing. In this case, chunks can be grouped in one packet, reducing the protocol overhead.
An example of the temporal evolution of a team using this behaviour has been described in the Figures 3, 4 and 5.
Chunks can be lost under bandwidth and buﬀering time constraints.11 A chunk is considered as lost when it is time to send it to the player and the chunk has not been received. In this situation, for each lost chunk, the peer sends a to a peer of the team, selected at random among the rest of the team. When a peer receives a from , adds to , where is the origin peer of the chunk stored in the position of its buﬀer.
In this situation, it can happen that some peers request redundant routes between an origin peer and itself, and therefore, some chunks could be received more than once. If this case, for each duplicate chunk, a peer should send a message to those neighbors that have sent to it the duplicate chunk (notice that the faster neighbor to send the chunk will not receive such prunning message). Neighbors receiving this message from peer should remove the from , where is the origin peer of the duplicate chunk.
As a consequence of these rules, the neighborhood degree of peers can decrease or increase. A decrement is produced if the is sent to a neighbor peer, and a duplicate is received from another neighbor. An increment is produced if the is sent to a peer that is not a neighbor, and a duplicate is received from a neighbor, that still keeps sending chunks at least from a diﬀerent origin.
As peers select (randomly) ’s destinations of the team, the list of known peers of the team should be up to date. This list is populated in ﬁrst instance when peers join a team. Peers also add to this list the sender of a or a chunk. Unfortunately, these rules do not guarantee that that all peers know the end-points of the rest of the team, because all the s and the chunks sent to a already incorporated peer can be lost. To solve this, peers also add to the list the new origin peers of the all received chunks, except for those received from the splitter.
An outgoing peer must to: (1) say to the splitter and the neighbor peers (in this order), (2) relay any pending (received but yet not sent) chunks, and (3) wait for a from the splitter. In case of timeout12 , the leaving procedure is reset a number of times.
When a peer of the team receives a , removes the sender from its table. The splitter removes the outgoing peer from the set of peers as soon as the is received.
The splitter remembers which chunk, of a list of the last transmitted chunks, was sent to each peer of the team. Notice that, in order to remember the chunk that was sent to each peer in each round, must be hold that . See in splitter_dbs.py.
Monitor peers (which are trusted peers) complain to their splitter with a for each lost chunk. The splitter only considers these type of messages if they come from a monitor.
Note: This last functionality has not been implemented, at least, as it has been explained here. The forget() thread is controlled by a timer, not by a counter of rounds.
DBS does not imposes any control over the grade of solidarity of the peers. This means that selﬁsh peers (or simply peers with reduced connectivity) can stay in the team thanks to the generosity of the rest of peers, even if they never achive to deliver a chunk to any peer of the team. This set or rules preclude this possible behavior, by impossing a minimum degree of solidarity between neighbor peers.
To know the level of solidarity between neighbor peers, each peer uses a table of chunk debts, . Every time a peer sends a chunk to , increments , and on the contrary, decrements when receives a chunk from .
Peers forward chunks to their neighbors in the order in which the entries of are accessed (a round-robing scheduling in DBS). Considering this, FCS modiﬁes this behavior:
Using FCS, supportive peers will be served ﬁrst, incrementing the QoE of the corresponding peers. On the other hand, those peers with a higher chunk debt will tend to be unserved if no enough bandwidth is available.
Note: The prioritized round-robin neighbor selection has not yet been implemented as it has been explained here. The structure exists, but is used for a diﬀerent purporse.
IPM is available by default in LANs (Local Are Network)s and VLANs (Virtual LANs) , but not in the Internet . IMS runs on the top of DBS and provides eﬃcient native IPM, where available.
All peers in the same LAN or VLAN have the same network address. When a joining peer receives the list of peers from its splitter, ﬁrst checks if there are neighbors in the same subnet. For all those peers, uses the IP address (all systems on this subnet), (default) port , to multicast (only) the chunks received from the splitter. Therefore, all peers in the same local network communicate using this multicast group address and port. The rest of external peers will be referenced using their public end-points.
In TAS, the splitter request to each peer of the team the list of neighbors (peers that send chunks directly, in one hop). This communication is reliable (TCP) and transmits the lists as a collection of end-points. The number of requests per round is limited by the available bandwidth in the overlay, and by the request-ratio deﬁned at the splitter. Obviously, the higher the ratio, a more accurate description of the real connectivity in the overlay will be obtained.
After knowing the connectivity degree of each peer, the slitter can adapt the round-robin scheduling of the origin peers by sending a number of chunks proportional to the inverse of the degree of the origin peer.
MRS extends DBS (or an extension of it) to retransmit massively-lost chunks. MRS should be implemented if error-prone communications are expected, specially if these channels are used by the splitter. MRS is based on the use of monitors (see Sec: 2.8). The idea is: the splitter will resend lost chunks to one or more the monitors when all monitors report their loss. To increase the probability of receiving on time the resent chunk (by normal peers), monitors halves the number of chunks in their buﬀers in relation to common peers. Notice that MRS only modiﬁes the behavior of the splitters and the monitors (normal peers does no need to implement LRS or its extensions).
ACS relaxes the peer’s upload requirements imposed by DBS. It should be used in if it is known that some peers can provide the capacity than others cannot, or when we want to mix the CS and P2P models, sending more chunks from the splitter to one or more monitors controlled by the contents provider.
ACS is based on the idea of using the information that the splitter knows about the number of chunks that each peer has lost (see Sec 2.8), to send to those more reliable peers a higher number of chunks than to the others. In other words, ACS adapts the round-time of each peer to its capacity.
Notice that ACS only aﬀects the behavior of the splitter.
Most of the peers run inside of “private” networks, i.e. behind NAT devices. NTS13 is an DBS extension which provides peer connectivity for some NAT conﬁgurations where DBS can not provide direct peer communication.14
Peers behind the same NAT will use the same external (also called “public”, because in most cases we have not nested NAT conﬁgurations) IP address of the NAT. Basically, there exist two diﬀerent types of NATs: (1) cone, and (2) symmetric. At the same time, NATs can implement diﬀerent ﬁltering strategies for the packets that comes from the external side: (a) no ﬁltering, (b) source IP ﬁltering, and (c) source end-point ﬁltering. Finally, NATs can use several port allocation algorithms, among which, the most frequent are: (i) port preservation and (ii) random port. Notice that in this discussion, only UDP transmissions will be considered.
Lets suppose a team in which, for the sake of simplicity, there is only one external (public) peer , and that a new internal (private) peer has sent the sequence of ’s (see Sec 2.3). Lets denote ’s NAT as . When no ﬁltering is used at all, forwards to any external packet that arrives to it (obviously, if it was sent to the entry in ’s translation table that was created during the transmission of the sequence of ’s), independently on the source end-points of the packets. In the case of source (IP) address ﬁltering, will forward the packets only if they come from ’s host. When source end-point ﬁltering is used, also checks the source port, i.e., that the packets were originated at ’s end-point.
Cone NATs use the same external end-point for every packet that comes from the same internal end-point, independently on the destination of the packets (see Fig. 9). For the external peer , the situation is identical to the case in which the NATed peer would be running in a public host.
Symmetric NATs use diﬀerent external end-points for diﬀerent packets that comes from the same internal end-point, when these packets have diﬀerent destination end-points (see Fig. ??). Thus, two diﬀerent external peers will see two diﬀerent public end-points of .
In the case of port preservation, if : is the private end-point (IP address:port) of a UDP packet, the NAT will use the public port , if available (notice that cound have been assigned to a previous communcation). If were unavailable, the NAT usually will assign the closer free port (this is called “sequentially port allocation”), usually by increasing the port value, although this behavior has not been standarized at all.
When random port allocation is implemented, the public port will be assigned at random. Notice that, even in SN-PPA conﬁgurations, in most of the real situations (where peers must compete with the rest of processes that use the network for the same NAT resources,) some kind of randomization should be always expected during a the port assignment.
An incoming peer can determine its NAT behavior using the following steps:
for , the port distances gathered by .
where GCD is the Greatest Common Divisor operator.
Table 2 shows the theoretical traversal success of DBS (or an extension of it) for diﬀerent NAT type combinations. Peer1 represents to a peer already joined to the team, and Peer2 to an incoming peer. The entries labeled with “DBS” are those that will be handled by DBS, out-of-the-box. An explanation of why the DBS handshake works for such conﬁgurations is shown in Fig. 10. Notice that source end-point ﬁltering has been used in this example, although a similar results should be obtained for simple source address ﬁltering. On the other hand, the combinations labeled with “-” or “NTS” will not work with DBS (see Fig.11). In fact, only the “NTS” entries should work, in general, with NTS, depending on the port prediction algorithm and the number of tries.
Fig. 12 shows an example of an NTS (NAT traversal) success. When the new NATed peers, and , arrive at the team, the following events happen:
Summarizing, NTS can provide connectivity for those peers that are behind port-preservation symmetric NATs with sequential port allocation.
When both peers, Peer1 and Peer2, are behind symmetric NATs, both must predict the port that the NAT of the interlocutor peer will use to send the packets towards it. And obviously, this must be performed by each already incorporated peer that is behind a symmetric NAT.
The list of predicted ports that a a peer performs is determined by:
where “” denotes the concatenation of lists and is the number of guessed ports, is the ﬁrst external port (the port used to communicate with ) assigned to the incoming peer and is the (maximum) port step measured for the incoming peer’s NAT.
When using MDC , SVC , or for emulating the CS model, it can be interesting for peers to belong to more than one team. To implement MCS, peers must replicate the P2PSP modules (DBS at least) for each team (channel), except the buﬀer.
The use of MDC is trivial: the higher the number received descriptions (channels), even partially, the higher the quality of the playback. However, when transmitting SVC media, peers should prioritize the reception of the most important layers.
When a peer belongs to more than one team, and the teams broadcast exactly the same stream (the same chunks and headers), it could move between teams seamless (without losts of signal).
A pure CS service could be provided if the corresponding splitter announces one empty team and sends each chunk so many times as teams (with one peer/team) there are.
A variety of techniques to ﬁght pollution in P2P live streaming systems are available in the literature, including hash-based signature and data encryption techniques.
 Pierpaolo Baccichet, Jeonghun Noh, Eric Setton, and Bernd Girod. Content-aware p2p video streaming with low latency. In Multimedia and Expo, 2007 IEEE International Conference on, pages 400–403. IEEE, 2007.
 Xiaowen Chu, Kaiyong Zhao, Zongpeng Li, and Anirban Mahanti. Auction-based on-demand p2p min-cost media streaming with network coding. IEEE Transactions on Parallel and Distributed Systems, 20(12):1816–1829, 2009.